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TCP Maintenance and Minor                                        F. Gont
Extensions (tcpm)                                                UK CPNI
Internet-Draft                                          January 21, 2011
Intended status: BCP
Expires: July 25, 2011


     Security Assessment of the Transmission Control Protocol (TCP)
                  draft-ietf-tcpm-tcp-security-02.txt

Abstract

   This document contains a security assessment of the specifications of
   the Transmission Control Protocol (TCP), and of a number of
   mechanisms and policies in use by popular TCP implementations.
   Additionally, it contains best current practices for hardening a TCP
   implementation.

Status of this Memo

   This Internet-Draft is submitted to IETF in full conformance with the
   provisions of BCP 78 and BCP 79.

   Internet-Drafts are working documents of the Internet Engineering
   Task Force (IETF).  Note that other groups may also distribute
   working documents as Internet-Drafts.  The list of current Internet-
   Drafts is at http://datatracker.ietf.org/drafts/current/.

   Internet-Drafts are draft documents valid for a maximum of six months
   and may be updated, replaced, or obsoleted by other documents at any
   time.  It is inappropriate to use Internet-Drafts as reference
   material or to cite them other than as "work in progress."

   This Internet-Draft will expire on July 25, 2011.

Copyright Notice

   Copyright (c) 2011 IETF Trust and the persons identified as the
   document authors.  All rights reserved.

   This document is subject to BCP 78 and the IETF Trust's Legal
   Provisions Relating to IETF Documents
   (http://trustee.ietf.org/license-info) in effect on the date of
   publication of this document.  Please review these documents
   carefully, as they describe your rights and restrictions with respect
   to this document.  Code Components extracted from this document must
   include Simplified BSD License text as described in Section 4.e of
   the Trust Legal Provisions and are provided without warranty as



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   described in the Simplified BSD License.


Table of Contents

   1.  Preface . . . . . . . . . . . . . . . . . . . . . . . . . . .   5
     1.1.  Introduction  . . . . . . . . . . . . . . . . . . . . . .   5
     1.2.  Scope of this document  . . . . . . . . . . . . . . . . .   6
     1.3.  Organization of this document . . . . . . . . . . . . . .   8
   2.  The Transmission Control Protocol . . . . . . . . . . . . . .   8
   3.  TCP header fields . . . . . . . . . . . . . . . . . . . . . .   9
     3.1.  Source Port and Destination Port  . . . . . . . . . . . .  10
     3.2.  Sequence number . . . . . . . . . . . . . . . . . . . . .  12
     3.3.  Acknowledgement Number  . . . . . . . . . . . . . . . . .  14
     3.4.  Data Offset . . . . . . . . . . . . . . . . . . . . . . .  15
     3.5.  Control bits  . . . . . . . . . . . . . . . . . . . . . .  15
       3.5.1.  Reserved (four bits)  . . . . . . . . . . . . . . . .  15
       3.5.2.  CWR (Congestion Window Reduced) . . . . . . . . . . .  16
       3.5.3.  ECE (ECN-Echo)  . . . . . . . . . . . . . . . . . . .  16
       3.5.4.  URG . . . . . . . . . . . . . . . . . . . . . . . . .  17
       3.5.5.  ACK . . . . . . . . . . . . . . . . . . . . . . . . .  17
       3.5.6.  PSH . . . . . . . . . . . . . . . . . . . . . . . . .  17
       3.5.7.  RST . . . . . . . . . . . . . . . . . . . . . . . . .  19
       3.5.8.  SYN . . . . . . . . . . . . . . . . . . . . . . . . .  19
       3.5.9.  FIN . . . . . . . . . . . . . . . . . . . . . . . . .  20
     3.6.  Window  . . . . . . . . . . . . . . . . . . . . . . . . .  20
     3.7.  Checksum  . . . . . . . . . . . . . . . . . . . . . . . .  22
     3.8.  Urgent pointer  . . . . . . . . . . . . . . . . . . . . .  23
     3.9.  Options . . . . . . . . . . . . . . . . . . . . . . . . .  24
     3.10. Padding . . . . . . . . . . . . . . . . . . . . . . . . .  28
     3.11. Data  . . . . . . . . . . . . . . . . . . . . . . . . . .  28
   4.  Common TCP Options  . . . . . . . . . . . . . . . . . . . . .  29
     4.1.  End of Option List (Kind = 0) . . . . . . . . . . . . . .  29
     4.2.  No Operation (Kind = 1) . . . . . . . . . . . . . . . . .  29
     4.3.  Maximum Segment Size (Kind = 2) . . . . . . . . . . . . .  29
     4.4.  Selective Acknowledgement Option  . . . . . . . . . . . .  32
       4.4.1.  SACK-permitted Option (Kind = 4)  . . . . . . . . . .  32
       4.4.2.  SACK Option (Kind = 5)  . . . . . . . . . . . . . . .  33
     4.5.  MD5 Option (Kind=19)  . . . . . . . . . . . . . . . . . .  35
     4.6.  Window scale option (Kind = 3)  . . . . . . . . . . . . .  36
     4.7.  Timestamps option (Kind = 8)  . . . . . . . . . . . . . .  37
       4.7.1.  Generation of timestamps  . . . . . . . . . . . . . .  37
       4.7.2.  Vulnerabilities . . . . . . . . . . . . . . . . . . .  38
   5.  Connection-establishment mechanism  . . . . . . . . . . . . .  39
     5.1.  SYN flood . . . . . . . . . . . . . . . . . . . . . . . .  40
     5.2.  Connection forgery  . . . . . . . . . . . . . . . . . . .  44
     5.3.  Connection-flooding attack  . . . . . . . . . . . . . . .  45
       5.3.1.  Vulnerability . . . . . . . . . . . . . . . . . . . .  45



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       5.3.2.  Countermeasures . . . . . . . . . . . . . . . . . . .  46
     5.4.  Firewall-bypassing techniques . . . . . . . . . . . . . .  48
   6.  Connection-termination mechanism  . . . . . . . . . . . . . .  49
     6.1.  FIN-WAIT-2 flooding attack  . . . . . . . . . . . . . . .  49
       6.1.1.  Vulnerability . . . . . . . . . . . . . . . . . . . .  49
       6.1.2.  Countermeasures . . . . . . . . . . . . . . . . . . .  50
   7.  Buffer management . . . . . . . . . . . . . . . . . . . . . .  52
     7.1.  TCP retransmission buffer . . . . . . . . . . . . . . . .  52
       7.1.1.  Vulnerability . . . . . . . . . . . . . . . . . . . .  52
       7.1.2.  Countermeasures . . . . . . . . . . . . . . . . . . .  53
     7.2.  TCP segment reassembly buffer . . . . . . . . . . . . . .  56
     7.3.  Automatic buffer tuning mechanisms  . . . . . . . . . . .  59
       7.3.1.  Automatic send-buffer tuning mechanisms . . . . . . .  59
       7.3.2.  Automatic receive-buffer tuning mechanism . . . . . .  61
   8.  TCP segment reassembly algorithm  . . . . . . . . . . . . . .  63
     8.1.  Problems that arise from ambiguity in the reassembly
           process . . . . . . . . . . . . . . . . . . . . . . . . .  63
   9.  TCP Congestion Control  . . . . . . . . . . . . . . . . . . .  64
     9.1.  Congestion control with misbehaving receivers . . . . . .  66
       9.1.1.  ACK division  . . . . . . . . . . . . . . . . . . . .  66
       9.1.2.  DupACK forgery  . . . . . . . . . . . . . . . . . . .  66
       9.1.3.  Optimistic ACKing . . . . . . . . . . . . . . . . . .  67
     9.2.  Blind DupACK triggering attacks against TCP . . . . . . .  68
       9.2.1.  Blind throughput-reduction attack . . . . . . . . . .  70
       9.2.2.  Blind flooding attack . . . . . . . . . . . . . . . .  70
       9.2.3.  Difficulty in performing the attacks  . . . . . . . .  71
       9.2.4.  Modifications to TCP's loss recovery algorithms . . .  72
       9.2.5.  Countermeasures . . . . . . . . . . . . . . . . . . .  74
     9.3.  TCP Explicit Congestion Notification (ECN)  . . . . . . .  79
       9.3.1.  Possible attacks by a compromised router  . . . . . .  79
       9.3.2.  Possible attacks by a malicious TCP endpoint  . . . .  80
   10. TCP API . . . . . . . . . . . . . . . . . . . . . . . . . . .  81
     10.1. Passive opens and binding sockets . . . . . . . . . . . .  81
     10.2. Active opens and binding sockets  . . . . . . . . . . . .  82
   11. Blind in-window attacks . . . . . . . . . . . . . . . . . . .  84
     11.1. Blind TCP-based connection-reset attacks  . . . . . . . .  84
       11.1.1. RST flag  . . . . . . . . . . . . . . . . . . . . . .  85
       11.1.2. SYN flag  . . . . . . . . . . . . . . . . . . . . . .  86
       11.1.3. Security/Compartment  . . . . . . . . . . . . . . . .  88
       11.1.4. Precedence  . . . . . . . . . . . . . . . . . . . . .  89
       11.1.5. Illegal options . . . . . . . . . . . . . . . . . . .  90
     11.2. Blind data-injection attacks  . . . . . . . . . . . . . .  90
   12. Information leaking . . . . . . . . . . . . . . . . . . . . .  91
     12.1. Remote Operating System detection via TCP/IP stack
           fingerprinting  . . . . . . . . . . . . . . . . . . . . .  91
       12.1.1. FIN probe . . . . . . . . . . . . . . . . . . . . . .  91
       12.1.2. Bogus flag test . . . . . . . . . . . . . . . . . . .  92
       12.1.3. TCP ISN sampling  . . . . . . . . . . . . . . . . . .  92



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       12.1.4. TCP initial window  . . . . . . . . . . . . . . . . .  92
       12.1.5. RST sampling  . . . . . . . . . . . . . . . . . . . .  93
       12.1.6. TCP options . . . . . . . . . . . . . . . . . . . . .  94
       12.1.7. Retransmission Timeout (RTO) sampling . . . . . . . .  94
     12.2. System uptime detection . . . . . . . . . . . . . . . . .  94
   13. Covert channels . . . . . . . . . . . . . . . . . . . . . . .  95
   14. TCP Port scanning . . . . . . . . . . . . . . . . . . . . . .  95
     14.1. Traditional connect() scan  . . . . . . . . . . . . . . .  96
     14.2. SYN scan  . . . . . . . . . . . . . . . . . . . . . . . .  96
     14.3. FIN, NULL, and XMAS scans . . . . . . . . . . . . . . . .  96
     14.4. Maimon scan . . . . . . . . . . . . . . . . . . . . . . .  98
     14.5. Window scan . . . . . . . . . . . . . . . . . . . . . . .  98
     14.6. ACK scan  . . . . . . . . . . . . . . . . . . . . . . . .  99
   15. Processing of ICMP error messages by TCP  . . . . . . . . . .  99
   16. TCP interaction with the Internet Protocol (IP) . . . . . . .  99
     16.1. TCP-based traceroute  . . . . . . . . . . . . . . . . . .  99
     16.2. Blind TCP data injection through fragmented IP traffic  . 100
     16.3. Broadcast and multicast IP addresses  . . . . . . . . . . 102
   17. Security Considerations . . . . . . . . . . . . . . . . . . . 102
   18. Acknowledgements  . . . . . . . . . . . . . . . . . . . . . . 102
   19. References  . . . . . . . . . . . . . . . . . . . . . . . . . 103
   20. References  . . . . . . . . . . . . . . . . . . . . . . . . . 113
     20.1. Normative References  . . . . . . . . . . . . . . . . . . 113
     20.2. Informative References  . . . . . . . . . . . . . . . . . 113
   Appendix A.  TODO list  . . . . . . . . . . . . . . . . . . . . . 113
   Appendix B.  Change log (to be removed by the RFC Editor
                before publication of this document as an RFC) . . . 113
     B.1.  Changes from draft-ietf-tcpm-tcp-security-01  . . . . . . 113
   Author's Address  . . . . . . . . . . . . . . . . . . . . . . . . 114






















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1.  Preface

1.1.  Introduction

   The TCP/IP protocol suite was conceived in an environment that was
   quite different from the hostile environment they currently operate
   in.  However, the effectiveness of the protocols led to their early
   adoption in production environments, to the point that, to some
   extent, the current world's economy depends on them.

   While many textbooks and articles have created the myth that the
   Internet protocols were designed for warfare environments, the top
   level goal for the DARPA Internet Program was the sharing of large
   service machines on the ARPANET [Clark, 1988].  As a result, many
   protocol specifications focus only on the operational aspects of the
   protocols they specify, and overlook their security implications.

   While the Internet technology evolved since it early inception, the
   Internet's building blocks are basically the same core protocols
   adopted by the ARPANET more than two decades ago.  During the last
   twenty years, many vulnerabilities have been identified in the TCP/IP
   stacks of a number of systems.  Some of them were based on flaws in
   some protocol implementations, affecting only a reduced number of
   systems, while others were based in flaws in the protocols
   themselves, affecting virtually every existing implementation
   [Bellovin, 1989].  Even in the last couple of years, researchers were
   still working on security problems in the core protocols [NISCC,
   2004] [NISCC, 2005].

   The discovery of vulnerabilities in the TCP/IP protocol suite usually
   led to reports being published by a number of CSIRTs (Computer
   Security Incident Response Teams) and vendors, which helped to raise
   awareness about the threats and the best mitigations known at the
   time the reports were published.  Unfortunately, this also led to the
   documentation of the discovered protocol vulnerabilities being spread
   among a large number of documents, which are sometimes difficult to
   identify.

   For some reason, much of the effort of the security community on the
   Internet protocols did not result in official documents (RFCs) being
   issued by the IETF (Internet Engineering Task Force).  This basically
   led to a situation in which "known" security problems have not always
   been addressed by all vendors.  In addition, in many cases vendors
   have implemented quick "fixes" to the identified vulnerabilities
   without a careful analysis of their effectiveness and their impact on
   interoperability [Silbersack, 2005].

   Producing a secure TCP/IP implementation nowadays is a very difficult



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   task, in part because of the lack of a single document that serves as
   a security roadmap for the protocols.  Implementers are faced with
   the hard task of identifying relevant documentation and
   differentiating between that which provides correct advice, and that
   which provides misleading advice based on inaccurate or wrong
   assumptions.

   There is a clear need for a companion document to the IETF
   specifications that discusses the security aspects and implications
   of the protocols, identifies the existing vulnerabilities, discusses
   the possible countermeasures, and analyzes their respective
   effectiveness.

   This document is the result of a security assessment of the IETF
   specifications of the Transmission Control Protocol (TCP), from a
   security point of view.  Possible threats are identified and, where
   possible, countermeasures are proposed.  Additionally, many
   implementation flaws that have led to security vulnerabilities have
   been referenced in the hope that future implementations will not
   incur the same problems.

   This document does not aim to be the final word on the security
   aspects of TCP.  On the contrary, it aims to raise awareness about a
   number of TCP vulnerabilities that have been faced in the past, those
   that are currently being faced, and some of those that we may still
   have to deal with in the future.

   Feedback from the community is more than encouraged to help this
   document be as accurate as possible and to keep it updated as new
   vulnerabilities are discovered.

   This document is heavily based on the "Security Assessment of the
   Transmission Control Protocol (TCP)" released by the UK Centre for
   the Protection of National Infrastructure (CPNI), available at: http:
   //www.cpni.gov.uk/Products/technicalnotes/
   Feb-09-security-assessment-TCP.aspx .

1.2.  Scope of this document

   While there are a number of protocols that may affect the way TCP
   operates, this document focuses only on the specifications of the
   Transmission Control Protocol (TCP) itself.

   The following IETF RFCs were selected for assessment as part of this
   work:






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   o  RFC 793, "Transmission Control Protocol.  DARPA Internet Program.
      Protocol Specification" (91 pages)

   o  RFC 1122, "Requirements for Internet Hosts -- Communication
      Layers" (116 pages)

   o  RFC 1191, "Path MTU Discovery" (19 pages)

   o  RFC 1323, "TCP Extensions for High Performance" (37 pages)

   o  RFC 1948, "Defending Against Sequence Number Attacks" (6 pages)

   o  RFC 1981, "Path MTU Discovery for IP version 6" (15 pages)

   o  RFC 2018, "TCP Selective Acknowledgment Options" (12 pages)

   o  RFC 2385, "Protection of BGP Sessions via the TCP MD5 Signature
      Option" (6 pages)

   o  RFC 2581, "TCP Congestion Control" (14 pages)

   o  RFC 2675, "IPv6 Jumbograms" (9 pages)

   o  RFC 2883, "An Extension to the Selective Acknowledgement (SACK)
      Option for TCP" (17 pages)

   o  RFC 2884, "Performance Evaluation of Explicit Congestion
      Notification (ECN) in IP Networks" (18 pages)

   o  RFC 2988, "Computing TCP's Retransmission Timer" (8 pages)

   o  RFC 3168, "The Addition of Explicit Congestion Notification (ECN)
      to IP" (63 pages)

   o  RFC 3465, "TCP Congestion Control with Appropriate Byte Counting
      (ABC)" (10 pages)

   o  RFC 3517, "A Conservative Selective Acknowledgment (SACK)-based
      Loss Recovery Algorithm for TCP" (13 pages)

   o  RFC 3540, "Robust Explicit Congestion Notification (ECN) Signaling
      with Nonces" (13 pages)

   o  RFC 3782, "The NewReno Modification to TCP's Fast Recovery
      Algorithm" (19 pages)






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1.3.  Organization of this document

   This document is basically organized in two parts.  The first part
   contains a discussion of each of the TCP header fields, identifies
   their security implications, and discusses the possible
   countermeasures.  The second part contains an analysis of the
   security implications of the mechanisms and policies implemented by
   TCP, and of a number of implementation strategies in use by a number
   of popular TCP implementations.


2.  The Transmission Control Protocol

   The Transmission Control Protocol (TCP) is a connection-oriented
   transport protocol that provides a reliable byte-stream data transfer
   service.

   Very few assumptions are made about the reliability of underlying
   data transfer services below the TCP layer.  Basically, TCP assumes
   it can obtain a simple, potentially unreliable datagram service from
   the lower level protocols.  Figure 1 illustrates where TCP fits in
   the DARPA reference model.

                             +---------------+
                             |  Application  |
                             +---------------+
                             |      TCP      |
                             +---------------+
                             |      IP       |
                             +---------------+
                             |    Network    |
                             +---------------+

                Figure 1: TCP in the DARPA reference model

   TCP provides facilities in the following areas:

   o  Basic Data Transfer

   o  Reliability

   o  Flow Control

   o  Multiplexing

   o  Connections





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   o  Precedence and Security

   o  Congestion Control

   The core TCP specification, RFC 793 [Postel, 1981c], dates back to
   1981 and standardizes the basic mechanisms and policies of TCP.  RFC
   1122 [Braden, 1989] provides clarifications and errata for the
   original specification.  RFC 2581 [Allman et al, 1999] specifies TCP
   congestion control and avoidance mechanisms, not present in the
   original specification.  Other documents specify extensions and
   improvements for TCP.

   The large amount of documents that specify extensions, improvements,
   or modifications to existing TCP mechanisms has led the IETF to
   publish a roadmap for TCP, RFC 4614 [Duke et al, 2006], that
   clarifies the relevance of each of those documents.


3.  TCP header fields

   RFC 793 [Postel, 1981c] defines the syntax of a TCP segment, along
   with the semantics of each of the header fields.  Figure 2
   illustrates the syntax of a TCP segment.

        0                   1                   2                   3
        0 1 2 3 4 5 6 7 8 9 0 1 2 3 4 5 6 7 8 9 0 1 2 3 4 5 6 7 8 9 0 1
       +-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
       |          Source Port          |       Destination Port        |
       +-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
       |                        Sequence Number                        |
       +-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
       |                    Acknowledgment Number                      |
       +-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
       |  Data |       |C|E|U|A|P|R|S|F|                               |
       | Offset|Resrved|W|C|R|C|S|S|Y|I|            Window             |
       |       |       |R|E|G|K|H|T|N|N|                               |
       +-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
       |           Checksum            |         Urgent Pointer        |
       +-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
       |                    Options                    |    Padding    |
       +-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
       |                             data                              |
       +-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+


   Note that one tick mark represents one bit position

           Figure 2: Transmission Control Protocol header format



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   The minimum TCP header size is 20 bytes, and corresponds to a TCP
   segment with no options and no data.  However, a TCP module might be
   handed an (illegitimate) "TCP segment" of less than 20 bytes.
   Therefore, before doing any processing of the TCP header fields, the
   following check should be performed by TCP on the segments handed by
   the internet layer:

                             Segment.Size >= 20

   If a segment does not pass this check, it should be dropped.

   The following subsections contain further sanity checks that should
   be performed on TCP segments.

3.1.  Source Port and Destination Port

   The Source Port field contains a 16-bit number that identifies the
   TCP end-point that originated this TCP segment.  The TCP Destination
   Port contains a 16-bit number that identifies the destination TCP
   end-point of this segment.  In most of the discussion we refer to
   client-side (or "ephemeral") port-numbers and server-side port
   numbers, since that distinction is what usually affects the
   interpretation of a port number.

   TCP SHOULD randomize its ephemeral (client-side) ports, to improve
   its resistance to off-path attacks.  For the purpose of ephemeral
   port selection, the largest posible port range SHOULD be used
   (ideally 1024-65535) I-D.ietf-tsvwg-port-randomization.

   DISCUSSION:

      [I-D.ietf-tsvwg-port-randomization] provides advice on port
      randomization.

   TCP MUST NOT allocate port number 0, as its use could lead to
   interoperability problems.  If a segment is received with port 0 as
   the Source Port or the Destination Port, a RST segment SHOULD be sent
   in response (provided that the incomming segment does not have the
   RST flag set).

   DISCUSSION:

      While port 0 is a legitimate port number, it has a special meaning
      in the UNIX Sockets API.  For example, when a TCP port number of 0
      is passed as an argument to the bind() function, rather than
      binding port 0, an ephemeral port is selected for the
      corresponding TCP end-point.  As a result, the TCP port number 0
      is never actually used in TCP segments.



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      Different implementations have been found to respond differently
      to TCP segments that have a port number of 0 as the Source Port
      and/or the Destination Port.  As a result, TCP segments with a
      port number of 0 are usually employed for remote OS detection via
      TCP/IP stack fingerprinting [Jones, 2003].

      Since in practice TCP port 0 is not used by any legitimate
      application and is only used for fingerprinting purposes, a number
      of host implementations already reject TCP segments that use 0 as
      the Source Port and/or the Destination Port.  Also, a number
      firewalls filter (by default) any TCP segments that contain a port
      number of zero for the Source Port and/or the Destination Port.

      We therefore recommend that TCP implementations respond to
      incoming TCP segments that have a Source Port or a Destination
      Port of 0 with an RST (provided these incoming segments do not
      have the RST bit set).

      Responding with an RST segment to incoming segments that have the
      RST bit would open the door to RST-war attacks.

   TCP MUST be able to grecefully handle the case where the source end-
   point (IP Source Address, TCP Source Port) is the same as the
   destination end-point (IP Destination Address, TCP Destination Port).

   DISCUSSION:

      Some systems have been found to be unable to process TCP segments
      in which the source endpoint {Source Address, Source Port} is the
      same than the destination end-point {Destination Address,
      Destination Port}.  Such TCP segments have been reported to cause
      malfunction of a number of implementations [CERT, 1996], and have
      been exploited in the past to perform Denial of Service (DoS)
      attacks [Meltman, 1997].  While these packets are very very
      unlikely to exist in real and legitimate scenarios, TCP should
      nevertheless be able to process them without the need of any
      "extra" code.

      A SYN segment in which the source end-point {Source Address,
      Source Port} is the same as the destination end-point {Destination
      Address, Destination Port} will result in a "simultaneous open"
      scenario, such as the one described in page 32 of RFC 793 [Postel,
      1981c].  Therefore, those TCP implementations that correctly
      handle simultaneous opens should already be prepared to handle
      these unusual TCP segments.

   TCP SHOULD NOT allocate of port numbers that are in use by a TCP that
   is in the LISTEN or CLOSED states for use as ephemeral ports, as this



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   could allow attackers on the local system to "steal" incomming TCP
   connections.

   DISCUSSION:

      While the only requirement for a selected ephemeral port is that
      the resulting four-tuple (connection-id) is unique (i.e., not
      currently in use by any other TCP connection), in practice it may
      be necessary to not allow the allocation of port numbers that are
      in use by a TCP that is in the LISTEN or CLOSED states for use as
      ephemeral ports, as this might allow an attacker to "steal"
      incoming connections from a local server application.  Therefore,
      TCP SHOULD NOT allocate port numbers that are in use by a TCP in
      the LISTEN or CLOSED states for use as ephemeral ports.  Section
      10.2 of this document provides a detailed discussion of this
      issue.

   While some systems restrict use of the port numbers in the range
   0-1024 to privileged users, applications SHOULD NOT grant any trust
   based on the port numbers used for a TCP connection.

   DISCUSSION:

      Not all systems require superuser privileges to bind port numbers
      in that range.  Besides, with desktop computers such "distinction"
      has generally become irrelevant.

   Middle-boxes such as packet filters MUST NOT assume that clients use
   port numbers from only the Dynamic or Registered port ranges.

   DISCUSSION:

      It should also be noted that some clients, such as DNS resolvers,
      are known to use port numbers from the "Well Known Ports" range.
      Therefore, middle-boxes such as packet filters MUST NOT assume
      that clients use port number from only the Dynamic or Registered
      port ranges.

3.2.  Sequence number

   TCP SHOULD select its Initial Sequence Numbers (ISNs) with the
   following expression:

   ISN = M + F(localhost, localport, remotehost, remoteport, secret_key)

   where M is a monotonically increasing counter maintained within TCP,
   and F() is a Pseudo-Random Function (PRF).  As it is vital that F()
   not be computable from the outside, F() could be a PRF of the



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   connection-id and some secret data.  HMAC-SHA-256 would be a good
   choice for F()

   DISCUSSION:

      The choice of the Initial Sequence Number of a connection is not
      arbitrary, but aims to minimize the chances of a stale segment
      from being accepted by a new incarnation of a previous connection.
      RFC 793 [Postel, 1981c] suggests the use of a global 32-bit ISN
      generator, whose lower bit is incremented roughly every 4
      microseconds.

      However, use of such an ISN generator makes it trivial to predict
      the ISN that a TCP will use for new connections, thus allowing a
      variety of attacks against TCP, such as those described in Section
      5.2 and Section 11 of this document.  This vulnerability was first
      described in [Morris, 1985], and its exploitation was widely
      publicized about 10 years later [Shimomura, 1995].

      As a matter of fact, protection against old stale segments from a
      previous incarnation of the connection comes from allowing the
      creation of a new incarnation of a previous connection only after
      2*MSL have passed since a segment corresponding to the old
      incarnation was last seen.  This is accomplished by the TIME-WAIT
      state, and TCP's "quiet time" concept.  However, as discussed in
      Section 3.1 and Section 11.1.2 of this document, the ISN can be
      used to perform some heuristics meant to avoid an interoperability
      problem that may arise when two systems establish connections at a
      high rate.  In order for such heuristics to work, the ISNs
      generated by a TCP should be monotonically increasing.

      The ISN generation scheme recommended in this section was
      originally proposed in RFC 1948 [Bellovin, 1996], such that the
      chances of an attacker from guessing the ISN of a TCP are reduced,
      while still producing a monotonically-increasing sequence that
      allows implementation of the optimization described in Section 3.1
      and Section 11.1.2 of this document.

      [CERT, 2001] and [US-CERT, 2001] are advisories about the security
      implications of weak ISN generators.  [Zalewski, 2001a] and
      [Zalewski, 2002] contain a detailed analysis of ISN generators,
      and a survey of the algorithms in use by popular TCP
      implementations.

      Another security consideration that should be made about TCP
      sequence numbers is that they might allow an attacker to count the
      number of systems behind a Network Address Translator (NAT)
      [Srisuresh and Egevang, 2001].  Depending on the ISN generators



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      implemented by each of the systems behind the NAT, an attacker
      might be able to count the number of systems behind the NAT by
      establishing a number of TCP connections (using the public address
      of the NAT) and indentifying the number of different sequence
      number "spaces".  This information leakage could be eliminated by
      rewriting the contents of all those header fields and options that
      make use of sequence numbers (such as the Sequence Number and the
      Acknowledgement Number fields, and the SACK Option) at the NAT.
      [Gont and Srisuresh, 2008] provides a detailed discussion of the
      security implications of NATs and of the possible mitigations for
      this and other issues.

3.3.  Acknowledgement Number

   TCP SHOULD set the Acknowledgement Number to zero when sending a TCP
   segment that does not have the ACK bit set (i.e., a SYN segment).

   TCP MUST check that, on segments that have the ACK bit set, the
   Acknowledgment Number satisfies the expression:

                SND.UNA - SND.MAX.WND <= SEG.ACK <= SND.NXT

   If a TCP segment does not pass this check, the segment MUST be
   dropped, and an ACK segment SHOULD be sent in response.

   DISCUSSION:

      If the ACK bit is on, the Acknowledgement Number contains the
      value of the next sequence number the sender of this segment is
      expecting to receive.  According to RFC 793, the Acknowledgement
      Number is considered valid as long as it does not acknowledge the
      receipt of data that has not yet been sent.

      However, as a result of recent concerns on forgery attacks against
      TCP (see Section 11 of this document), ongoing work at the IETF
      [Ramaiah et al, 2008] has proposed to enforce a more strict check
      on the Acknowledgement Number of segments that have the ACK bit
      set:

                SND.UNA - SND.MAX.WND <= SEG.ACK <= SND.NXT

      If the ACK bit is off, the Acknowledgement Number field is not
      valid.  We recommend TCP implementations to set the
      Acknowledgement Number to zero when sending a TCP segment that
      does not have the ACK bit set (i.e., a SYN segment).  Some TCP
      implementations have been known to fail to set the Acknowledgement
      Number to zero, thus leaking information.




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      TCP Acknowledgements are also used to perform heuristics for loss
      recovery and congestion control.  Section 9 of this document
      describes a number of ways in which these mechanisms can be
      exploited.

3.4.  Data Offset

   TCP MUST enforce the following checks on the Data Offset field:

                              Data Offset >= 5


                   Data Offset * 4 <= TCP segment length

   If a TCP segment does not pass these checks, it should be silently
   dropped.

      The TCP segment length should be obtained from the IP layer, as
      TCP does not include a TCP segment length field.

   DISCUSSION:

      The Data Offset field indicates the length of the TCP header in
      32-bit words.  As the minimum TCP header size is 20 bytes, the
      minimum legal value for this field is 5.

      For obvious reasons, the TCP header cannot be larger than the
      whole TCP segment it is part of.

3.5.  Control bits

   The following subsections provide a discussion of the different
   control bits in the TCP header.  TCP segments with unusual
   combinations of flags set have been known in the past to cause
   malfunction of some implementations, sometimes to the extent of
   causing them to crash [Postel, 1987] [Braden, 1992].  These packets
   are still usually employed for the purpose of TCP/IP stack
   fingerprinting.  Section 12.1 contains a discussion of TCP/IP stack
   fingerprinting.

3.5.1.  Reserved (four bits)

   TCP MUST ignore the Reserved field of incoming TCP segments.

   DISCUSSION:






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      These four bits are reserved for future use, and must be zero.  As
      with virtually every field, the Reserved field could be used as a
      covert channel.  While there exist intermediate devices such as
      protocol scrubbers that clear these bits, and firewalls that drop/
      reject segments with any of these bits set, these devices should
      consider the impact of these policies on TCP interoperability.
      For example, as TCP continues to evolve, all or part of the bits
      in the Reserved field could be used to implement some new
      functionality.  If some middle-box or end-system implementation
      were to drop a TCP segment merely because some of these bits are
      not set to zero, interoperability problems would arise.

3.5.2.  CWR (Congestion Window Reduced)

   DISCUSSION:

      The CWR flag, defined in RFC 3168 [Ramakrishnan et al, 2001], is
      used as part of the Explicit Congestion Notification (ECN)
      mechanism.  For connections in any of the synchronized states,
      this flag indicates, when set, that the TCP sending this segment
      has reduced its congestion window.

      An analysis of the security implications of ECN can be found in
      Section 9.3 of this document.

3.5.3.  ECE (ECN-Echo)

   DISCUSSION:

      The ECE flag, defined in RFC 3168 [Ramakrishnan et al, 2001], is
      used as part of the Explicit Congestion Notification (ECN)
      mechanism.

      Once a TCP connection has been established, an ACK segment with
      the ECE bit set indicates that congestion was encountered in the
      network on the path from the sender to the receiver.  This
      indication of congestion should be treated just as a congestion
      loss in non-ECN-capable TCP [Ramakrishnan et al, 2001].
      Additionally, TCP should not increase the congestion window (cwnd)
      in response to such an ACK segment that indicates congestion, and
      should also not react to congestion indications more than once
      every window of data (or once per round-trip time).

      An analysis of the security implications of ECN can be found in
      Section 9.3 of this document.






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3.5.4.  URG

   DISCUSSION:

      When the URG flag is set, the Urgent Pointer field contains the
      current value of the urgent pointer.

      Receipt of an "urgent" indication generates, in a number of
      implementations (such as those in UNIX-like systems), a software
      interrupt (signal) that is delivered to the corresponding process.

      In UNIX-like systems, receipt of an urgent indication causes a
      SIGURG signal to be delivered to the corresponding process.

      A number of applications handle TCP urgent indications by
      installing a signal handler for the corresponding signal (e.g.,
      SIGURG).  As discussed in [Zalewski, 2001b], some signal handlers
      can be maliciously exploited by an attacker, for example to gain
      remote access to a system.  While secure programming of signal
      handlers is out of the scope of this document, we nevertheless
      raise awareness that TCP urgent indications might be exploited to
      abuse poorly-written signal handlers.

      Section 3.9 discusses the security implications of the TCP urgent
      mechanism.

3.5.5.  ACK

   DISCUSSION:

      When the ACK bit is one, the Acknowledgment Number field contains
      the next sequence number expected, cumulatively acknowledging the
      receipt of all data up to the sequence number in the
      Acknowledgement Number, minus one.  Section 3.4 of this document
      describes sanity checks that should be performed on the
      Acknowledgement Number field.

      TCP Acknowledgements are also used to perform heuristics for loss
      recovery and congestion control.  Section 9 of this document
      describes a number of ways in which these mechanisms can be
      exploited.

3.5.6.  PSH

   As a result of a SEND call, TCP SHOULD send all queued data (provided
   that TCP's flow control and congestion control algorithms allow it).

   Received data SHOULD be immediately delivered to an application



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   calling the RECEIVE function, even if the data already available are
   less than those requested by the application.

   DISCUSSION:

      RFC 793 [Postel, 1981c] contains (in pages 54-64) a functional
      description of a TCP Application Programming Interface (API).  One
      of the parameters of the SEND function is the PUSH flag which,
      when set, signals the local TCP that it must send all unsent data.
      The TCP PSH (PUSH) flag will be set in the last outgoing segment,
      to signal the push function to the receiving TCP.  Upon receipt of
      a segment with the PSH flag set, the receiving user's buffer is
      returned to the user, without waiting for additional data to
      arrive.

      There are two security considerations arising from the PUSH
      function.  On the sending side, an attacker could cause a large
      amount of data to be queued for transmission without setting the
      PUSH flag in the SEND call.  This would prevent the local TCP from
      sending the queued data, causing system memory to be tied to those
      data for an unnecessarily long period of time.

      An analogous consideration should be made for the receiving TCP.
      TCP is allowed to buffer incoming data until the receiving user's
      buffer fills or a segment with the PSH bit set is received.  If
      the receiving TCP implements this policy, an attacker could send a
      large amount of data, slightly less than the receiving user's
      buffer size, to cause system memory to be tied to these data for
      an unnecessarily long period of time.  Both of these issues are
      discussed in Section 4.2.2.2 of RFC 1122 [Braden, 1989].

      In order to mitigate these potential vulnerabilities, we suggest
      assuming an implicit "PUSH" in every SEND call.  On the sending
      side, this means that as a result of a SEND call TCP should try to
      send all queued data (provided that TCP's flow control and
      congestion control algorithms allow it).  On the receiving side,
      this means that the received data will be immediately delivered to
      an application calling the RECEIVE function, even if the data
      already available are less than those requested by the
      application.

      It is interesting to note that popular TCP APIs (such as
      "sockets") do not provide a PUSH flag in any of the interfaces
      they define, but rather perform some kind of "heuristics" to set
      the PSH bit in outgoing segments.  As a result, the value of the
      PSH bit in the received TCP segments is usually a policy of the
      sending TCP, rather than a policy of the sending application.  All
      robust applications that make use of those APIs (such as the



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      sockets API) properly handle the case of a RECEIVE call returning
      less data (e.g., zero) than requested, usually by performing
      subsequent RECEIVE calls.

      Another potential malicious use of the PSH bit would be for an
      attacker to send small TCP segments (probably with zero bytes of
      data payload) to cause the receiving application to be
      unnecessarily woken up (increasing the CPU load), or to cause
      malfunction of poorly-written applications that may not handle
      well the case of RECEIVE calls returning less data than requested.

3.5.7.  RST

   TCP MUST process RST segments (i.e., segments with the RST bit set)
   as follows:

   o  If the Sequence Number of the RST segment is not valid (i.e.,
      falls outside of the receive window), silently drop the segment.

   o  If the Sequence Number of the RST segment matches the next
      expected sequence number (RCV.NXT), abort the corresponding
      connection.

   o  If the Sequence Number is valid (i.e., falls within the receive
      window) but is not exactly RCV.NXT, send an ACK segment (a
      "challenge ACK") of the form: <SEQ=SND.NXT><ACK=RCV.NXT><CTL=ACK>.
      TCP SHOULD rate-limit these challenge ACK segments.

   DISCUSSION:

      The RST bit is used to request the abortion (abnormal close) of a
      TCP connection.  RFC 793 [Postel, 1981c] suggests that an RST
      segment should be considered valid if its Sequence Number is valid
      (i.e., falls within the receive window).  However, in response to
      the security concerns raised by [Watson, 2004] and [NISCC, 2004],
      [Ramaiah et al, 2008] proposec the aforementioned stricter
      validity checks.

      Section 11.1 of this document describes TCP-based connection-reset
      attacks, along with a number of countermeasures to mitigate their
      impact.

3.5.8.  SYN

   DISCUSSION:






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      The SYN bit is used during the connection-establishment phase, to
      request the synchronization of sequence numbers.

      There are basically four different vulnerabilities that make use
      of the SYN bit: SYN-flooding attacks, connection forgery attacks,
      connection flooding attacks, and connection-reset attacks.  They
      are described in Section 5.1, Section 5.2, Section 5.3, and
      Section 11.1.2, respectively, along with the possible
      countermeasures.

3.5.9.  FIN

   DISCUSSION:

      The FIN flag is used to signal the remote end-point the end of the
      data transfer in this direction.  Receipt of a valid FIN segment
      (i.e., a TCP segment with the FIN flag set) causes the transition
      in the connection state, as part of what is usually referred to as
      the "connection termination phase".

      The connection-termination phase can be exploited to perform a
      number of resource-exhaustion attacks.  Section 6 of this document
      describes a number of attacks that exploit the connection-
      termination phase along with the possible countermeasures.

3.6.  Window

   DISCUSSION:

      The TCP Window field advertises how many bytes of data the remote
      peer is allowed to send before a new advertisement is made.
      Theoretically, the maximum transfer rate that can be achieved by
      TCP is limited to:

      Maximum Transfer Rate = Window / RTT

      This means that, under ideal network conditions (e.g., no packet
      loss), the TCP Window in use should be at least:

                       Window = 2 * Bandwidth * Delay

      Using a larger Window than that resulting from the previous
      equation will not provide any improvements in terms of
      performance.







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      In practice, selection of the most convenient Window size may also
      depend on a number of other parameters, such as: packet loss rate,
      loss recovery mechanisms in use, etc.

      Security implications of the maximum TCP window size

      An aspect of the TCP Window that is usually overlooked is the
      security implications of its size.  Increasing the TCP window
      increases the sequence number space that will be considered
      "valid" for incoming segments.  Thus, use of unnecessarily large
      TCP Window sizes increases TCP's vulnerability to forgery attacks
      unnecessarily.

      In those scenarios in which the network conditions are known
      and/or can be easily predicted, it is recommended that the TCP
      Window is never set to a value larger than that resulting from the
      equations above.  Additionally, the nature of the application
      running on top of TCP should be considered when tuning the TCP
      window.  As an example, an H.245 signaling application certainly
      does not have high requirements on throughput, and thus a window
      size of around 4 KBytes will usually fulfill its needs, while
      keeping TCP's resistance to off-path forgery attacks at a decent
      level.  Some rough measurements seem to indicate that a TCP window
      of 4Kbytes is common practice for TCP connections servicing
      applications such as BGP.

      In principle, a possible approach to avoid requiring
      administrators to manually set the TCP window would be to
      implement an automatic buffer tuning mechanism, such as that
      described in [Heffner, 2002].  However, as discussed in Section
      7.3.2 of this document these mechanisms can be exploited to
      perform other types of attacks.

      Security implications arising from closed windows

      The TCP window is a flow-control mechanism that prevents a fast
      data sender application from overwhelming a "slow" receiver.  When
      a TCP end-point is not willing to receive any more data (before
      some of the data that have already been received are consumed), it
      will advertise a TCP window of zero bytes.  This will effectively
      stop the sender from sending any new data to the TCP receiver.
      Transmission of new data will resume when the TCP receiver
      advertises a nonzero TCP window, usually with a TCP segment that
      contains no data ("an ACK").







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      This segment is usually referred to as a "window update", as the
      only purpose of this segment is to update the server regarding the
      new window.

      To accommodate those scenarios in which the ACK segment that
      "opens" the window is lost, TCP implements a "persist timer" that
      causes the TCP sender to query the TCP receiver periodically if
      the last segment received advertised a window of zero bytes.  This
      probe simply consists of sending one byte of new data that will
      force the TCP receiver to send an ACK segment back to the TCP
      sender, containing the current TCP window.  Similarly to the
      retransmission timeout timer, an exponential back-off is used when
      calculating the retransmission timer, so that the spacing between
      probes increases exponentially.

      A fundamental difference between the "persist timer" and the
      retransmission timer is that there is no limit on the amount of
      time during which a TCP can advertise a zero window.  This means
      that a TCP end-point could potentially advertise a zero window
      forever, thus keeping kernel memory at the TCP sender tied to the
      TCP retransmission buffer.  This could clearly be exploited as a
      vector for performing a Denial of Service (DoS) attack against
      TCP, such as that described in Section 7.1 of this document.

      Section 7.1 of this document describes a Denial of Service attack
      that aims at exhausting the kernel memory used for the TCP
      retransmission buffer, along with possible countermeasures.

3.7.  Checksum

   Middleboxes that process TCP segments MUST validate the Checksum
   field, and silently discard the TCP segment if such validation fails.

   DISCUSSION:

      The Checksum field is an error detection mechanism meant for the
      contents of the TCP segment and a number of important fields of
      the IP header.  It is computed over the full TCP header pre-pended
      with a pseudo header that includes the IP Source Address, the IP
      Destination Address, the Protocol number, and the TCP segment
      length.  While in principle there should not be security
      implications arising from this field, due to non-RFC-compliant
      implementations, the Checksum can be exploited to detect
      firewalls, evade network intrusion detection systems (NIDS),
      and/or perform Denial of Service attacks.






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      If a stateful firewall does not check the TCP Checksum in the
      segments it processes, an attacker can exploit this situation to
      perform a variety of attacks.  For example, he could send a flood
      of TCP segments with invalid checksums, which would nevertheless
      create state information at the firewall.  When each of these
      segments is received at its intended destination, the TCP checksum
      will be found to be incorrect, and the corresponding will be
      silently discarded.  As these segments will not elicit a response
      (e.g., an RST segment) from the intended recipients, the
      corresponding connection state entries at the firewall will not be
      removed.  Therefore, an attacker may end up tying all the state
      resources of the firewall to TCP connections that will never
      complete or be terminated, probably leading to a Denial of Service
      to legitimate users, or forcing the firewall to randomly drop
      connection state entries.

      If a NIDS does not check the Checksum of TCP segments, an attacker
      may send TCP segments with an invalid checksum to cause the NIDS
      to obtain a TCP data stream different from that obtained by the
      system being monitored.  In order to "confuse" the NIDS, the
      attacker would send TCP segments with an invalid Checksum and a
      Sequence Number that would overlap the sequence number space being
      used for his malicious activity.  FTester [Barisani, 2006] is a
      tool that can be used to assess NIDS on this issue.

      Finally, an attacker performing port-scanning could potentially
      exploit intermediate systems that do not check the TCP Checksum to
      detect whether a given TCP port is being filtered by an
      intermediate firewall, or the port is actually closed by the host
      being port-scanned.  If a given TCP port appeared to be closed,
      the attacker would then send a SYN segment with an invalid
      Checksum.  If this segment elicited a response (either an ICMP
      error message or a TCP RST segment) to this packet, then that
      response should come from a system that does not check the TCP
      checksum.  Since normal host implementations of the TCP protocol
      do check the TCP checksum, such a response would most likely come
      from a firewall or some other middle-box.

      [Ed3f, 2002] describes the exploitation of the TCP checksum for
      performing the above activities.  [US-CERT, 2005d] provides an
      example of a TCP implementation that failed to check the TCP
      checksum.

3.8.  Urgent pointer

                     Segment.Size - Data Offset * 4 > 0

   If a TCP segment with the URG bit set does not pass this check, it



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   MUST be silently dropped.

   For TCP segments that have the URG bit set to zero, sending TCP TCP
   SHOULD set the Urgent Pointer to zero.

   A receiving TCP MUST ignore the Urgent Pointer field of TCP segments
   for which the URG bit is zero.

   DISCUSSION:

      Section 3.7 of RFC 793 [Postel, 1981c] states (in page 42) that to
      send an urgent indication the user must also send at least one
      byte of data.

      If the URG bit is zero, the Urgent Pointer is not valid, and thus
      should not be processed by the receiving TCP.  Nevertheless, we
      recommend TCP implementations to set the Urgent Pointer to zero
      when sending a TCP segment that does not have the URG bit set, and
      to ignore the Urgent Pointer (as required by RFC 793) when the URG
      bit is zero.

      Some stacks have been known to fail to set the Urgent Pointer to
      zero when the URG bit is zero, thus leaking out the corresponding
      system memory contents.  [Zalewski, 2008] provides further details
      about this issue.

      Some implementations have been found to be unable to process TCP
      urgent indications correctly.  [Myst, 1997] originally described
      how TCP urgent indications could be exploited to perform a Denial
      of Service (DoS) attack against some TCP/IP implementations,
      usually leading to a system crash.

3.9.  Options

   [IANA, 2007] contains the official list of the assigned option
   numbers.  TCP Options have been specified in the past both within the
   IETF and by other groups.  [Hnes, 2007] contains an un-official
   updated version of the IANA list of assigned option numbers.  The
   following table contains a summary of the assigned TCP option
   numbers, which is based on [Hnes, 2007].











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   +--------+----------------------+-----------------------------------+
   |  Kind  |        Meaning       |              Summary              |
   +--------+----------------------+-----------------------------------+
   |    0   |  End of Option List  |      Discussed in Section 4.1     |
   +--------+----------------------+-----------------------------------+
   |    1   |     No-Operation     |      Discussed in Section 4.2     |
   +--------+----------------------+-----------------------------------+
   |    2   | Maximum Segment Size |      Discussed in Section 4.3     |
   +--------+----------------------+-----------------------------------+
   |    3   | WSOPT - Window Scale |      Discussed in Section 4.6     |
   +--------+----------------------+-----------------------------------+
   |    4   |    SACK Permitted    |     Discussed in Section 4.4.1    |
   +--------+----------------------+-----------------------------------+
   |    5   |         SACK         |     Discussed in Section 4.4.2    |
   +--------+----------------------+-----------------------------------+
   |    6   |  Echo (obsoleted by  |  Obsolete.  Specified in RFC 1072 |
   |        |       option 8)      |    [Jacobson and Braden, 1988]    |
   +--------+----------------------+-----------------------------------+
   |    7   |      Echo Reply      |  Obsolete.  Specified in RFC 1072 |
   |        | (obsoleted by option |    [Jacobson and Braden, 1988]    |
   |        |          8)          |                                   |
   +--------+----------------------+-----------------------------------+
   |    8   |  TSOPT - Time Stamp  |      Discussed in Section 4.7     |
   |        |        Option        |                                   |
   +--------+----------------------+-----------------------------------+
   |    9   |     Partial Order    |  Historic.  Specified in RFC 1693 |
   |        | Connection Permitted |       [Connolly et al, 1994]      |
   +--------+----------------------+-----------------------------------+
   |   10   |     Partial Order    |  Historic.  Specified in RFC 1693 |
   |        |    Service Profile   |       [Connolly et al, 1994]      |
   +--------+----------------------+-----------------------------------+
   |   11   |          CC          |  Historic.  Specified in RFC 1644 |
   |        |                      |           [Braden, 1994]          |
   +--------+----------------------+-----------------------------------+
   |   12   |        CC.NEW        |  Historic.  Specified in RFC 1644 |
   |        |                      |           [Braden, 1994]          |
   +--------+----------------------+-----------------------------------+
   |   13   |        CC.ECHO       |  Historic.  Specified in RFC 1644 |
   |        |                      |           [Braden, 1994]          |
   +--------+----------------------+-----------------------------------+
   |   14   |     TCP Alternate    |  Historic.  Specified in RFC 1146 |
   |        |   Checksum Request   |    [Zweig and Partridge, 1990]    |
   +--------+----------------------+-----------------------------------+
   |   15   |     TCP Alternate    |  Historic.  Specified in RFC 1145 |
   |        |     Checksum Data    |    [Zweig and Partridge, 1990]    |
   +--------+----------------------+-----------------------------------+
   |   16   |        Skeeter       |              Historic             |
   +--------+----------------------+-----------------------------------+



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   +--------+----------------------+-----------------------------------+
   |   17   |         Bubba        |              Historic             |
   +--------+----------------------+-----------------------------------+
   |   18   |   Trailer Checksum   |              Historic             |
   |        |        Option        |                                   |
   +--------+----------------------+-----------------------------------+
   |   19   | MD5 Signature Option |      Discussed in Section 4.5     |
   +--------+----------------------+-----------------------------------+
   |   20   |   SCPS Capabilities  |     Specified in [CCSDS, 2006]    |
   +--------+----------------------+-----------------------------------+
   |   21   |  Selective Negative  |     Specified in [CCSDS, 2006]    |
   |        |   Acknowledgements   |                                   |
   +--------+----------------------+-----------------------------------+
   |   22   |   Record Boundaries  |     Specified in [CCSDS, 2006]    |
   +--------+----------------------+-----------------------------------+
   |   23   |      Corruption      |     Specified in [CCSDS, 2006]    |
   |        |      experienced     |                                   |
   +--------+----------------------+-----------------------------------+
   |   24   |         SNAP         |              Historic             |
   +--------+----------------------+-----------------------------------+
   |   25   | Unassigned (released |             Unassigned            |
   |        |      2000-12-18)     |                                   |
   +--------+----------------------+-----------------------------------+
   |   26   |    TCP Compression   |              Historic             |
   |        |        Filter        |                                   |
   +--------+----------------------+-----------------------------------+
   |   27   | Quick-Start Response |  Specified in RFC 4782 [Floyd et  |
   |        |                      |             al, 2007]             |
   +--------+----------------------+-----------------------------------+
   | 28-252 |      Unassigned      |             Unassigned            |
   +--------+----------------------+-----------------------------------+
   |   253  |     RFC3692-style    |   Described by RFC 4727 [Fenner,  |
   |        |     Experiment 1     |               2006]               |
   +--------+----------------------+-----------------------------------+
   |   254  |     RFC3692-style    |   Described by RFC 4727 [Fenner,  |
   |        |     Experiment 2     |               2006]               |
   +--------+----------------------+-----------------------------------+

                           Table 1: TCP Options

   There are two cases for the format of a TCP option:

   o  Case 1: A single byte of option-kind.

   o  Case 2: An option-kind byte, followed by an option-length byte,
      and the actual option-data bytes.

   In options of the Case 2 above, the option-length byte counts the



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   option-kind byte and the option-length byte, as well as the actual
   option-data bytes.

   All options except "End of Option List" (Kind = 0) and "No Operation"
   (Kind = 1), are of "Case 2".

   For options that belong to the "Case 2" described above, the
   following checks MUST be performed:

                             option-length >= 2


              option-offset + option-length <= Data Offset * 4

   Where option-offset is the offset of the first byte of the option
   within the TCP header, with the first byte of the TCP header being
   assigned an offset of 0.

   If a TCP segment fails to pass any of these checks, it SHOULD be
   silently dropped.

   TCP MUST ignore unknown TCP options, provided they pass the
   validation checks specified above.  In the same way, middle-boxes
   such as packet filters SHOULD NOT reject TCP segments containing
   "unknown" TCP options that pass the validation checks described
   earlier in this Section.

   DISCUSSION:

      The value "2" in the first equation accounts for the option-kind
      byte and the option-length byte, and assumes zero bytes of option-
      data.  This check prevents, among other things, loops in option
      processing that may arise from incorrect option lengths.

      The second equation takes into account the limit on the legitimate
      option length imposed by the syntax of the TCP header, and is
      meant to detect forged option-length values that might make an
      option overlap with the TCP payload, or even go past the actual
      end of the TCP segment carrying the option.

   Middle-boxes such as packet filters should not reject TCP segments
   containing unknown options solely because these options have not been
   present in the SYN/SYN-ACK handshake.

   DISCUSSION:






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      There is renewed interest in defining new TCP options for purposes
      like improved connection management and maintenance, advanced
      congestion control schemes, and security features.  The evolution
      of the TCP/IP protocol suite would be severely impacted by
      obstacles to deploying such new protocol mechanisms.

   Middle-boxes such as packet filters SHOULD NOT reject TCP segments
   containing unknown options solely because these options have not been
   present in the SYN/SYN-ACK handshake.

   DISCUSSION:

      In the past, TCP enhancements based on TCP options regularly have
      specified the exchange of a specific "enabling" option during the
      initial SYN/SYN-ACK handshake.  Due to the severely limited TCP
      option space which has already become a concern, it should be
      expected that future specifications might introduce new options
      not negotiated or enabled in this way.  Therefore, middle-boxes
      such as packet filters should not reject TCP segments containing
      unknown options solely because these options have not been present
      in the SYN/SYN-ACK handshake.

   TCP MUST NOT "echo" in any way unknown TCP options received in
   inbound TCP segments.

   DISCUSSION:

      Some TCP implementations have been known to "echo" unknown TCP
      options received in incoming segments.  Here we stress that TCP
      must not "echo" in any way unknown TCP options received in inbound
      TCP segments.  This is at the foundation for the introduction of
      new TCP options, ensuring unambiguous behavior of systems not
      supporting a new specification.

   Section 4 discusses the security implications of common TCP options.

3.10.  Padding

   The TCP header padding is used to ensure that the TCP header ends and
   data begins on a 32-bit boundary.  The padding is composed of zeros.

3.11.  Data

   The data field contains the upper-layer packet being transmitted by
   means of TCP.  This payload is processed by the application process
   making use of the transport services of TCP.  Therefore, the security
   implications of this field are out of the scope of this document.




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4.  Common TCP Options

4.1.  End of Option List (Kind = 0)

   TCP implementations MUST be able to gracefully handle those TCP
   segments in which the End of Option List should have been present,
   but is missing.

   DISCUSSION:

      This option is used to indicate the "end of options" in those
      cases in which the end of options would not coincide with the end
      of the TCP header.

      TCP implementations are required to ignore those options they do
      not implement, and to be able to handle options with illegal
      lengths.  Therefore, TCP implementations should be able to
      gracefully handle those TCP segments in which the End of Option
      List should have been present, but is missing.

      It is interesting to note that some TCP implementations do not use
      the "End of Option List" option for indicating the "end of
      options", but simply pad the TCP header with several "No
      Operation" (Kind = 1) options to meet the header length specified
      by the Data Offset header field.

4.2.  No Operation (Kind = 1)

   The no-operation option is basically used to allow the sending system
   to align subsequent options in, for example, 32-bit boundaries.

   This option does not have any known security implications.

4.3.  Maximum Segment Size (Kind = 2)

   The Maximum Segment Size (MSS) option is used to indicate to the
   remote TCP endpoint the maximum segment size this TCP is willing to
   receive.

   The following check MUST be performed on a TCP segment that carries a
   MSS option:

                                  SYN == 1

   If the segment does not pass this check, it MUST be silently dropped.

   DISCUSSION:




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      As stated in Section 3.1 of RFC 793 [Postel, 1981c], this option
      can only be sent in the initial connection request (i.e., in
      segments with the SYN control bit set).

   TCP MUST check that the option length is 4.  If the option does not
   pass this check, it MUST be dropped.

   The received MSS SHOULD be sanitized as follows:

                       Sanitized_MSS = max(MSS, 536)

   This "sanitized" MSS value SHOULD be used to compute the "effective
   send MSS" by the expression included in Section 4.2.2.6 of RFC 1122
   [Braden, 1989], as follows:

   Eff.snd.MSS = min(Sanitized_MSS+20, MMS_S) - TCPhdrsize - IPoptionsize

   where:

   Sanitized_MSS:
      sanitized MSS value (the value received in the MSS option, with an
      enforced minimum value)

   MMS_S:
      maximum size for a transport-layer message that TCP may send

   TCPhdrsize:
      size of the TCP header, which typically was 20, but may be larger
      if TCP options are to be sent.

   IPoptionsize
      size of any IP options that TCP will pass to the IP layer with the
      current message.

   DISCUSSION:

      The advertised maximum segment size may be the result of the
      consideration of a number of factors.  Firstly, if fragmentation
      is employed, the size of the IP reassembly buffer may impose a
      limit on the maximum TCP segment size that can be received.
      Considering that the minimum IP reassembly buffer size is 576
      bytes, if an MSS option is not present included in the connection-
      establishment phase, an MSS of 536 bytes should be assumed.
      Secondly, if Path-MTU Discovery (specified in RFC 1191 [Mogul and
      Deering, 1990] and RFC 1981 [McCann et al, 1996]) is expected to
      be used for the connection, an artificial maximum segment size may
      be enforced by a TCP to prevent the remote peer from sending TCP
      segments which would be too large to be transmitted without



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      fragmentation.  Finally, a system connected by a low-speed link
      may choose to introduce an artificial maximum segment size to
      enforce an upper limit on the network latency that would otherwise
      negatively affect its interactive applications [Stevens, 1994].

      The TCP specifications do not impose any requirements on the
      maximum segment size value that is included in the MSS option.
      However, there are a number of values that may cause undesirable
      results.  Firstly, an MSS of 0 could possible "freeze" the TCP
      connection, as it would not allow data to be included in the
      payload of the TCP segments.  Secondly, low values other than 0
      would degrade the performance of the TCP connection (wasting more
      bandwidth in protocol headers than in actual data), and could
      potentially exhaust processing cycles at the sending TCP and/or
      the receiving TCP by producing an increase in the interrupt rate
      caused by the transmitted (or received) packets.

      The problems that might arise from low MSS values were first
      described by [Reed, 2001].  However, the community did not reach
      consensus on how to deal with these issues at that point.

      RFC 791 [Postel, 1981a] requires IP implementations to be able to
      receive IP datagrams of at least 576 bytes.  Assuming an IPv4
      header of 20 bytes, and a TCP header of 20 bytes, there should be
      room in each IP packet for 536 application data bytes.

      There are two cases to analyze when considering the possible
      interoperability impact of sanitizing the received MSS value: TCP
      connections relying on IP fragmentation and TCP connections
      implementing Path-MTU Discovery.  In case the corresponding TCP
      connection relies on IP fragmentation, given that the minimum
      reassembly buffer size is required to be 576 bytes by RFC 791
      [Postel, 1981a], the adoption of 536 bytes as a lower limit is
      safe.

      In case the TCP connection relies on Path-MTU Discovery, imposing
      a lower limit on the adopted MSS may ignore the advice of the
      remote TCP on the maximum segment size that can possibly be
      transmitted without fragmentation.  As a result, this could lead
      to the first TCP data segment to be larger than the Path-MTU.
      However, in such a scenario, the TCP segment should elicit an ICMP
      Unreachable "fragmentation needed and DF bit set" error message
      that would cause the "effective send MSS" (E_MSS) to be decreased
      appropriately.  Thus, imposing a lower limit on the accepted MSS
      will not cause any interoperability problems.






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      A possible scenario exists in which the proposed enforcement of a
      lower limit in the received MSS might lead to an interoperability
      problem.  If a system was attached to the network by means of a
      link with an MTU of less than 576 bytes, and there was some
      intermediate system which either silently dropped (i.e., without
      sending an ICMP error message) those packets equal to or larger
      than that 576 bytes, or some intermediate system simply filtered
      ICMP "fragmentation needed and DF bit set" error messages, the
      proposed behavior would not lead to an interoperability problem,
      when communication could have otherwise succeeded.  However, the
      interoperability problem would really be introduced by the network
      setup (e.g., the middle-box silently dropping packets), rather
      than by the mechanism proposed in this section.  In any case, TCP
      should nevertheless implement a mechanism such as that specified
      by RFC 4821 [Mathis and Heffner, 2007] to deal with this type of
      "network black-holes".

4.4.  Selective Acknowledgement Option

   The Selective Acknowledgement option provides an extension to allow
   the acknowledgement of individual segments, to enhance TCP's loss
   recovery.

   Two options are involved in the SACK mechanism.  The "Sack-permitted
   option" is sent during the connections-establishment phase, to
   advertise that SACK is supported.  If both TCP peers agree to use
   selective acknowledgements, the actual selective acknowledgements are
   sent, if needed, by means of "SACK options".

4.4.1.  SACK-permitted Option (Kind = 4)

   The SACK-permitted option is meant to advertise that the TCP sending
   this segment supports Selective Acknowledgements.

   The following check MUST be performed on a TCP segment that carries a
   MSS option:

                                  SYN == 1

   If a segment does not pass this check, it MUST be silently dropped.

   DISCUSSION:

      The SACK-permitted option can be sent only in SYN segments.

   TCP MUST check that the option length is 2.  If the option does not
   pass this check it MUST be silently dropped.




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4.4.2.  SACK Option (Kind = 5)

   The SACK option is used to convey extended acknowledgment information
   from the receiver to the sender over an established TCP connection.
   The option consists of an option-kind byte (which must be 5), an
   option-length byte, and a variable number of SACK blocks.

   TCP MUST silently discard those TCP segments carrying a SACK option
   that does not pass the following check:

              option-offset + option-length <= Data Offset * 4

   TCP MUST silently discard those TCP segments carrying a SACK option
   that does not pass the following check:

                            option-length >= 10

   DISCUSSION:

      A SACK Option with zero SACK blocks is nonsensical.  The value
      "10" accounts for the option-kind byte, the option-length byte, a
      4-byte left-edge field, and a 4-byte right-edge field.

   TCP MUST silently discard those TCP segments carrying a SACK option
   that does not pass the following check:

                        (option-length - 2) % 8 == 0

   DISCUSSION:

      As stated in Section 3 of RFC 2018 [Mathis et al, 1996], a SACK
      option that specifies n blocks will have a length of 8*n+2.

   TCP MUST silently discard those TCP segments carrying a SACK option
   that contains a SACK block that does not pass the following check:

                  Left Edge of Block < Right Edge of Block

   As in all the other occurrences in this document, all comparisons
   between sequence numbers should be performed using sequence number
   arithmetic.

   DISCUSSION:








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      Each block included in a SACK option represents a number of
      received data bytes that are contiguous and isolated; that is, the
      bytes just below the block, (Left Edge of Block - 1), and just
      above the block, (Right Edge of Block), have not yet been
      received.

   TCP MUST enforce a limit on the number of SACK blocks that a TCP will
   store in memory for each connection at any time.

   DISCUSSION:

      The TCP receiving a SACK option is expected to keep track of the
      selectively-acknowledged blocks.  Even when space in the TCP
      header is limited (and thus each TCP segment can selectively-
      acknowledge at most four blocks of data), an attacker could try to
      perform a buffer overflow or a resource-exhaustion attack by
      sending a large number of SACK options.

      For example, an attacker could send a large number of SACK
      options, each of them acknowledging one byte of data.
      Additionally, for the purpose of wasting resources on the attacked
      system, each of these blocks would be separated from each other by
      one byte, to prevent the attacked system from coalescing two (or
      more) contiguous SACK blocks into a single SACK block.  If the
      attacked system kept track of each SACKed block by storing both
      the Left Edge and the Right Edge of the block, then for each
      window of data, the attacker could waste up to 4 * Window bytes of
      memory at the attacked TCP.

      The value "4 * Window" results from the expression "(Window / 2) *
      8", in which the value "2" accounts for the 1-byte block
      selectively-acknowledged by each SACK block and 1 byte that would
      be used to separate each SACK blocks from each other, and the
      value "8" accounts for the 8 bytes needed to store the Left Edge
      and the Right Edge of each SACKed block.

      Therefore, it is clear that a limit should be imposed on the
      number of SACK blocks that a TCP will store in memory for each
      connection at any time.  Measurements in [Dharmapurikar and
      Paxson, 2005] indicate that in the vast majority of cases
      connections have a single hole in the data stream at any given
      time.  Thus, a limit of 16 SACK blocks for each connection would
      handle even most of the more unusual cases in which there is more
      than one simultaneous hole at a time.







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4.5.  MD5 Option (Kind=19)

   The TCP MD5 option provides a mechanism for authenticating TCP
   segments with a 18-byte digest produced by the MD5 algorithm.  The
   option consists of an option-kind byte (which must be 19), an option-
   length byte (which must be 18), and a 16-byte MD5 digest.

   TCP MUST silently drop a TCP segment that carries a TCP MD5 option
   that does not pass the following checks:

              option-offset + option-length <= Data Offset * 4


                            option-length == 18

   DISCUSSION:

      The TCP MD5 option is of "Case 2", and has a fixed length.

   DISCUSSION:

      A basic weakness on the TCP MD5 option is that the MD5 algorithm
      itself has been known (for a long time) to be vulnerable to
      collision search attacks.

      [Bellovin, 2006] argues that it has two other weaknesses, namely
      that it does not provide a key identifier, and that it has no
      provision for automated key management.  However, it is generally
      accepted that while a Key-ID field can be a good approach for
      providing smooth key rollover, it is not actually a requirement.
      For instance, most systems implementing the TCP MD5 option include
      a "keychain" mechanism that fully supports smooth key rollover.
      Additionally, with some further work, ISAKMP/IKE could be used to
      configure the MD5 keys.

      It is interesting to note that while the TCP MD5 option, as
      specified by RFC 2385 [Heffernan, 1998], addresses the TCP-based
      forgery attacks against TCP discussed in Section 11, it does not
      address the ICMP-based connection-reset attacks discussed in
      Section 15.  As a result, while a TCP connection may be protected
      from TCP-based forgery attacks by means of the MD5 option, an
      attacker might still be able to successfully perform the ICMP-
      based counter-part.

      The TCP MD5 option has been obsoleted by the TCP-AO.






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4.6.  Window scale option (Kind = 3)

   The window scale option provides a mechanism to expand the definition
   of the TCP window to 32 bits, such that the performance of TCP can be
   improved in some network scenarios.  The Window scale option consists
   of an option-kind byte (which must be 3), followed by an option-
   length byte (which must be 3), and a shift count (shift.cnt) byte
   (the actual option-data).

   The option may be sent only in the initial SYN segment, but may also
   be sent in a SYN/ACK segment if the option was received in the
   initial SYN segment.  If the option is received in any other segment,
   it MUST be silently dropped.

   TCP MUST silently discard TCP segments that contain a Window scale
   option whose option-length is not 3.

   DISCUSSION:

      This option has a fixed length.

   TCP MUST silently discard TCP segments that contain a Window scale
   option that does not pass the following check:

                              shift.cnt <= 14

   DISCUSSION:

      As discussed in Section 2.3 of RFC 1323 [Jacobson et al, 1992], in
      order to prevent new data from being mistakenly considered as old
      and vice versa, the resulting window should be equal to or smaller
      than 2^32.

   DISCUSSION:

      [Welzl, 2008] describes major problems with the use of the Window
      scale option in the Internet due to faulty equipment.

      While there are not known security implications arising from the
      window scale mechanism itself, the size of the TCP window has a
      number of security implications.  In general, larger window sizes
      increase the chances of an attacker from successfully performing
      forgery attacks against TCP, such as those described in Section 11
      of this document.  Additionally, large windows can exacerbate the
      impact of resource exhaustion attacks such as those described in
      Section 7 of this document.





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      Section 3.7 provides a general discussion of the security
      implications of the TCP window size.  Section 7.3.2 discusses the
      security implications of Automatic receive-buffer tuning
      mechanisms.

4.7.  Timestamps option (Kind = 8)

   The Timestamps option, specified in RFC 1323 [Jacobson et al, 1992],
   is used to perform two functions: Round-Trip Time Measurement (RTTM),
   and Protection Against Wrapped Sequence Numbers (PAWS).

   TCP MUST silently discard TCP segments that contain a Timestamps
   option that does not pass the following check:

                            option-length == 10

   DISCUSSION:

      As specified by RFC 1323, the option-length must be 10.

4.7.1.  Generation of timestamps

   TCP SHOULD generate timestamps with the following expression:

   timestamp = T() + F(localhost, localport, remotehost, remoteport, secret_key)

   where the result of T() is a global system clock that complies with
   the requirements of Section 4.2.2 of RFC 1323 [Jacobson et al, 1992],
   and F() is a function that should not be computable from the outside.
   Therefore, we suggest F() to be a cryptographic hash function of the
   connection-id and some secret data.

   DISCUSSION:

      For the purpose of PAWS, the timestamps sent on a connection are
      required to be monotonically increasing.  While there is no
      requirement that timestamps are monotonically increasing across
      TCP connections, the generation of timestamps such that they are
      monotonically increasing across connections between the same two
      endpoints allows the use of timestamps for improving the handling
      of SYN segments that are received while the corresponding four-
      tuple is in the TIME-WAIT state.  This is discussed in Section
      11.1.2 of this document.

      F() provides an offset that will be the same for all incarnations
      of a connection between the same two endpoints, while T() provides
      the monotonically increasing values that are needed for PAWS.




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      Further discussion about this algorithm is available in
      [I-D.gont-timestamps-generation].

   TCP SHOULD NOT initialize a global timestamp counter to a fixed value
   when the system is bootstrapped.

   DISCUSSION:

      Some implementations are known to initialize their global
      timestamp clock to zero when the system is bootstrapped.  This is
      undesirable, as the timestamp clock would disclose the system
      uptime.

   TCP SHOULD set the Timestamp Echo Reply (TSecr) field to zero when
   sending a TCP segment that does not have the ACK bit set (i.e., a SYN
   segment).

   DISCUSSION:

      Some TCP implementations have been found to fail to set the
      Timestamp Echo Reply field (TSecr) to zero in TCP segments that do
      not have the ACK bit set, thus potentially leaking information.

4.7.2.  Vulnerabilities

   Blind In-Window Attacks

   Segments that contain a timestamp option smaller than the last
   timestamp option recorded by TCP are silently dropped.  This allows
   for a subtle attack against TCP that would allow an attacker to cause
   one direction of data transfer of the attacked connection to freeze
   [US-CERT, 2005c].  An attacker could forge a TCP segment that
   contains a timestamp that is much larger than the last timestamp
   recorded for that direction of the data transfer of the connection.
   The offending segment would cause the recorded timestamp (TS.Recent)
   to be updated and, as a result, subsequent segments sent by the
   impersonated TCP peer would be simply dropped by the receiving TCP.
   This vulnerability has been documented in [US-CERT, 2005d].  However,
   it is worth noting that exploitation of this vulnerability requires
   an attacker to guess (or know) the four-tuple {IP Source Address, IP
   Destination Address, TCP Source Port, TCP Destination Port}, as well
   a valid Sequence Number and a valid Acknowledgement Number.  If an
   attacker has such detailed knowledge about a TCP connection, unless
   TCP segments are protected by proper authentication mechanisms (such
   as IPsec [Kent and Seo, 2005]), he can perform a variety of attacks
   against the TCP connection, even more devastating than the one just
   described.




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   Information leaking

   Some implementations are known to maintain a global timestamp clock,
   which is used for all connections.  This is undesirable, as an
   attacker that can establish a connection with a host would learn the
   timestamp used for all the other connections maintained by that host,
   which could be useful for performing any attacks that require the
   attacker to forge TCP segments.  A timestamps generator such as the
   one recommended in Section 4.7.1 of this document would prevent this
   information leakage, as it separates the "timestamps space" among the
   different TCP connections.

   Some implementations are known to initialize their global timestamp
   clock to zero when the system is bootstrapped.  This is undesirable,
   as the timestamp clock would disclose the system uptime.  A
   timestamps generator such as the one recommended in Section 4.7.1 of
   this document would prevent this information leakage, as the function
   F() introduces an "offset" that does not disclose the system uptime.

   As discussed in Section 3.2 of RFC 1323 [Jacobson et al, 1992], the
   Timestamp Echo Reply field (TSecr) is only valid if the ACK bit of
   the TCP header is set, and its value must be zero when it is not
   valid.  However, some TCP implementations have been found to fail to
   set the Timestamp Echo Reply field (TSecr) to zero in TCP segments
   that do not have the ACK bit set, thus potentially leaking
   information.  We stress that TCP implementations should comply with
   RFC 1323 by setting the Timestamp Echo Reply field (TSecr) to zero in
   those TCP segments that do not have the ACK bit set, thus eliminating
   this potential information leakage.

   Finally, it should be noted that the Timestamps option can be
   exploited to count the number of systems behind NATs (Network Address
   Translators) [Srisuresh and Egevang, 2001].  An attacker could count
   the number of systems behind a NAT by establishing a number of TCP
   connections (using the public address of the NAT) and indentifying
   the number of different timestamp sequences.  This information
   leakage could be eliminated by rewriting the contents of the
   Timestamps option at the NAT.  [Gont and Srisuresh, 2008] provides a
   detailed discussion of the security implications of NATs, and
   proposes mitigations for this and other issues.


5.  Connection-establishment mechanism

   The following subsections describe a number of attacks that can be
   performed against TCP by exploiting its connection-establishment
   mechanism.




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5.1.  SYN flood

   TCP SHOULD implement (and enable by default) a syn-cache [Lemon,
   2002].

   TCP SHOULD implement syn-cookies, and SHOULD enable them only after a
   specified number of TCBs has been allocated for connections in the
   SYN-RECEIVED state.

   DISCUSSION:

      TCP uses a mechanism known as the "three-way handshake" for the
      establishment of a connection between two TCP peers.  RFC 793
      [Postel, 1981c] states that when a TCP that is in the LISTEN state
      receives a SYN segment (i.e., a TCP segment with the SYN flag
      set), it must transition to the SYN-RECEIVED state, record the
      control information (e.g., the ISN) contained in the SYN segment
      in a Transmission Control Block (TCB), and respond with a SYN/ACK
      segment.

      A Transmission Control Block is the data structure used to store
      (usually within the kernel) all the information relevant to a TCP
      connection.  The concept of "TCB" is introduced in the core TCP
      specification RFC 793 [Postel, 1981c].

      In practice, virtually all existing implementations do not modify
      the state of the TCP that was in the LISTEN state, but rather
      create a new TCP (i.e., a new "protocol machine"), and perform all
      the state transitions on this newly-created TCP.  This allows the
      application running on top of TCP to service to more than one
      client at the same time.  As a result, each connection request
      results in the allocation of system memory to store the TCB
      associated with the newly created TCB.

      If TCP was implemented strictly as described in RFC 793, the
      application running on top of TCP would have to finish servicing
      the current client before being able to service the next one in
      line, or should instead be able to perform some kind of connection
      hand-off.

      An attacker could exploit TCP's connection-establishment mechanism
      to perform a Denial of Service (DoS) attack, by sending a large
      number of connection requests to the target system, with the
      intent of exhausting the system memory destined for storing TCBs
      (or related kernel data structures), thus preventing the attacked
      system from establishing new connections with legitimate users.
      This attack is widely known as "SYN flood", and has received a lot
      of attention during the late 90's [CERT, 1996].



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      Given that the attacker does not need to complete the three-way
      handshake for the attacked system to tie system resources to the
      newly created TCBs, he will typically forge the source IP address
      of the malicious SYN segments he sends, thus concealing his own IP
      address.

      If the forged IP addresses corresponded to some reachable system,
      the impersonated system would receive the SYN/ACK segment sent by
      the attacked host (in response to the forged SYN segment), which
      would elicit an RST segment.  This RST segment would be delivered
      to the attacked system, causing the corresponding connection to be
      aborted, and the corresponding TCB to be removed.

      As the impersonated host would not have any state information for
      the TCP connection being referred to by the SYN/ACK segment, it
      would respond with a RST segment, as specified by the TCP segment
      processing rules of RFC 793 [Postel, 1981c].

      However, if the forged IP source addresses were unreachable, the
      attacked TCP would continue retransmitting the SYN/ACK segment
      corresponding to each connection request, until timing out and
      aborting the connection.  For this reason, a number of widely
      available attack tools first check whether each of the (forged) IP
      addresses are reachable by sending an ICMP echo request to them.
      The receipt of an ICMP echo response is considered an indication
      of the IP address being reachable (and thus results in the
      corresponding IP address not being used for performing the
      attack), while the receipt of an ICMP unreachable error message is
      considered an indication of the IP address being unreachable (and
      thus results in the corresponding IP address being used for
      performing the attack).

      [Gont, 2008b] describes how the so-called ICMP soft errors could
      be used by TCP to abort connections in any of the non-synchronized
      states.  While implementation of the mechanism described in that
      document would certainly not eliminate the vulnerability of TCP to
      SYN flood attacks (as the attacker could use addresses that are
      simply "black-holed"), it provides an example of how signaling
      information such as that provided by means of ICMP error messages
      can provide valuable information that a transport protocol could
      use to perform heuristics.

      In order to mitigate the impact of this attack, the amount of
      information stored for non-established connections should be
      reduced (ideally, non-synchronized connections should not require
      any state information to be maintained at the TCP performing the
      passive OPEN).  There are basically two mitigation techniques for
      this vulnerability: a syn-cache and syn-cookies.



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      [Borman, 1997] and RFC 4987 [Eddy, 2007] contain a general
      discussion of SYN-flooding attacks and common mitigation
      approaches.

      The syn-cache [Lemon, 2002] approach aims at reducing the amount
      of state information that is maintained for connections in the
      SYN-RECEIVED state, and allocates a full TCB only after the
      connection has transited to the ESTABLISHED state.

      The syn-cookie [Bernstein, 1996] approach aims at completely
      eliminating the need to maintain state information at the TCP
      performing the passive OPEN, by encoding the most elementary
      information required to complete the three-way handshake in the
      Sequence Number of the SYN/ACK segment that is sent in response to
      the received SYN segment.  Thus, TCP is relieved from keeping
      state for connections in the SYN-RECEIVED state.

      The syn-cookie approach has a number of drawbacks:

      *  Firstly, given the limited space in the Sequence Number field,
         it is not possible to encode all the information included in
         the initial segment, such as, for example, support of Selective
         Acknowledgements (SACK).

      *  Secondly, in the event that the Acknowledgement segment sent in
         response to the SYN/ACK sent by the TCP that performed the
         passive OPEN (i.e., the TCP server) were lost, the connection
         would end up in the ESTABLISHED state on the client-side, but
         in the CLOSED state on the server side.  This scenario is
         normally handled in TCP by having the TCP server retransmit its
         SYN/ACK.  However, if syn-cookies are enabled, there would be
         no connection state information on the server side, and thus
         the SYN/ACK would never be retransmitted.  This could lead to a
         scenario in which the connection could remain in the
         ESTABLISHED state on the client side, but in the CLOSED state
         at the server side, indefinitely.  If the application protocol
         was such that it required the client to wait for some data from
         the server (e.g., a greeting message) before sending any data
         to the server, a deadlock would take place, with the client
         application waiting for such server data, and the server
         waiting for the TCP three-way handshake to complete.

      *  Thirdly, unless the function used to encode information in the
         SYN/ACK packet is cryptographically strong, an attacker could
         forge TCP connections in the ESTABLISHED state by forging ACK
         segments that would be considered as "legitimate" by the
         receiving TCP.




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      *  Fourthly, in those scenarios in which establishment of new
         connections is blocked by simply dropping segments with the SYN
         bit set, use of SYN cookies could allow an attacker to bypass
         the firewall rules, as a connection could be established by
         forging an ACK segment with the correct values, without the
         need of setting the SYN bit.

      As a result, syn-cookies are usually not employed as a first line
      of defense against SYN-flood attacks, but are only as the last
      resort to cope with them.  For example, some TCP implementations
      enable syn-cookies only after a certain number of TCBs has been
      allocated for connections in the SYN-RECEIVED state.  We recommend
      this implementation technique, with a syn-cache enabled by
      default, and use of syn-cookies triggered, for example, when the
      limit of TCBs for non-synchronized connections with a given port
      number has been reached.

      It is interesting to note that a SYN-flood attack should only
      affect the establishment of new connections.  A number of books
      and online documents seem to assume that TCP will not be able to
      respond to any TCP segment that is meant for a TCP port that is
      being SYN-flooded (e.g., respond with an RST segment upon receipt
      of a TCP segment that refers to a non-existent TCP connection).
      While SYN-flooding attacks have been successfully exploited in the
      past for achieving such a goal [Shimomura, 1995], as clarified by
      RFC 1948 [Bellovin, 1996] the effectiveness of SYN flood attacks
      to silence a TCP implementation arose as a result of a bug in the
      4.4BSD TCP implementation [Wright and Stevens, 1994], rather than
      from a theoretical property of SYN-flood attacks themselves.
      Therefore, those TCP implementations that do not suffer from such
      a bug should not be silenced as a result of a SYN-flood attack.

      [Zquete, 2002] describes a mechanism that could theoretically
      improve the functionality of SYN cookies.  It exploits the TCP
      "simultaneous open" mechanism, as illustrated in Figure 5.

             See Figure 5, in page 46 of the UK CPNI document.

             Use of TCP simultaneous open for handling SYN floods

      In line 1, TCP A initiates the connection-establishment phase by
      sending a SYN segment to TCP B. In line 2, TCP B creates a SYN
      cookie as described by [Bernstein, 1996], but does not set the ACK
      bit of the segment it sends (thus really sending a SYN segment,
      rather than a SYN/ACK).  This "fools" TCP A into thinking that
      both SYN segments "have crossed each other in the network" as if a
      "simultaneous open" scenario had taken place.  As a result, in
      line 3 TCP A sends a SYN/ACK segment containing the same options



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      that were contained in the original SYN segment.  In line 4, upon
      receipt of this segment, TCP processes the cookie encoded in the
      ACK field as if it had been the result of a traditional SYN cookie
      scenario, and moves the connection into the ESTABLISHED state.  In
      line 5, TCP B sends a SYN/ACK segment, which causes the connection
      at TCP A to move into the ESTABLISHED state.  In line 6, TCP A
      sends a data segment on the connection.

      While this mechanism would work in theory, unfortunately there are
      a number of factors that prevent it from being usable in real
      network environments:

      *  Some systems are not able to perform the "simultaneous open"
         operation specified in RFC 793, and thus the connection
         establishment will fail.

      *  Some firewalls might prevent the establishment of TCP
         connections that rely on the "simultaneous open" mechanism
         (e.g., a given firewall might be allowing incoming SYN/ACK
         segments, but not outgoing SYN/ACK segments).

      Therefore, we do not recommend implementation of this mechanism
      for mitigating SYN-flood attacks.

5.2.  Connection forgery

   The process of causing a TCP connection to be illegitimately
   established between two arbitrary remote peers is usually referred to
   as "connection spoofing" or "connection forgery".  This can have a
   great negative impact when systems establish some sort of trust
   relationships based on the IP addresses used to establish a TCP
   connection [daemon9 et al, 1996].

   It should be stressed that hosts should not establish trust
   relationships based on the IP addresses [CPNI, 2008] or on the TCP
   ports in use for the TCP connection (see Section 3.1 and Section 3.2
   of this document).

   One of the underlying weaknesses that allow this vulnerability to be
   more easily exploited is the use of an inadequate Initial Sequence
   Number (ISN) generator, as explained back in the 80's in [Morris,
   1985].  As discussed in Section 3.3.1 of this document, any TCP
   implementation that makes use of an inadequate ISN generator will be
   more vulnerable to this type of attack.  A discussion of approaches
   for a more careful generation of Initial Sequence Numbers (ISNs) can
   be found in Section 3.3.1 of this document.

   Another attack vector for performing connection-forgery attacks is



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   the use of IP source routing.  By forging the Source Address of the
   IP packets that encapsulate the TCP segments of a connection, and
   carefully crafting an IP source route option (i.e., either LSSR or
   SSRR) that includes a system whose traffic he can monitor, an
   attacker could cause the packets sent by the attacked system (e.g.,
   the SYN/ACK segment sent in response to the attacker's SYN segment)
   to be illegitimately directed to him [CPNI, 2008].  Thus, the
   attacker would not even need to guess valid sequence numbers for
   forging a TCP connection, as he would simply have direct access to
   all this information.  As discussed in [CPNI, 2008], it is strongly
   recommended that systems disable IP Source Routing by default, or at
   the very least, they disable source routing for IP packets that
   encapsulate TCP segments.

   The IPv6 Routing Header Type 0, which provides a similar
   functionality to that provided by IPv4 source routing, has been
   officially deprecated by RFC 5095 [Abley et al, 2007].

5.3.  Connection-flooding attack

5.3.1.  Vulnerability

   The creation and maintenance of a TCP connection requires system
   memory to maintain shared state between the local and the remote TCP.
   As system memory is a finite resource, there is a limit on the number
   of TCP connections that a system can maintain at any time.  When the
   TCP API is employed to create a TCP connection with a remote peer, it
   allocates system memory for maintaining shared state with the remote
   TCP peer, and thus the resulting connection would tie a similar
   amount of resources at the remote host as at the local host.
   However, if special packet-crafting tools are employed to forge TCP
   segments to establish TCP connections with a remote peer, the local
   kernel implementation of TCP can be bypassed, and the allocation of
   resources on the attacker's system for maintaining shared state can
   be avoided.  Thus, a malicious user could create a large number of
   TCP connections, and subsequently abandon them, thus tying system
   resources only at the remote peer.  This allows an attacker to create
   a large number of TCP connections at the attacked system with the
   intent of exhausting its kernel memory, without exhausting the
   attacker's own resources.  [CERT, 2000] discusses this vulnerability,
   which is usually referred to as the "Naptha attack".

   This attack is similar in nature to the "Netkill" attack discussed in
   Section 7.1.1.  However, while Netkill ties both TCBs and TCP send
   buffers to the abandoned connections, Naptha only ties TCBs (and
   related kernel structures), as it doesn't issue any application
   requests.




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   The symptom of this attack is an extremely large number of TCP
   connections in the ESTABLISHED state, which would tend to exhaust
   system resources and deny service to new clients (or possibly cause
   the system to crash).

   It should be noted that it is possible for an attacker to perform the
   same type of attack causing the abandoned connections to remain in
   states other than ESTABLISHED.  This might be interesting for an
   attacker, as it is usually the case that connections in states other
   than ESTABLISHED usually have no controlling user-space process (that
   is, the former controlling process for the connection has already
   closed the corresponding file descriptor).

   A particularly interesting case of a connection-flooding attack that
   aims at abandoning connections in a state other than ESTABLISHED is
   discussed in Section 6.1 of this document.

5.3.2.  Countermeasures

   As with many other resource exhaustion attacks, the problem in
   generating countermeasures for this attack is that it may be
   difficult to differentiate between an actual attack and a legitimate
   high-load scenario.  However, there are a number of countermeasures
   which, when tuned for each particular network environment, could
   allow a system to resist this attack and continue servicing
   legitimate clients.

   Hosts SHOULD enforce limits on the number of TCP connections with no
   user-space controlling process.

   DISCUSSION:

      Connections in states other than ESTABLISHED usually have no user-
      space controlling process.  This prevents the application making
      use of those connections from enforcing limits on the maximum
      number of ongoing connections (either on a global basis or a
      per-IP address basis).  When resource exhaustion is imminent or
      some threshold of ongoing connections is reached, the operating
      system should consider freeing system resources by aborting
      connections that have no user-space controlling process.  A number
      of such connections could be aborted on a random basis, or based
      on some heuristics performed by the operating system (e.g., first
      abort connections with peers that have the largest number of
      ongoing connections with no user-space controlling process).

   Hosts SHOULD enforce per-process and per-user limits on maximum
   kernel memory that can be used at any time.




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   Hosts SHOULD enforce per-process and per-user limits on the number of
   existent TCP connections at any time.

   DISCUSSION:

      While the Naphta attack is usually targeted at a service such as
      HTTP, its impact is usually system-wide.  This is particularly
      undesirable, as an attack against a single service might affect
      the system as a whole (for example, possibly precluding remote
      system administration).

      In order to avoid an attack to a single service from affecting
      other services, we advise TCP implementations to enforce per-
      process and per-user limits on maximum kernel memory that can be
      used at any time.  Additionally, we recommend implementations to
      enforce per-process and per-user limits on the number of existent
      TCP connections at any time.

   Applications SHOULD enforce limits on the number of simultaneous
   connections that can be established from a single IP address or
   network prefix at any given time.

   DISCUSSION:

      An application could limit the number of simultaneous connections
      that can be established from a single IP address or network prefix
      at any given time.  Once that limit has been reached, some other
      connection from the same IP address or network prefix would be
      aborted, thus allowing the application to service this new
      incoming connection.

      There are a number of factors that should be taken into account
      when defining the specific limit to enforce.  For example, in the
      case of protocols that have an authentication phase (e.g., SSH,
      POP3, etc.), this limit could be applied to sessions that have not
      yet been authenticated.  Additionally, depending on the nature and
      use of the application, it might or might not be normal for a
      single system to have multiple connections to the same server at
      the same time.

      For many network services, the limit of maximum simultaneous
      connections could be kept very low.  For example, an SMTP server
      could limit the number of simultaneous connections from a single
      IP address to 10 or 20 connections.

      While this limit could work in many network scenarios, we
      recommend network operators to measure the maximum number of
      concurrent connections from a single IP address during normal



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      operation, and set the limit accordingly.

      In the case of web servers, this limit will usually need to be set
      much higher, as it is common practice for web clients to establish
      multiple simultaneous connections with a single web server to
      speed up the process of loading a web page (e.g., multiple graphic
      files can be downloaded simultaneously using separate TCP
      connections).

      NATs (Network Address Translators) [Srisuresh and Egevang, 2001]
      are widely deployed in the Internet, and may exacerbate this
      situation, as a large number of clients behind a NAT might each
      establish multiple TCP connections with a given web server, which
      would all appear to be originate from the same IP address (that of
      the NAT box).

   Firewalls MAY enforce limits on the number of simultaneous
   connections that can be established from a single IP address or
   network prefix at any given time.

   DISCUSSION:

      Some firewalls can be configured to limit the number of
      simultaneous connections that any system can maintain with a
      specific system and/or service at any given time.  Limiting the
      number of simultaneous connections that each system can establish
      with a specific system and service would effectively limit the
      possibility of an attacker that controls a single IP address to
      exhaust system resources at the attacker system/service.

5.4.  Firewall-bypassing techniques

   TCP MUST silently drop those TCP segments that have both the SYN and
   the RST flags set.

   DISCUSSION:

      Some firewalls block incoming TCP connections by blocking only
      incoming SYN segments.  However, there are inconsistencies in how
      different TCP implementations handle SYN segments that have
      additional flags set, which may allow an attacker to bypass
      firewall rules [US-CERT, 2003b].

      For example, some firewalls have been known to mistakenly allow
      incoming SYN segments if they also have the RST bit set.  As some
      TCP implementations will create a new connection in response to a
      TCP segment with both the SYN and RST bits set, an attacker could
      bypass the firewall rules and establish a connection with a



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      "protected" system by setting the RST bit in his SYN segments.

      Here we advise TCP implementations to silently drop those TCP
      segments that have both the SYN and the RST flags set.


6.  Connection-termination mechanism

6.1.  FIN-WAIT-2 flooding attack

6.1.1.  Vulnerability

   TCP implements a connection-termination mechanism that is employed
   for the graceful termination of a TCP connection.  This mechanism
   usually consists of the exchange of four-segments.  Figure 6
   illustrates the usual segment exchange for this mechanism.

   Figure 6: TCP connection-termination mechanism

             See Figure 6, in page 50 of the UK CPNI document.

                   TCP connection-termination mechanism

   A potential problem may arise as a result of the FIN-WAIT-2 state:
   there is no limit on the amount of time that a TCP can remain in the
   FIN-WAIT-2 state.  Furthermore, no segment exchange is required to
   maintain the connection in that state.

   As a result, an attacker could establish a large number of
   connections with the target system, and cause it close each of them.
   For each connection, once the target system has sent its FIN segment,
   the attacker would acknowledge the receipt of this segment, but would
   send no further segments on that connection.  As a result, an
   attacker could cause the corresponding system resources (e.g., the
   system memory used for storing the TCB) without the need to send any
   further packets.

   While the CLOSE command described in RFC 793 [Postel, 1981c] simply
   signals the remote TCP end-point that this TCP has finished sending
   data (i.e., it closes only one direction of the data transfer), the
   close() system-call available in most operating systems has different
   semantics: it marks the corresponding file descriptor as closed (and
   thus it is no longer usable), and assigns the operating system the
   responsibility to deliver any queued data to the remote TCP peer and
   to terminate the TCP connection.  This makes the FIN-WAIT-2 state
   particularly attractive for performing memory exhaustion attacks, as
   even if the application running on top of TCP were imposing limits on
   the maximum number of ongoing connections, and/or time limits on the



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   function calls performed on TCP connections, that application would
   be unable to enforce these limits on the FIN-WAIT-2 state.

6.1.2.  Countermeasures

   A number of countermeasures can be implemented to mitigate FIN-WAIT-2
   flooding attacks.  Some of these countermeasures require changes in
   the TCP implementations, while others require changes in the
   applications running on top of TCP.

   TCP SHOULD enforce limits on the duration of the FIN-WAIT-2 state.

   DISCUSSION:

      In order to avoid the risk of having connections stuck in the FIN-
      WAIT-2 state indefinitely, a number of systems incorporate a
      timeout for the FIN-WAIT-2 state.  For example, the Linux kernel
      version 2.4 enforces a timeout of 60 seconds [Linux, 2008].  If
      the connection-termination mechanism does not complete before that
      timeout value, it is aborted.

   Enabling applications to enforce limits on ongoing connections

   As discussed in Section 6.1.1, the fact that the close() system call
   marks the corresponding file descriptor as closed prevents the
   application running on top of TCP from enforcing limits on the
   corresponding connection.

   While it is common practice for applications to terminate their
   connections by means of the close() system call, it is possible for
   an application to initiate the connection-termination phase without
   closing the corresponding file descriptor (hence keeping control of
   the connection).

   In order to achieve this, an application performing an active close
   (i.e., initiating the connection-termination phase) should replace
   the system-call close(sockfd) with the following code sequence:

   o  A call to shutdown(sockfd, SHUT_WR), to close the sending
      direction of this connection

   o  Successive calls to read(), until it returns 0, thus indicating
      that the remote TCP peer has finished sending data.

   o  A call to setsockopt(sockfd, SOL_SOCKET, SO_LINGER, &l,
      sizeof(l)), where l is of type struct linger (with its members
      l.l_onoff=1 and l.l_linger=90).




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   o  A call to close(sockfd), to close the corresponding file
      descriptor.

   The call to shutdown() (instead of close()) allows the application to
   retain control of the underlying TCP connection while the connection
   transitions through the FIN-WAIT-1 and FIN-WAIT-2 states.  However,
   the application will not retain control of the connection while it
   transitions through the CLOSING and TIME-WAIT states.

   It should be noted that, strictly speaking, close(sockfd) decrements
   the reference count for the descriptor sockfd, and initiates the
   connection termination phase only when the reference count reaches 0.
   On the other hand, shutdown(sockfd, SHUT_WR) initiates the
   connection-termination phase, regardless of the reference count for
   the sockfd descriptor.  This should be taken into account when
   performing the code replacement described above.  For example, it
   would be a bug for two processes (e.g., parent and child) that share
   a descriptor to both call shutdown(sockfd, SHUT_WR).

   An application performing a passive close should replace the call to
   close(sockfd) with the following code sequence:

   o  A call to setsockopt(sockfd, SOL_SOCKET, SO_LINGER, &l,
      sizeof(l)), where l is of type struct linger (with its members
      l.l_onoff=1 and l.l_linger=90).

   o  A call to close(sockfd), to close the corresponding file
      descriptor.

   It is assumed that if the application is performing a passive close,
   the application already detected that the remote TCP peer finished
   sending data by means as a result of a call to read() returning 0.

   In this scenario, the application will not retain control of the
   underlying connection when it transitions through the LAST_ACK state.

   Enforcing limits on the number of connections with no user-space
   controlling process

   The considerations and recommendations in Section 5.3.2 for enforcing
   limits on the number of connections with no user-space controlling
   process are applicable to mitigate this vulnerability.

   Limiting the number of simultaneous connections at the application

   The considerations and recommendations in Section 5.3.2 for limiting
   the number of simultaneous connections at the application are to
   mitigate this vulnerability.  We note, however, that unless



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   applications are implemented to retain control of the underlying TCP
   connection while the connection transitions through the FIN-WAIT-1
   and FIN-WAIT-2 states, enforcing such limits may prove to be a
   difficult task.

   Limiting the number of simultaneous connections at firewalls

   The considerations and recommendations in Section 5.3.2 for enforcing
   limiting the number of simultaneous connections at firewalls are
   applicable to mitigate this vulnerability.


7.  Buffer management

7.1.  TCP retransmission buffer

7.1.1.  Vulnerability

   [Shalunov, 2000] describes a resource exhaustion attack (Netkill)
   that can be performed against TCP.  The attack aims at exhausting
   system memory by creating a large number of TCP connections which are
   then abandoned.  The attack is usually performed as follows:

   o  The attacker creates a TCP connection to a service in which a
      small client request can result in a large server response (e.g.,
      HTTP).  Rather than relying on his kernel implementation of TCP,
      the attacker creates his TCP connections by means of a specialized
      packet-crafting tool.  This allows the attacker to create the TCP
      connections and later abandon them, exhausting the resources at
      the attacked system, while not tying his own system resources to
      the abandoned connections.

   o  When the connection is established (i.e., the three-way handshake
      has completed), an application request is sent, and the TCP
      connection is subsequently abandoned.  At this point, any state
      information kept by the attack tool is removed.

   o  The attacked server allocates TCP send buffers for transmitting
      the response to the client's request.  This causes the victim TCP
      to tie resources not only for the Transmission Control Block
      (TCB), but also for the application data that needs to be
      transferred.

   o  Once the application response is queued for transmission, the
      application closes the TCP connection, and thus TCP takes the
      responsibility to deliver the queued data.  Having the application
      close the connection has the benefit for the attacker that the
      application is not able to keep track of the number of TCP



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      connections in use, and thus it is not able to enforce limits on
      the number of connections.

   o  The attacker repeats the above steps a large number of times, thus
      causing a large amount of system memory at the victim host to be
      tied to the abandoned connections.  When the system memory is
      exhausted, the victim host denies service to new connections, or
      possibly crashes.

   There are a number of factors that affect the effectiveness of this
   attack that are worth considering.  Firstly, while the attack is
   typically targeted at a service such as HTTP, the consequences of the
   attack are usually system-wide.  Secondly, depending on the size of
   the server's response, the underlying TCP connection may or may not
   be closed: if the response is larger than the TCP send buffer size at
   the server, the application will usually block in a call to write()
   or send(), and would therefore not close the TCP connection, thus
   allowing the application to enforce limits on the number of ongoing
   connections.  Consequently, the attacker will usually try to elicit a
   response that is equal to or slightly smaller than the send buffer of
   the attacked TCP.  Thirdly, while [Shalunov, 2000] notes that one
   visible effect of this attack is a large number of connections in the
   FIN-WAIT-1 state, this will not usually be the case.  Given that the
   attacker never acknowledges any segment other than the SYN/ACK
   segment that is part of the three-way handshake, at the point in
   which the attacked TCP tries to send the application's response the
   congestion window (cwnd) will usually be 4*SMSS (four maximum-sized
   segments).  If the application's response were larger than 4*SMSS,
   even if the application had closed the connection, the FIN segment
   would never be sent, and thus the connection would still remain in
   the ESTABLISHED state (rather than transit to the FIN-WAIT-1 state).

7.1.2.  Countermeasures

   The resource exhaustion attack described in Section 7.1.1 does not
   necessarily differ from a legitimate high-load scenario, and
   therefore is hard to mitigate without negatively affecting the
   robustness of TCP.  However, complementary mitigations can still be
   implemented to limit the impact of these attacks.

   Enforcing limits on the number of connections with no user-space
   controlling process

   The considerations and recommendations in Section 5.3.2 for enforcing
   limits on the number of connections with no user-space controlling
   process are applicable to mitigate this vulnerability.

   Enforcing per-user and per-process limits



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   While the Netkill attack is usually targeted at a service such as
   HTTP, its impact is usually system-wide.  This is particularly
   undesirable, as an attack against a single service might affect the
   system as a whole (for example possibly precluding remote system
   administration).

   In order to avoid an attack against a single service from affecting
   other services, we advise TCP implementations to enforce per-process
   and per-user limits on maximum kernel memory that can be used at any
   time.  Additionally, we recommend implementations to enforce per-
   process and per-user limits on the number of existent TCP connections
   at any time.

   Limiting the number of ongoing connections at the application

   The considerations and recommendations in Section 5.3.2 for enforcing
   limits on the number of ongoing connections at the application are
   applicable to mitigate this vulnerability.

   Enabling applications to enforce limits on ongoing connections

   As discussed in Section 6.1.1, the fact that the close() system call
   marks the corresponding file descriptor as closed prevents the
   application running on top of TCP from enforcing limits on the
   corresponding connection.

   While it is common practice for applications to terminate their
   connections by means of the close() system call, it is possible for
   an application to initiate the connection-termination phase without
   closing the corresponding file descriptor (hence keeping control of
   the connection).

   In order to achieve this, an application performing an active close
   (i.e., initiating the connection-termination phase) should replace
   the call to close(sockfd) with the following code sequence:

   o  A call to shutdown(sockfd, SHUT_WR), to close the sending
      direction of this connection

   o  Successive calls to read(), until it returns 0, thus indicating
      that the remote TCP peer has finished sending data.

   o  A call to setsockopt(sockfd, SOL_SOCKET, SO_LINGER, &l,
      sizeof(l)), where l is of type struct linger (with its members
      l.l_onoff=1 and l.l_linger=90).

   o  A call to close(sockfd), to close the corresponding file
      descriptor.



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   The call to shutdown() (instead of close()) allows the application to
   retain control of the underlying TCP connection while the connection
   transitions through the FIN-WAIT-1 and FIN-WAIT-2 states.  However,
   the application will not retain control of the connection while it
   transitions through the CLOSING and TIME-WAIT states.  Nevertheless,
   in these states TCP should not have any pending data to send to the
   remote TCP peer or to be received by the application running on top
   of it, and thus these states are less of a concern for this
   particular vulnerability (Netkill).

   It should be noted that, strictly speaking, close(sockfd) decrements
   the reference count for the descriptor sockfd, and initiates the
   connection termination phase only when the reference count reaches 0.
   On the other hand, shutdown(sockfd, SHUT_WR) initiates the
   connection-termination phase, regardless of the reference count for
   the sockfd descriptor.  This should be taken into account when
   performing the code replacement described above.  For example, it
   would be a bug for two processes (e.g., parent and child) that share
   a descriptor to both call shutdown(sockfd, SHUT_WR).

   An application performing a passive close should replace the call to
   close(sockfd) with the following code sequence:

   o  A call to setsockopt(sockfd, SOL_SOCKET, SO_LINGER, &l,
      sizeof(l)), where l is of type struct linger (with its members
      l.l_onoff=1 and l.l_linger=90).

   o  A call to close(sockfd), to close the corresponding file
      descriptor.

   It is assumed that if the application is performing a passive close,
   the application already detected that the remote TCP peer finished
   sending data by means as a result of a call to read() returning 0.

   In this scenario, the application will not retain control of the
   underlying connection when it transitions through the LAST_ACK state.
   However, in this state TCP should not have any pending data to send
   to the remote TCP peer or to be received by the application running
   on top of TCP, and thus this state is less of a concern for this
   particular vulnerability (Netkill).

   Limiting the number of simultaneous connections at firewalls

   The considerations and recommendations in Section 5.3.2 for enforcing
   limiting the number of simultaneous connections at firewalls are
   applicable to mitigate this vulnerability.

   Performing heuristics on ongoing TCP connections



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   Some heuristics could be performed on TCP connections that may
   possibly help if scarce system requirements such as memory become
   exhausted.  A number of parameters may be useful to perform such
   heuristics.

   In the case of the Netkill attack described in [Shalunov, 2000],
   there are two parameters that are characteristic of a TCP being
   attacked:

   o  A large amount of data queued in the TCP retransmission buffer
      (e.g., the socket send buffer).

   o  Only small amount of data has been successfully transferred to the
      remote peer.

   Clearly, these two parameters do not necessarily indicate an ongoing
   attack.  However, if exhaustion of the corresponding system resources
   was imminent, these two parameters (among others) could be used to
   perform heuristics when considering aborting ongoing connections.

   It should be noted that while an attacker could advertise a zero
   window to cause the target system to tie system memory to the TCP
   retransmission buffer, it is hard to perform any useful statistics
   from the advertised window.  While it is tempting to enforce a limit
   on the length of the persist state (see Section 3.7.2 of this
   document), an attacker could simply open the window (i.e., advertise
   a TCP window larger than zero) from time to time to prevent this
   enforced limit from causing his malicious connections to be aborted.

7.2.  TCP segment reassembly buffer

   TCP MAY discard out-of-order data when system-memory exhaustion is
   imminent.

   DISCUSSION:

      TCP buffers out-of-order segments to more efficiently handle the
      occurrence of packet reordering and segment loss.  When out-of-
      order data are received, a "hole" momentarily exists in the data
      stream which must be filled before the received data can be
      delivered to the application making use of TCP's services.  This
      situation can be exploited by an attacker, which could
      intentionally create a hole in the data stream by sending a number
      of segments with a sequence number larger than the next sequence
      number expected (RCV.NXT) by the attacked TCP.  Thus, the attacked
      TCP would tie system memory to buffer the out-of-order segments,
      without being able to hand the received data to the corresponding
      application.



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      If a large number of such connections were created, system memory
      could be exhausted, precluding the attacked TCP from servicing new
      connections and/or continue servicing TCP connections previously
      established.

      Fortunately, these attacks can be easily mitigated, at the expense
      of degrading the performance of possibly legitimate connections.
      When out-of-order data is received, an Acknowledgement segment is
      sent with the next sequence number expected (RCV.NXT).  This means
      that receipt of the out-of-order data will not be actually
      acknowledged by the TCP's cumulative Acknowledgement Number.  As a
      result, a TCP is free to discard any data that have been received
      out-of-order, without affecting the reliability of the data
      transfer.  Given the performance implications of discarding out-
      of-order segments for legitimate connections, this pruning policy
      should be applied only if memory exhaustion is imminent.

      As a result of discarding the out-of-order data, these data will
      need to be unnecessarily retransmitted.  Additionally, a loss
      event will be detected by the sending TCP, and thus the slow start
      phase of TCP's congestion control will be entered, thus reducing
      the data transfer rate of the connection.

      It is interesting to note that this pruning policy could be
      applied even if Selective Acknowledgements (SACK) (specified in
      RFC 2018 [Mathis et al, 1996]) are in use, as SACK provides only
      advisory information, and does not preclude the receiving TCP from
      discarding data that have been previously selectively-acknowledged
      by means of TCP's SACK option, but not acknowledged by TCP's
      cumulative Acknowledgement Number.

      There are a number of ways in which the pruning policy could be
      triggered.  For example, when out of order data are received, a
      timer could be set, and the sequence number of the out-of-order
      data could be recorded.  If the hole were filled before the timer
      expires, the timer would be turned off.  However, if the timer
      expired before the hole were filled, all the out-of-order segments
      of the corresponding connection would be discarded.  This would be
      a proactive counter-measure for attacks that aim at exhausting the
      receive buffers.

      In addition, an implementation could incorporate reactive
      mechanisms for more carefully controlling buffer allocation when
      some predefined buffer allocation threshold was reached.  At such
      point, pruning policies would be applied.

      A number of mechanisms can aid in the process of freeing system
      resources.  For example, a table of network prefixes corresponding



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      to the IP addresses of TCP peers that have ongoing TCP connections
      could record the aggregate amount of out-of-order data currently
      buffered for those connections.  When the pruning policy was
      triggered, TCP connections with hosts that have network prefixes
      with large aggregate out-of-order buffered data could be selected
      first for pruning the out-of-order segments.

      Alternatively, if TCP segments were de-multiplexed by means of a
      hash table (as it is currently the case in many TCP
      implementations), a counter could be held at each entry of the
      hash table that would record the aggregate out-of-order data
      currently buffered for those connections belonging to that hash
      table entry.  When the pruning policy is triggered, the out-of-
      order data corresponding to those connections linked by the hash
      table entry with largest amount of aggregate out-of-order data
      could be pruned first.  It is important that this hash is not
      computable by an attacker, as this would allow him to maliciously
      cause the performance of specific connections to be degraded.
      That is, given a four-tuple that identifies a connection, an
      attacker should not be able to compute the corresponding hash
      value used by the target system to de-multiplex incoming TCP
      segments to that connection.

      Another variant of a resource exhaustion attack against TCP's
      segment reassembly mechanism would target the data structures used
      to link the different holes in a data stream.  For example, an
      attacker could send a burst of 1 byte segments, leaving a one-byte
      hole between each of the data bytes sent.  Depending on the data
      structures used for holding and linking together each of the data
      segments, such an attack might waste a large amount of system
      memory by exploiting the overhead needed store and link together
      each of these one-byte segments.

      For example, if a linked-list is used for holding and linking each
      of the data segments, each of the involved data structures could
      involve one byte of kernel memory for storing the received data
      byte (the TCP payload), plus 4 bytes (32 bits) for storing a
      pointer to the next node in the linked-list.  Additionally, while
      such a data structure would require only a few bytes of kernel
      memory, it could result in the allocation of a whole memory page,
      thus consuming much more memory than expected.

      Therefore, implementations should enforce a limit on the number of
      holes that are allowed in the received data stream at any given
      time.  When such a limit is reached, incoming TCP segments which
      would create new holes would be silently dropped.  Measurements in
      [Dharmapurikar and Paxson, 2005] indicate that in the vast
      majority of TCP connections have at most a single hole at any



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      given time.  A limit of 16 holes for each connection would
      accommodate even most of the very unusual cases in which there can
      be more than hole in the data stream at a given time.

      [US-CERT, 2004a] is a security advisory about a Denial of Service
      vulnerability resulting from a TCP implementation that did not
      enforce limits on the number of segments stored in the TCP
      reassembly buffer.

      Section 8 of this document describes the security implications of
      the TCP segment reassembly algorithm.

7.3.  Automatic buffer tuning mechanisms

7.3.1.  Automatic send-buffer tuning mechanisms

   A TCP implementing an automatic send-buffer tuning mechanism SHOULD
   enforce the following limit on the size of the send buffer of each
   TCP connection:

   send_buffer_size <= send_buffer_pool / (min_buffer_size * max_connections)

   where

   send_buffer_size:
      Maximum send buffer size to be used for this connection

   send_buffer_pool:
      Total amount of system memory meant for TCP send buffers

   min_buffer_size:
      Minimum send buffer size for each TCP connection

   max_connections:
      Maximum number of TCP connections this system is expected to
      handle at a time

   max_connections may be an artificial limit enforced by the system
   administrator specifically on the number of TCP connections, or may
   be derived from some other system limit (e.g., the maximum number of
   file descriptors)

   DISCUSSION:

      A number of TCP implementations incorporate automatic tuning
      mechanisms for the TCP send buffer size.  In most of them, the
      underlying idea is to set the send buffer to some multiple of the
      congestion window (cwnd).  This type of mechanism usually improves



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      TCP's performance, by preventing the socket send buffer from
      becoming a bottleneck, while avoiding the need to simply
      overestimate the TCP send buffer size (i.e., make it arbitrarily
      large).  [Semke et al, 1998] discusses such an automatic buffer
      tuning mechanism.

      Unfortunately, automatic tuning mechanisms can be exploited by
      attackers to amplify the impact of other resource exhaustion
      attacks.  For example, an attacker could establish a TCP
      connection with a victim host, and cause the congestion window to
      be increased (either legitimately or illegitimately).  Once the
      congestion window (and hence the TCP send buffer) is increased, he
      could cause the corresponding system memory to be tied up by
      advertising a zero-byte TCP window (see Section 3.7) or simply not
      acknowledging any data, thus amplifying the effect of resource
      exhaustion attacks such as that discussed in Section 7.1.1.

      When an automatic buffer tuning mechanism is implemented, a number
      of countermeasures should be incorporated to prevent the mechanism
      from being exploited to amplify other resource exhaustion attacks.

      Firstly, appropriate policies should be applied to guarantee fair
      use of the available system memory by each of the established TCP
      connections.  Secondly, appropriate policies should be applied to
      avoid existing TCP connections from consuming all system
      resources, thus preventing service to new TCP connections.

      Appendix A of [Semke et al, 1998] proposes an algorithm for the
      fair share of the available system memory among the established
      connections.  However, there are a number of limits that should be
      enforced on the system memory assigned for the send buffer of each
      connection.  Firstly, each connection should always be assigned
      some minimum send buffer space that would enable TCP to perform at
      an acceptable performance.  Secondly, some system memory should be
      reserved for future connections, according to the maximum number
      of concurrent TCP connections that are expected to be successfully
      handled at any given time.

      These limits preclude the automatic tuning algorithm from
      assigning all the available memory buffers to ongoing connections,
      thus preventing the establishment of new connections.

      Even if these limits are enforced, an attacker could still create
      a large number of TCP connections, each of them tying valuable
      system resources.  Therefore, in scenarios in which most of the
      system memory reserved for TCP send buffers is allocated to
      ongoing connections, it may be necessary for TCP to enforce some
      policy to free resources to either service more TCP connections,



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      or to be able to improve the performance of other existing
      connections, by allocating more resources to them.

      When needing to free memory in use for send buffers, particular
      attention should be paid to TCP's that have a large amount of data
      in the socket send buffer, and that at the same time fall into any
      of these categories:



      *  The remote TCP peer that has been advertising a small (possibly
         zero) window for a considerable period of time.

      *  There have been a large number of retransmissions of segments
         corresponding to the first few windows of data.

      *  Connections that fall into one of the previous categories, for
         which only a reduced amount of data have been successfully
         transferred to the peer TCP since the connection was
         established.

      Unfortunately, all these cases are valid scenarios for the TCP
      protocol, and thus aborting connections that fall in any of these
      categories has the potential of causing interoperability problems.
      However, in scenarios in which all system resources are allocated,
      it may make sense to free resources allocated to TCP connections
      which are tying a considerable amount of system resources and that
      have not made progress in a considerable period of time.

7.3.2.  Automatic receive-buffer tuning mechanism

   A number of TCP implementations include automatic tuning mechanisms
   for the receive buffer size.  These mechanisms aim at setting the
   socket buffer to a size that is large enough to avoid the TCP window
   from becoming a bottleneck that would limit TCP's throughput, without
   wasting system memory by over-sizing it.

   [Heffner, 2002] describes a mechanism for the automatic tuning of the
   socket receive buffer.  Basically, the mechanism aims at measuring
   the amount of data received during a RTT (Round-Trip Time), and
   setting the socket receive buffer to some multiple of that value.

   A TCP implementing an automatic receive-buffer tuning mechanism
   SHOULD enforce the following limit on the size of the receive buffer
   of each TCP connection:

   recv_buffer_size <= recv_buffer_pool / (min_buffer_size * max_connections)




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   where:

   recv_buffer_size:
      Maximum receive buffer size to be used for this connection

   recv_buffer_pool:
      Total amount of system memory meant for TCP receive buffers

   min_buffer_size:
      Minimum receive buffer size for each TCP connection

   max_connections:
      Maximum number of TCP connections this system is expected to
      handle at a time

   max_connections may be an artificial limit enforced by the system
   administrator specifically on the number of TCP connections, or may
   be derived from some other system limit (e.g., the maximum number of
   file descriptors).

   DISCUSSION:

      Unfortunately, automatic tuning mechanisms for the socket receive
      buffer can be exploited to perform a resource exhaustion attack.
      An attacker willing to exploit the automatic buffer tuning
      mechanism would first establish a TCP connection with the victim
      host.  Subsequently, he would start a bulk data transfer to the
      victim host.  By carefully responding to the peer's TCP segments,
      the attacker could cause the peer TCP to measure a large data/RTT
      value, which would lead to the adoption of an unnecessarily large
      socket receive buffer.  For example, the attacker could
      optimistically send more data than those allowed by the TCP window
      advertised by the remote TCP.  Those extra data would cross in the
      network with the window updates sent by the remote TCP, and could
      lead the TCP receiver to measure a data/RTT twice as big as the
      real one.  Alternatively, if the TCP timestamp option (specified
      in RFC 1323 [Jacobson et al, 1992]) is used for RTT measurement,
      the attacker could lead the TCP receiver to measure a small RTT
      (and hence a large Data/RTT rate) by "optimistically" echoing
      timestamps that have not yet been received.

      Finally, once the TCP receiver is led to increase the size of its
      receive buffer, the attacker would transmit a large amount of
      data, filling the whole peer's receive buffer except for a few
      bytes at the beginning of the window (RCV.NXT).  This gap would
      prevent the peer application from reading the data queued by TCP,
      thus tying system memory to the received data segments until (if
      ever) the peer application times out.



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      A number of limits should be enforced on the amount of system
      memory assigned to any given connection.  Firstly, each connection
      should always be assigned some minimum receive buffer space that
      would enable TCP to perform at a minimum acceptable performance.
      Additionally, some system memory should be reserved for future
      connections, according to the maximum number of concurrent TCP
      connections that are expected to be successfully handled at any
      given time.

      These limits preclude the automatic tuning algorithm from
      assigning all the available memory buffers to existing
      connections, thus preventing the establishment of new connections.

      It is interesting to note that a TCP sender will always try to
      retransmit any data that have not been acknowledged by TCP's
      cumulative acknowledgement.  Therefore, if memory exhaustion is
      imminent, a system should consider freeing those memory buffers
      used for TCP segments that were received out of order,
      particularly when a given connection has been keeping a large
      number of out-of-order segments in the receive buffer for a
      considerable period of time.

      It is worth noting that TCP Selective Acknowledgements (SACK) are
      advisory, in the sense that a TCP that has SACKed (but not ACKed)
      a block of data is free to discard that block, and expect the TCP
      sender to retransmit them when the retransmission timer of the
      peer TCP expires.


8.  TCP segment reassembly algorithm

8.1.  Problems that arise from ambiguity in the reassembly process

   If a TCP segment is received containing some data bytes that had
   already been received, the first copy of those data SHOULD be used
   for reassembling the application data stream.

   DISCUSSION:

      A security consideration that should be made for the TCP segment
      reassembly algorithm is that of data stream consistency between
      the host performing the TCP segment reassembly, and a Network
      Intrusion Detection System (NIDS) being employed to monitor the
      host in question.

      In the event a TCP segment was unnecessarily retransmitted, or
      there was packet duplication in any of the intervening networks, a
      TCP might get more than one copy of the same data.  Also, as TCP



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      segments can be re-packetized when they are retransmitted, a given
      TCP segment might partially overlap data already received in
      earlier segments.  In all these cases, the question arises about
      which of the copies of the received data should be used when
      reassembling the data stream.  In legitimate and normal
      circumstances, all copies would be identical, and the same data
      stream would be obtained regardless of which copy of the data was
      used.  However, an attacker could maliciously send overlapping
      segments containing different data, with the intent of evading a
      Network Intrusion Detection Systems (NIDS), which might reassemble
      the received TCP segments differently than the monitored system.
      [Ptacek and Newsham, 1998] provides a detailed discussion of these
      issues.

      As suggested in Section 3.9 of RFC 793 [Postel, 1981c], if a TCP
      segment arrives containing some data bytes that have already been
      received, the first copy of those data should be used for
      reassembling the application data stream.  It should be noted that
      while convergence to this policy might prevent some cases of
      ambiguity in the reassembly process, there are a number of other
      techniques that an attacker could still exploit to evade a NIDS
      [CPNI, 2008].  These techniques can generally be defeated if the
      NIDS is placed in-line with the monitored system, thus allowing
      the NIDS to normalize the network traffic or apply some other
      policy that could ensure consistency between the result of the
      segment reassembly process obtained by the monitored host and that
      obtained by the NIDS.

      [CERT, 2003] and [CORE, 2003] are advisories about a heap buffer
      overflow in a popular Network Intrusion Detection System resulting
      from incorrect sequence number calculations in its TCP stream-
      reassembly module.


9.  TCP Congestion Control

   TCP implements two algorithms, "slow start" and "congestion
   avoidance", for controlling the rate at which data is transmitted on
   a TCP connection [Allman et al, 1999].  These algorithms require the
   addition of two variables as part of TCP per-connection state: cwnd
   and ssthresh.

   The congestion window (cwnd) is a sender-side limit on the amount of
   outstanding data that the sender can have at any time, while the
   receiver's advertised window (rwnd) is a receiver-side limit on the
   amount of outstanding data.  The minimum of cwnd and rwnd governs
   data transmission.




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   Another state variable, the slow-start threshold (ssthresh), is used
   to determine whether it is the slow start or the congestion avoidance
   algorithm that should control data transmission.  When cwnd <
   ssthresh, "slow start" governs data transmission, and the congestion
   window (cwnd) is exponentially increased.  When cwnd > ssthresh,
   "congestion avoidance" governs data transmission, and the congestion
   window (cwnd) is only linearly increased.

   As specified in RFC 2581 [Allman et al, 1999], when cwnd and ssthresh
   are equal the sender may use either slow start or congestion
   avoidance.

   During slow start, TCP increments cwnd by at most SMSS bytes for each
   ACK received that acknowledges new data.  During congestion
   avoidance, cwnd is incremented by 1 full-sized segment per round-trip
   time (RTT), until congestion is detected.

   Additionally, TCP uses two algorithms, Fast Retransmit and Fast
   Recovery, to mitigate the effects of packet loss.  The "Fast
   Retransmit" algorithm infers packet loss when three Duplicate
   Acknowledgements (DupACKs) are received.

   The value "three" is meant to allow for fast-retransmission of
   "missing" data, while avoiding network packet reordering from
   triggering loss recovery.

   Once packet loss is detected by the receipt of three duplicate-ACKs,
   the "Fast Recovery" algorithm governs the transfer of new data until
   a non-duplicate ACK is received that acknowledges the receipt of new
   data.  The Fast Retransmit and Fast Recovery algorithms are usually
   implemented together, as follows (from RFC 2581):

   o  When the third duplicate ACK is received, set ssthresh to no more
      than the value given in the equation: ssthresh = max (FlightSize /
      2, 2*SMSS)

   o  Retransmit the lost segment and set cwnd to ssthresh plus 3*SMSS.
      This artificially "inflates" the congestion window by the number
      of segments (three) that have left the network and which the
      receiver has buffered.

   o  For each additional duplicate ACK received, increment cwnd by
      SMSS.  This artificially inflates the congestion window in order
      to reflect the additional segment that has left the network.

   o  Transmit a segment, if allowed by the new value of cwnd and the
      receiver's advertised window.




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   o  When the next ACK arrives that acknowledges new data, set cwnd to
      ssthresh (the value set in step 1).  This is termed "deflating"
      the window.

9.1.  Congestion control with misbehaving receivers

   [Savage et al, 1999] describes a number of ways in which TCP's
   congestion control mechanisms can be exploited by a misbehaving TCP
   receiver to obtain more than its fair share of bandwidth.  The
   following subsections provide a brief discussion of these
   vulnerabilities, along with the possible countermeasures.

9.1.1.  ACK division

   TCP SHOULD increase cwnd by one SMSS only when a valid ACK covers the
   entire data segment sent

   (note: or should we recommend the other counter-measure (i.e.,
   implementation of ABC?)

   DISCUSSION:

      Given that TCP updates cwnd based on the number of duplicate ACKs
      it receives, rather than on the amount of data that each ACK is
      actually acknowledging, a malicious TCP receiver could cause the
      TCP sender to illegitimately increase its congestion window by
      acknowledging a data segment with a number of separate
      Acknowledgements, each covering a distinct piece of the received
      data segment.

             See Figure 7, in page 64 of the UK CPNI document.

                             ACK division attack

      [Savage et al, 1999] describes two possible countermeasures for
      this vulnerability.  One of them is to increment cwnd not by a
      full SMSS, but proportionally to the amount of data being
      acknowledged by the received ACK, similarly to the policy
      described in RFC 3465 [Allman, 2003].  Another alternative is to
      increase cwnd by one SMSS only when a valid ACK covers the entire
      data segment sent.

9.1.2.  DupACK forgery

   TCP SHOULD keep track of the number of outstanding segments (o_seg),
   and accept only up to (o_seg -1) duplicate Acknowledgements.

   DISCUSSION:



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      The second vulnerability discussed in [Savage et al, 1999] allows
      an attacker to cause the TCP sender to illegitimately increase its
      congestion window by forging a number of duplicate
      Acknowledgements (DupACKs).  Figure 8 shows a sample scenario.
      The first three DupACKs trigger the Fast Recovery mechanism, while
      the rest of them cause the congestion window at the TCP sender to
      be illegitimately inflated.  Thus, the attacker is able to
      illegitimately cause the TCP sender to increase its data
      transmission rate.

             See Figure 8, in page 65 of the UK CPNI document.

                            DupACK forgery attack

      Fortunately, a number of sender-side heuristics can be implemented
      to mitigate this vulnerability.  First, the TCP sender could keep
      track of the number of outstanding segment (o_seg), and accept
      only up to (o_seg -1) DupACKs.  Secondly, a TCP sender might, for
      example, refuse to enter Fast Recovery multiple times in some
      period of time (e.g., one RTT).

      [Savage et al, 1999] also describes a modification to TCP to
      implement a nonce protocol that would eliminate this
      vulnerability.  However, this would require modification of all
      implementations, which makes this counter-measure hard to deploy.

9.1.3.  Optimistic ACKing

   Another alternative for an attacker to exploit TCP's congestion
   control mechanisms is to acknowledge data that has not yet been
   received, thus causing the congestion window at the TCP sender to be
   incremented faster than it should.

             See Figure 9, in page 66 of the UK CPNI document.

                         Optimistic ACKing attack

   [Savage et al, 1999] describes a number of mitigations for this
   vulnerability.  Firstly, it describes a countermeasure based on the
   concept of "cumulative nonce", which would allow a receiver to prove
   that it has received all the segments it is acknowledging.  However,
   this countermeasure requires the introduction of two new fields to
   the TCP header, thus requiring a modification to all the
   communicating TCPs, makes this counter-measure hard to deploy.
   Secondly, it describes a possible way to encode the nonce in a TCP
   segment by carefully modifying its size.  While this countermeasure
   could be easily deployed (as it is just sender side policy), we
   believe that middle-boxes such as protocol-scrubbers might prevent



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   this counter-measure from working as expected.  Finally, it suggests
   that a TCP sender might penalize a TCP receiver that acknowledges
   data not yet sent by resetting the corresponding connection.  Here we
   discourage the implementation of this policy, as it would provide an
   attack vector for a TCP-based connection-reset attack, similar to
   those described in Section 11.

   [US-CERT, 2005a] is a vulnerability advisory about this issue.

9.2.  Blind DupACK triggering attacks against TCP

   While all of the attacks discussed in [Savage et al, 1999] have the
   goal of increasing the performance of the attacker's TCP connections,
   TCP congestion control mechanisms can be exploited with a variety of
   goals.

   Firstly, if bursts of many duplicate-ACKs are sent to the "sending
   TCP", the third duplicate-ACK will cause the "lost" segment to be
   retransmitted, and each subsequent duplicate-ACK will cause cwnd to
   be artificially inflated.  Thus, the "sending TCP" might end up
   injecting more packets into the network than it really should, with
   the potential of causing network congestion.  This is a potential
   consequence of the "Duplicate-ACK spoofing attack" described in
   [Savage et al, 1999].

   Secondly, if bursts of three duplicate ACKs are sent to the TCP
   sender, the attacked system would infer packet loss, and ssthresh and
   cwnd would be reduced.  As noted in RFC 2581 [Allman et al, 1999],
   causing two congestion control events back-to-back will often cut
   ssthresh and cwnd to their minimum value of 2*SMSS, with the
   connection immediately entering the slower-performing congestion
   avoidance phase.  While it would not be attractive for an attacker to
   perform this attack against one of his TCP connections, the attack
   might be attractive when the TCP connection to be attacked is
   established between two other parties.

   It is usually assumed that in order for an off-path attacker to
   perform attacks against a third-party TCP connection, he should be
   able to guess a number of values, including a valid TCP Sequence
   Number and a valid TCP Acknowledgement Number.  While this is true if
   the attacker tries to "inject" valid packets into the connection by
   himself, a feature of TCP can be exploited to fool one of the TCP
   endpoints to transmit valid duplicate Acknowledgements on behalf of
   the attacker, hence relieving the attacker of the hard task of
   forging valid values for the Sequence Number and Acknowledgement
   Number TCP header fields.

   Section 3.9 of RFC 793 [Postel, 1981c] describes the processing of



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   incoming TCP segments as a function of the connection state and the
   contents of the various header fields of the received segment.  For
   connections in the ESTABLISHED state, the first check that is
   performed on incoming segments is that they contain "in window" data.
   That is,

                 RCV.NXT <= SEG.SEQ <= RCV.NXT+RCV.WND, or


               RCV.NXT <= SEG.SEQ+SEG.LEN-1 < RCV.NXT+RCV.WND

   If a segment does not pass this check, it is dropped, and an
   Acknowledgement is sent in response:

                    <SEQ=SND.NXT><ACK=RCV.NXT><CTL=ACK>

   The goal of this behavior is that, in the event data segments are
   received by the TCP receiver, but all the corresponding
   Acknowledgements are lost, when the TCP sender retransmits the
   supposedly lost data, the TCP receiver will send an Acknowledgement
   reflecting all the data received so far.  If "old" TCP segments were
   silently dropped, the scenario just described would lead to a
   "frozen" TCP connection, with the TCP sender retransmitting the data
   for which it has not yet received an Acknowledgement, and the TCP
   receiver silently ignoring these segments.  Additionally, it helps
   TCP to detect half-open connections.

   This feature implies that, provided the four-tuple that identifies a
   given TCP connection is known or can be easily guessed, an attacker
   could send a TCP segment with an "out of window" Sequence Number to
   one of the endpoints of the TCP connection to cause it to send a
   valid ACK to the other endpoint of the connection.  Figure 10
   illustrates such a scenario.

             See Figure 10, in page 68 of the UK CPNI document.

                       Blind Dup-ACK forgery attack

   As discussed in [Watson, 2004] and RFC 4953 [Touch, 2007], there are
   a number of scenarios in which the four-tuple that identifies a TCP
   connection is known or can be easily guessed.  In those scenarios, an
   attacker could perform any of the "blind" attacks described in the
   following subsections by exploiting the technique described above.

   The following subsections describe blind DupACK-triggering attacks
   that aim at either degrading the performance of an arbitrary
   connection, or causing a TCP sender to illegitimately increase the
   rate at which it transmits data, potentially leading to network



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   congestion.

9.2.1.  Blind throughput-reduction attack

   As discussed in Section 9, when three duplicate Acknowledgements are
   received, the congestion window is reduced to half the current amount
   of outstanding data (FlightSize).  Additionally, the slow-start
   threshold (ssthresh) is reduced to the same value, causing the
   connection to enter the slower-performing congestion avoidance phase.
   If two congestion-control events occur back to back, ssthresh and
   cwnd will often be reduced to their minimum value of 2*SMSS.

   An attacker could exploit the technique described in Section 9.2 to
   cause the throughput of the attacked TCP connection to be reduced, by
   eliciting three duplicate acknowledgements from the TCP receiver,
   which would cause the TCP sender to reduce its congestion window.  In
   principle, the attacker would need to send a burst of only three out-
   of-window segments.  However, in case the TCP receiver implements an
   acknowledgement policy such as "ACK every other segment", four out-
   of-window segments might be needed.  The first segment would cause
   the pending (delayed) Acknowledgement to be sent, and the next three
   segments would elicit the actual duplicate Acknowledgements.

   Figure 11 shows a time-line graph of a sample scenario.  The burst of
   DupACKs (in green) elicited by the burst of out-of-window segments
   (in red) sent by the attacker causes the TCP sender to retransmit the
   missing segment (in blue) and enter the loss recovery phase.  Once a
   segment that acknowledges new data is received by the TCP sender, the
   loss recovery phase ends, and cwnd and ssthresh are set to half the
   number of segments that were outstanding when the loss recovery phase
   was entered.

             See Figure 11, in page 69 of the UK CPNI document.

            Blind throughput-reduction attack (time-line graph)

   The graphic assumes that the TCP receiver sends an Acknowledgement
   for every other data segment it receives, and that the TCP sender
   implements Appropriate Byte Counting (specified in RFC 3465 [Allman,
   2003]) on the received Acknowledgement segments.  However,
   implementation of these policies is not required for the attack to
   succeed.

9.2.2.  Blind flooding attack

   As discussed in Section 9, when three duplicate Acknowledgements are
   received, the "lost" segment is retransmitted, and the congestion
   window is artificially inflated for each DupACK received, until the



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   loss recovery phase ends.  By sending a long burst of out-of-window
   segments to the TCP receiver of the attacked connection, an attacker
   could elicit a long burst of valid duplicate acknowledgements that
   would illegitimately cause the TCP sender of the attacked TCP
   connection to increase its data transmission rate.

   Figure 12 shows a time-line graph for this attack.  The long burst of
   DupACKs (in green) elicited by the long burst of out-of-window
   segments (in red) sent by the attacker causes the TCP sender to enter
   the loss recovery phase and illegitimately inflate the congestion
   window, leading to an increase in the data transmission rate.  Once a
   segment that acknowledges new data is received by the TCP sender, the
   loss recovery phase ends, and the data transmission rate is reduced.

             See Figure 12, in page 70 of the UK CPNI document.

                  Blind flooding attack (time-line graph)

   Figure 13 is a time-sequence graph produced from packet logs obtained
   from tests of the described attack in a real network.  A burst of
   segments is sent upon receipt of the burst of Duplicate
   Acknowledgements illegitimately elicited by the attacker.  Figure 14
   is an averaged-throughput graphic for the same time frame, which
   clearly shows the effect of the attack in terms of throughput.

             See Figure 13, in page 71 of the UK CPNI document.

                Blind flooding attack (time sequence graph)


             See Figure 14, in page 71 of the UK CPNI document.

             Blind flooding attack (averaged throughput graph)

   These graphics were produced with Shawn Ostermann's tcptrace tool
   [Ostermann, 2008].  An explanation of the format of the graphics can
   be found in tcptrace's manual (available at the project's web site:
   http://www.tcptrace.org).

9.2.3.  Difficulty in performing the attacks

   In order to exploit the technique described in Section 9.2 of this
   document, an attacker would need to know the four-tuple {IP Source
   Address, TCP Source Port, IP Destination Address, TCP Destination
   Port} that identifies the connection to be attacked.  As discussed by
   [Watson, 2004] and RFC 4953 [Touch, 2007], there are a number of
   scenarios in which these values may be known or easily guessed.




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   It is interesting to note that the attacks described in Section 9.2
   of this document will typically require a much smaller number of
   packets than other "blind" attacks against TCP, such as those
   described in [Watson, 2004] and RFC 4953 [Touch, 2007], as the
   technique discussed in Section 9.2 relieves the attacker from having
   to guess valid TCP Sequence Numbers and a TCP Acknowledgement
   numbers.

   The attacks described in Section 9.2.1 and Section 9.2.2 of this
   document require the attacker to forge the source address of the
   packets it sends.  Therefore, if ingress/egress filtering is
   performed by intermediate systems, the attacker's packets would not
   get to the intended recipient, and thus the attack would not succeed.
   However, we consider that ingress/egress filtering cannot be relied
   upon as the first line of defense against these attacks.

   Finally, it is worth noting that in order to successfully perform the
   blind attacks discussed in Section 9.2.1 and Section 9.2.2 of this
   document, the burst of out-of-sequence segments sent by the attacker
   should not be intermixed with valid data segments sent by the TCP
   sender, or else the Acknowledgement number of the illegitimately-
   elicited ACK segments would change, and the Acknowledgements would
   not be considered "Duplicate Acknowledgements" by the TCP sender.
   Tests performed in real networks seem to suggest that this
   requirement is not hard to fulfill, though.

9.2.4.  Modifications to TCP's loss recovery algorithms

   There are a number of algorithms that augment TCP's loss recovery
   mechanism that have been suggested by TCP researchers and have been
   specified by the IETF in the RFC series.  This section describes a
   number of these algorithms, and discusses how their implementation
   affects (or not) the vulnerability of TCP to the attacks discussed in
   Section 9.2.1 and Section 9.2.2 of this document.

   NewReno

   RFC 3782 [Floyd et al, 2004] specifies the NewReno algorithm, which
   is meant to improve TCP's performance in the presence of multiple
   losses in a single window of data.  The implication of this algorithm
   with respect to the attacks discussed in the previous sections is
   that whenever either of the attacks is performed against a connection
   with a NewReno TCP sender, a full-window (or half a window) of data
   will be unnecessarily retransmitted.  This is particularly
   interesting in the case of the blind-flooding attack, as the attack
   would elicit even more packets from the TCP sender.

   Whether a full-window or just half a window of data is retransmitted



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   depends on the Acknowledgement policy at the TCP receiver.  If the
   TCP receiver sends an Acknowledgement (ACK) for every segment, a
   full-window of data will be retransmitted.  If the TCP receiver sends
   an Acknowledgement (ACK) for every other segment, then only half a
   window of data will be retransmitted.

   Figure 15 is a time-sequence graph produced from packet logs obtained
   from tests performed in a real network.  Once loss recovery is
   illegitimately triggered by the duplicate-ACKs elicited by the
   attacker, an entire flight of data is unnecessarily retransmitted.
   Figure 16 is an averaged-throughput graphic for the same time-frame,
   which shows an increase in the throughput of the connection resulting
   from the retransmission of segments governed by NewReno's loss
   recovery.

             See Figure 15, in page 73 of the UK CPNI document.

                NewReno loss recovery (time-sequence graph)


             See Figure 16, in page 74 of the UK CPNI document.

             NewReno loss recovery (averaged throughput graph)

   Limited Transmit

   RFC 3042 [Allman et al, 2001] proposes an enhancement to TCP to more
   effectively recover lost segments when a connection's congestion
   window is small, or when a large number of segments are lost in a
   single transmission window.  The "Limited Transmit" algorithm calls
   for sending a new data segment in response to each of the first two
   Duplicate Acknowledgements that arrive at the TCP sender.  This would
   provide two additional transmitted packets that may be useful for the
   attacker in the case of the blind flooding attack described in
   Section 9.2.2 is performed.

   SACK-based loss recovery

   RFC 3517 [Blanton et al, 2003] specifies a conservative loss-recovery
   algorithm that is based on the use of the selective acknowledgement
   (SACK) TCP option.  The algorithm uses DupACKs as an indication of
   congestion, as specified in RFC 2581 [Allman et al, 1999].  However,
   a difference between this algorithm and the basic algorithm described
   in RFC 2581 is that it clocks out segments only with the SACK
   information included in the DupACKs.  That is, during the loss
   recovery phase, segments will be injected in the network only if the
   SACK information included in the received DupACKs indicates that one
   or more segments have left the network.  As a result, those systems



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   that implement SACK-based loss recovery will not be vulnerable to the
   blind flooding attack described in Section 9.2.2.  However, as RFC
   3517 does not actually require DupACKs to include new SACK
   information (corresponding to data that has not yet been acknowledged
   by TCP's cumulative Acknowledgement), systems that implement SACK-
   based loss-recovery may still remain vulnerable to the blind
   throughput-reduction attack described in Section 9.2.1.  SACK-based
   loss recovery implementations should be updated to implement the
   countermeasure ("Use of SACK information to validate DupACKs")
   described in Section 9.2.5.

9.2.5.  Countermeasures

   TCP SHOULD validate the Sequence Number of an incomming TCP segment
   as follows:

           RCV.NXT - MAX.RCV.WND <= SEG.SEQ <= RCV.NXT + RCV.WND

   where MAX.RCV.WND is the largest TCP window that has so far been
   advertised to the remote endpoint.

   If a segment passes this check, the processing rules specified in RFC
   793 [Postel, 1981c] MUST applied.  Otherwise, TCP SHOULD send an ACK
   (as specified by the processing rules in RFC 793 [Postel, 1981c]),
   applying rate-limiting to the Acknowledgement segments sent in
   response to out-of-window segments.

   DISCUSSION:

      As discussed in Section 9.2, TCP responds with an ACK when an out-
      of-window segment is received, to accommodate those scenarios in
      which the Acknowledgement segments that correspond to some
      received data are lost in the network, and to help discover half-
      open TCP connections.

      However, it is possible to restrict the sequence numbers that are
      considered acceptable, and have TCP respond with ACKs only when it
      is strictly necessary.

      A feature of TCP is that, in some scenarios, it can detect half-
      open connections.  If an implementation chose to silently drop
      those TCP segments that do not pass the check enforced by the
      equation above, it could prevent TCP from detecting half-open
      connections.  Figure 17 shows a scenario in which, provided that
      "TCP B" behaves as specified in RFC 793, a half-open connection
      would be discovered and aborted.





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      An established connection is said to be "half open" if one of the
      TCPs has closed or aborted the connection at its end without the
      knowledge of the other, or if the two ends of the connection have
      become desynchronized owing to a crash that resulted in loss of
      memory.

             See Figure 17, in page 76 of the UK CPNI document.

                        Half-Open Connection Discovery

      In the scenario illustrated by Figure 17, TCP A crashes losing the
      connection-state information of the TCP connection with TCP B. In
      line 3, TCP A tries to establish a new connection with TCP B,
      using the same four-tuple {IP Source Address, TCP source port, IP
      Destination Address, TCP destination port}.  In line 4, as the SYN
      segment is out of window, TCP B responds with an ACK.  This ACK
      elicits an RST segment from TCP A, which causes the half-open
      connection at TCP B to be aborted.

      If the SYN segment had been "in window", TCP B would have sent an
      RST segment instead, which would have closed the half-open
      connection.  Ongoing work at the TCPM WG of the IETF proposes to
      change this behavior, and make TCP respond to a SYN segment
      received for any of the synchronized states with an ACK segment,
      to avoid in-window SYN segments from being used to perform
      connection-reset attacks [Ramaiah et al, 2008].

      However, in case the out-of-window segment was silently dropped,
      the scenario in Figure 17 would change into that in Figure 18.

             See Figure 18, in page 76 of the UK CPNI document.

       Half-Open Connection Discovery with the proposed counter-measure

      In line 3, the SYN segment sent by TCP A is silently dropped by
      TCP B because it does not pass the check enforced by the equation
      above (i.e., it contains an out-of-window sequence number).  As a
      result, some time later (an RTO) TCP A retransmits its SYN
      segment.  Even after TCP A times out, the half-open connection at
      TCP B will remain in the same state.

      Thus, a conservative reaction to those segments that do not pass
      the check enforced by the equation above would be to respond with
      an Acknowledgement segment (as specified by RFC 793), applying
      rate-limiting to those Acknowledgement segments sent in response
      to segments that do not pass the check enforced by that equation.
      An implementation might choose to enforce a rate-limit of, e.g.,
      one ACK per five seconds, as a single ACK segment is needed for



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      the Half-Open Connection Discovery mechanism to work.

      As the only reason to respond with an ACK to those segments that
      do not pass the check enforced by the equation above is to allow
      TCP to discover half-open connections, an aggressive rate-limit
      can be enforced.  As long as the rate-limit prevents out-of-window
      segments from eliciting three Acknowledgment segments in a Round-
      trip Time (RTT), an attacker would not be able to trigger TCP's
      loss-recovery, and thus would not be able to perform the attacks
      described in the previous sections.

      It is interesting to note that RFC 793 [Postel, 1981c] itself
      states that half-open connections are expected to be unusual.
      Additionally, given that in many scenarios it may be unlikely for
      a TCP connection request to be issued with the same four-tuple as
      that of the half-open connection, a complete solution for the
      discovery of half-open connections cannot rely on the mechanism
      illustrated by Figure 17, either.  Therefore, some implementations
      might choose to sacrifice TCP's ability to detect half-open
      connections, and have a more aggressive reaction to those segments
      that do not pass the check enforced by the equation above by
      silently dropping them.

      This validation check can also help to avoid ACK wars in some
      scenarios that may arise from the use of transparent proxies.  In
      those scenarios, when the transparent proxy fails to wire (i.e.,
      is disabled), the sequence numbers of the two end-points of the
      TCP connection become desynchronized, and both TCPs begin to send
      duplicate Acknowledgements to each other, with the intention of
      re-synchronizing them.  As the sequence numbers never get re-
      synchronized, the ACK war can only be stopped by an external
      agent.

   TCP SHOULD limit the number of duplicate acknowledgements it will
   honour to:

                   Max_DupACKs = (FlightSize / SMSS) - 1

   Where FlightSize and SMSS are the values defined in RFC 2581 [Allman
   et al, 1999].  When more than Max_DupACKs duplicate acknowledgements
   are received, the exceeding DupACKs should be silently dropped.

   DISCUSSION:

      Note that duplicate acknowledgements should be elicited by out-of-
      order segments.

   In the case of TCP connections that have agreed to employ SACK, TCP



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   SHOULD validate duplicate ACKs with the following criteria: Valid
   Duplicate ACKs MUST contain new SACK information.  The SACK
   information MUST refer to data that has already been sent, but that
   has not yet been acknowledged by TCP's cumulative Acknowledgement.  A
   TCP segment that does not pass this check SHOULD NOT be considered as
   "duplicate Acknowledgement".

   DISCUSSION:

      SACK, specified in 2018 [Mathis et al, 1996], provides a mechanism
      for TCP to be able to acknowledge the receipt of out-of-order TCP
      segments.  For connections that have agreed to use SACK, each
      legitimate DupACK will contain new SACK information that reflects
      the data bytes contained in the out-of-order data segment that
      elicited the DupACK.

      RFC 3517 [Blanton et al, 2003] specifies a SACK-based loss
      recovery algorithm for TCP.  However, it does recommend TCP
      implementations to validate DupACKs by requiring that they contain
      new SACK information.  Results obtained from auditing a number of
      TCP implementations seem to indicate that most TCP implementations
      do not enforce this validation check on incoming DupACKs, either.

      In the case of TCP connections that have agreed to use SACK, a
      validation check should be performed on incoming ACK segments to
      completely eliminate the attacks described in Section 9.2.1 and
      Section 9.2.2 of this document: "Duplicate ACKs should contain new
      SACK information.  The SACK information should refer to data that
      has already been sent, but that has not yet been acknowledged by
      TCP's cumulative Acknowledgement".

      Those ACK segments that do not comply with this validation check
      should not be considered "duplicate ACKs", and thus should not
      trigger the loss-recovery phase.

      In case at least one segment in a window of data has been lost,
      the successive segments will elicit the generation of Duplicate
      ACKs containing new SACK information.  This SACK information will
      indicate the receipt of these successive segments by the TCP
      receiver.

      In the case of pure ACKs illegitimately elicited by out-of-window
      segments, however, the ACKs will not contain any SACK information.

      If DSACK (specified in 2883 [Floyd et al, 2000]) were implemented
      by the TCP receiver, then the illegitimately elicited DupACKs
      might contain out-of-window SACK information if the sequence
      number of the forged TCP segment (SEG.SEQ) is lower than the next



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      expected sequence number (RECV.NXT) at the TCP receiver.  Such
      segments should be considered to indicate the receipt of duplicate
      data, rather than an indication of lost data, and therefore should
      not trigger loss recovery.

   Other possible general mitigations are discussed in the following
   paragraphs:

   TCP port number randomization

   As in order to perform the blind attacks described in Section 9.2.1
   and Section 9.2.2 the attacker needs to know the TCP port numbers in
   use by the connection to be attacked, obfuscating the TCP source port
   used for outgoing TCP connections will increase the number of packets
   required to successfully perform these attacks.  Section 3.1 of this
   document discusses the use of port randomization.

   It must be noted that given that these blind DupACK triggering
   attacks do not require the attacker to forge valid TCP Sequence
   numbers and TCP Acknowledgement numbers, port randomization should
   not be relied upon as a first line of defense.

   Ingress and Egress filtering

   Ingress and Egress filtering reduces the number of systems in the
   global Internet that can perform attacks that rely on forged source
   IP addresses.  While protection from the blind attacks discussed in
   Section 9.2 should not rely only on Ingress and Egress filtering, its
   deployment is recommended to help prevent all attacks that rely on
   forged IP addresses.  RFC 3704 [Baker and Savola, 2004], RFC 2827
   [Ferguson and Senie, 2000], and [NISCC, 2006] provide advice on
   Ingress and Egress filtering.

   Generalized TTL Security Mechanism (GTSM)

   RFC 5082 [Gill et al, 2007] proposes a check on the TTL field of the
   IP packets that correspond to a given TCP connection to reduce the
   number of systems that could successfully attack the protected TCP
   connection.  It provides for the attacks discussed in this document
   the same level of protection than for the attacks described in
   [Watson, 2004] and RFC 4953 [Touch, 2007].  While implementation of
   this mechanism may be useful in some scenarios, it should be clear
   that countermeasures discussed in the previous sections provide a
   more effective and simpler solution than that provided by the GTSM.







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9.3.  TCP Explicit Congestion Notification (ECN)

   ECN (Explicit Congestion Notification) provides a mechanism for
   intermediate systems to signal congestion to the communicating
   endpoints that in some scenarios can be used as an alternative to
   dropping packets.

   RFC 3168 [Ramakrishnan et al, 2001] contains a detailed discussion of
   the possible ways and scenarios in which ECN could be exploited by an
   attacker.

   RFC 3540 [Spring et al, 2003] specifies an improvement to ECN based
   on nonces, that protects against accidental or malicious concealment
   of marked packets from the TCP sender.  The specified mechanism
   defines a "NS" ("Nonce Sum") field in the TCP header that makes use
   of one bit from the Reserved field, and requires a modification in
   both of the endpoints of a TCP connection to process this new field.
   This mechanism is still in "Experimental" status, and since it might
   suffer from the behavior of some middle-boxes such as firewalls or
   packet-scrubbers, we defer a recommendation of this mechanism until
   more experience is gained.

   There also is ongoing work in the research community and the IETF to
   define alternate semantics for the ECN field of the IP header (e.g.,
   see [PCNWG, 2009]).

   The following subsections try to summarize the security implications
   of ECN.

9.3.1.  Possible attacks by a compromised router

   Firstly, a router controlled by a malicious user could erase the CE
   codepoint (either by replacing it with the ECT(0), ECT(1), or non-ECT
   codepoints), effectively eliminating the congestion indication.  As a
   result, the corresponding TCP sender would not reduce its data
   transmission rate, possibly leading to network congestion.  This
   could also lead to unfairness, as this flow could experience better
   performance than other flows for which the congestion indication is
   not erased (and thus their transmission rate is reduced).

   Secondly, a router controlled by a malicious user could
   illegitimately set the CE codepoint, falsely indicating congestion,
   to cause the TCP sender to reduce its data transmission rate.
   However, this particular attack is no worse than the malicious router
   simply dropping the packets rather setting their CE codepoint.

   Thirdly, a malicious router could turn off the ECT codepoint of a
   packet, thus disabling ECN support.  As a result, if the packet later



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   arrives at a router that is experiencing congestion, it may be
   dropped rather than marked.  As with the previous scenario, though,
   this is no worse than the malicious router simply dropping the
   corresponding packet.

   It should be noted that a compromised on-path IP router could engage
   in a much broader range of attacks, with broader impacts, and at much
   lower attacker cost than the ones described here.  Such a compromised
   router is extremely unlikely to engage in the attack vectors
   discussed in this section, given the existence of more effective
   attack vectors that have lower attacker cost.

9.3.2.  Possible attacks by a malicious TCP endpoint

   If a packet with the ECT codepoint set arrives at an ECN-capable
   router that is experiencing moderate congestion, the router may
   decide to set its CE codepoint instead of dropping it.  If either of
   the TCP endpoints do not honour the congestion indication provided by
   an ECN-capable router, this would result in unfairness, as other
   (legitimate) ECN-capable flows would still reduce their sending rate
   in response to the ECN marking of packets.  Furthermore, under
   moderate congestion, non-ECN-capable flows would be subject to packet
   drops by the same router.  As a result, the flow with a malicious TCP
   end-point would obtain better service than the legitimate flows.

   As noted in RFC 3168 [Ramakrishnan et al, 2001], a TCP endpoint
   falsely indicating ECN capability could lead to unfairness, allowing
   the mis-beheaving flow to get more than its fair share of the
   bandwidth.  This could be the result of the mis-behavior of either of
   the TCP endpoints.  For example, the sending TCP could indicate ECN
   capability, but then send a CWR in response to an ECE without
   actually reducing its congestion window.  Alternatively (or in
   addition), the receiving TCP could simply ignore those packets with
   the CE codepoint set, thus avoiding the sending TCP from receiving
   the congestion indication.

   In the case of the sending TCP ignoring the ECN congestion
   indication, this would be no worse than the sending TCP ignoring the
   congestion indication provided by a lost segment.  However, the case
   of a TCP receiver ignoring the CE codepoint allows the TCP receiver
   to get more than its fair share of bandwidth in a way that was
   previously unavailable.  If congestion was kept "moderate", then the
   malicious TCP receiver could maintain the unfairness, as the router
   experiencing congestion would mark the offending packets of the
   misbehaving flow rather than dropping them.  At the same time,
   legitimate ECN-capable flows would respond to the congestion
   indication provided by the CE codepoint, while legitimate non-ECN-
   capable flows would be subject of packet dropping.  However, if



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   congestion turned to sufficiently heavy, the router experiencing
   congestion would switch from marking packets to dropping packets, and
   at that point the attack vector provided by ECN could no longer be
   exploited (until congestion returns to moderate state).

   RFC 3168 [Ramakrishnan et al, 2001] describes the use of "penalty
   boxes" which would act on flows that do not respond appropriately to
   congestion indications.  Section 10 of RFC 3168 suggests that a first
   action taken at a penalty box for an ECN-capable flow would be to
   switch to dropping packets (instead of marking them), and, if the
   flow does not respond appropriately to the congestion indication, the
   penalty box could reset the misbehaving connection.  Here we
   discourage implementation of such a policy, as it would create a
   vector for connection-reset attacks.  For example, an attacker could
   forge TCP segments with the same four-tuple as the targeted
   connection and cause them to transit the penalty box.  The penalty
   box would first switch from marking to dropping packets.  However,
   the attacker would continue sending forged segments, at a steady
   rate.  As a result, if the penalty box implemented such a severe
   policy of resetting connections for flows that still do not respond
   to end-to-end congestion control after switching from marking to
   dropping, the attacked connection would be reset.


10.  TCP API

   Section 3.8 of RFC 793 [Postel, 1981c] describes the minimum set of
   TCP User Commands required of all TCP Implementations.  Most
   operating systems provide an Application Programming Interface (API)
   that allows applications to make use of the services provided by TCP.
   One of the most popular APIs is the Sockets API, originally
   introduced in the BSD networking package [McKusick et al, 1996].

10.1.  Passive opens and binding sockets

   When there is already a pending passive OPEN for some local port
   number, TCP SHOULD NOT allow processes that do not belong to the same
   user to "reuse" the local port for another passive OPEN.
   Additionally, reuse of a local port SHOULD default to "off", and be
   enabled only by an explicit command (e.g., the setsockopt() function
   of the Sockets API).

   DISCUSSION:

      RFC 793 specifies the syntax of the "OPEN" command, which can be
      used to perform both passive and active opens.  The syntax of this
      command is as follows:




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      OPEN (local port, foreign socket, active/passive [, timeout] [,
      precedence] [, security/compartment] [, options]) -> local
      connection name

      When this command is used to perform a passive open (i.e., the
      active/passive flag is set to passive), the foreign socket
      parameter may be either fully-specified (to wait for a particular
      connection) or unspecified (to wait for any call).

      As discussed in Section 2.7 of RFC 793 [Postel, 1981c], if there
      are several passive OPENs with the same local socket (recorded in
      the corresponding TCB), an incoming connection will be matched to
      the TCB with the more specific foreign socket.  This means that
      when the foreign socket of a passive OPEN matches that of the
      incoming connection request, that passive OPEN takes precedence
      over those passive OPENs with an unspecified foreign socket.

      Popular implementations such as the Sockets API let the user
      specify the local socket as fully-specified {local IP address,
      local TCP port} pair, or as just the local TCP port (leaving the
      local IP address unspecified).  In the former case, only those
      connection requests sent to {local port, local IP address} will be
      accepted.  In the latter case, connection requests sent to any of
      the system's IP addresses will be accepted.  In a similar fashion
      to the generic API described in Section 2.7 of RFC 793, if there
      is a pending passive OPEN with a fully-specified local socket that
      matches that for which a connection establishment request has been
      received, that local socket will take precedence over those which
      have left the local IP address unspecified.  The implication of
      this is that an attacker could "steal" incoming connection
      requests meant for a local application by performing a passive
      OPEN that is more specific than that performed by the legitimate
      application.

10.2.  Active opens and binding sockets

   TCP SHOULD NOT allow port numbers that have been allocated for a TCP
   that is the LISTEN or CLOSED states to be specified as the "local
   port" argument of the "OPEN" command.

   An implementation MAY relax the aforementioned restriction when the
   process or system user requesting allocation of such a port number is
   the same that the process or system user controlling the TCP in the
   CLOSED or LISTEN states with the same port number.

   DISCUSSION:





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      As discussed in Section 10.1, the "OPEN" command specified in
      Section 3.8 of RFC 793 [Postel, 1981c] can be used to perform
      active opens.  In case of active opens, the parameter "local port"
      will contain a so-called "ephemeral port".  While the only
      requirement for such an ephemeral port is that the resulting
      connection-id is unique, port numbers that are currently in use by
      a TCP in the LISTEN state should not be allowed for use as
      ephemeral ports.  If this rule is not complied, an attacker could
      potentially "steal" an incoming connection to a local server
      application by issuing a connection request to the victim client
      at roughly the same time the client tries to connect to the victim
      server application.  If the SYN segment corresponding to the
      attacker's connection request and the SYN segment corresponding to
      the victim client "cross each other in the network", and provided
      the attacker is able to know or guess the ephemeral port used by
      the client, a TCP simultaneous open scenario would take place, and
      the incoming connection request sent by the client would be
      matched with the attacker's socket rather than with the victim
      server application's socket.

      As already noted, in order for this attack to succeed, the
      attacker should be able to guess or know (in advance) the
      ephemeral port selected by the victim client, and be able to know
      the right moment to issue a connection request to the victim
      client.  While in many scenarios this may prove to be a difficult
      task, some factors such as an inadequate ephemeral port selection
      policy at the victim client could make this attack feasible.

      It should be noted that most applications based on popular
      implementations of TCP API (such as the Sockets API) perform
      "passive opens" in three steps.  Firstly, the application obtains
      a file descriptor to be used for inter-process communication
      (e.g., by issuing a socket() call).  Secondly, the application
      binds the file descriptor to a local TCP port number (e.g., by
      issuing a bind() call), thus creating a TCP in the fictional
      CLOSED state.  Thirdly, the aforementioned TCP is put in the
      LISTEN state (e.g., by issuing a listen() call).  As a result,
      with such an implementation of the TCP API, even if port numbers
      in use for TCPs in the LISTEN state were not allowed for use as
      ephemeral ports, there is a window of time between the second and
      the third steps in which an attacker could be allowed to select a
      port number that would be later used for listening to incoming
      connections.  Therefore, these implementations of the TCP API
      should enforce a stricter requirement for the allocation of port
      numbers: port numbers that are in use by a TCP in the LISTEN or
      CLOSED states should not be allowed for allocation as ephemeral
      ports.




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      An implementation might choose to relax the aforementioned
      restriction when the process or system user requesting allocation
      of such a port number is the same that the process or system user
      controlling the TCP in the CLOSED or LISTEN states with the same
      port number.


11.  Blind in-window attacks

   In the last few years awareness has been raised about a number of
   "blind" attacks that can be performed against TCP by forging TCP
   segments that fall within the receive window [NISCC, 2004] [Watson,
   2004].

   The term "blind" refers to the fact that the attacker does not have
   access to the packets that belong to the attacked connection.

   The effects of these attacks range from connection resets to data
   injection.  While these attacks were known in the research community,
   they were generally considered unfeasible.  However, increases in
   bandwidth availability and the use of larger TCP windows raised
   concerns in the community.  The following subsections discuss a
   number of forgery attacks against TCP, along with the possible
   countermeasures to mitigate their impact.

11.1.  Blind TCP-based connection-reset attacks

   Blind connection-reset attacks have the goal of causing a TCP
   connection maintained between two TCP endpoints to be aborted.  The
   level of damage that the attack may cause usually depends on the
   application running on top of TCP, with the more vulnerable
   applications being those that rely on long-lived TCP connections.

   An interesting case of such applications is BGP [Rekhter et al,
   2006], in which a connection-reset usually results in the
   corresponding entries of the routing table being flushed.

   There are a variety of vectors for performing TCP-based connection-
   reset attacks against TCP.  [Watson, 2004] and [NISCC, 2004] raised
   awareness about connection-reset attacks that exploit the RST flag of
   TCP segments.  [Ramaiah et al, 2008] noted that carefully crafted SYN
   segments could also be used to perform connection-reset attacks.
   This document describes yet two previously undocumented vectors for
   performing connection-reset attacks: the Precedence field of IP
   packets that encapsulate TCP segments, and illegal TCP options.






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11.1.1.  RST flag

   TCP SHOULD implement the mitigation for RST-based attacks specified
   in [Ramaiah et al, 2008].

   DISCUSSION:

      The RST flag signals a TCP peer that the connection should be
      aborted.  In contrast with the FIN handshake (which gracefully
      terminates a TCP connection), an RST segment causes the connection
      to be abnormally closed.

      As stated in Section 3.4 of RFC 793 [Postel, 1981c], all reset
      segments are validated by checking their Sequence Numbers, with
      the Sequence Number considered valid if it is within the receive
      window.  In the SYN-SENT state, however, an RST is valid if the
      Acknowledgement Number acknowledges the SYN segment that
      supposedly elicited the reset.

      [Ramaiah et al, 2008] proposes a modification to TCP's transition
      diagram to address this attack vector.  The counter-measure is a
      combination of enforcing a more strict validation check on the
      sequence number of reset segments, and the addition of a
      "challenge" mechanism.  With the implementation of the proposed
      mechanism, TCP would behave as follows:

      If the Sequence Number of an RST segment is outside the receive
      window, the segment is silently dropped (as stated by RFC 793).
      That is, a reset segment is discarded unless it passes the
      following check:

                RCV.NXT <= Sequence Number < RCV.NXT+RCV.WND

      If the sequence number falls exactly on the left-edge of the
      receive window, the reset is honoured.  That is, the connection is
      reset if the following condition is true:

                         Sequence Number == RCV.NXT

      If an RST segment passes the first check (i.e., it is within the
      receive window) but does not pass the second check (i.e., it does
      not fall exactly on the left edge of the receive window), an
      Acknowledgement segment ("challenge ACK") is set in response:

                    <SEQ=SND.NXT><ACK=RCV.NXT><CTL=ACK>






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      This Acknowledgement segment is referred to as a "challenge ACK"
      as, in the event the RST segment that elicited it had been
      legitimate (but silently dropped as a result of enforcing the
      above checks), the challenge ACK would elicit a new reset segment
      that would fall exactly on the left edge of the window and would
      thus pass all the above checks, finally resetting the connection.

      We recommend the implementation of this countermeasure.  However,
      we are aware of patent claims on this counter-measure, and suggest
      vendors to research the consequences of the possible patents that
      may apply.

      [US-CERT, 2003a] is an advisory of a firewall system that was
      found particularly vulnerable to resets attack because of not
      validating the TCP Sequence Number of RST segments.  Clearly, all
      TCPs (including those in middle-boxes) should validate RST
      segments as discussed in this section.

11.1.2.  SYN flag

   Processing of SYN segments received for connections in the
   synchronized states SHOULD occur as follows:

   o  If a SYN segment is received for a connection in any synchronized
      state other than TIME-WAIT, respond with an ACK, applying rate-
      throttling.  [Ramaiah et al, 2008]

   o  If the corresponding connection is in the TIME-WAIT state, then
      process the incomming SYN as specified in
      [I-D.ietf-tcpm-tcp-timestamps].

   DISCUSSION:

      Section 3.9 (page 71) of RFC 793 [Postel, 1981c] states that if a
      SYN segment is received with a valid (i.e., "in window") Sequence
      Number, an RST segment should be sent in response, and the
      connection should be aborted.

      The IETF has published an RFC, "Improving TCP's Resistance to
      Blind In-Window Attacks" [Ramaiah et al, 2008] which addresses,
      among others, this variant of TCP-based connection-reset attack.
      This section describes the counter-measure proposed by the IETF, a
      problem that may arise from the implementation of that solution,
      and a workaround to it.

      In order to mitigate this attack vector, [Ramaiah et al, 2008]
      proposes to change TCP's reaction to SYN segments as follows.
      When a SYN segment is received for a connection in any of the



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      synchronized states, an Acknowledgement (ACK) segment is sent in
      response.

      As discussed in [Ramaiah et al, 2008], there is a corner-case that
      would not be properly handled by this mechanism.  If a host (TCP
      A) establishes a TCP connection with a remote peer (TCP B), and
      then crashes, reboots and tries to initiate a new incarnation of
      the same connection (i.e., a connection with the same four-tuple
      as the previous connection) using an Initial Sequence Number equal
      to the RCV.NXT value at the remote peer (TCP B), the ACK segment
      sent by TCP B in response to the SYN segment would contain an
      Acknowledgement number that would be considered valid by TCP A,
      and thus an RST segment would not be sent in response to the
      Acknowledgement (ACK) segment.  As this ACK would not have the SYN
      bit set, TCP A (being in the SYN-SENT state) would silently drop
      it (as stated on page 68 of RFC 793).  After a Retransmission
      Timeout (RTO), TCP A would retransmit its SYN segment, which would
      lead to the same sequence of events as before.  Eventually, TCP A
      would timeout, and the connection would be aborted.  This is a
      corner case in which the introduced change would lead to a non-
      desirable behavior.  However, we consider this scenario to be
      extremely unlikely and, in the event it ever took place, the
      connection would nevertheless be aborted after retrying for a
      period of USER TIMEOUT seconds.

      However, when this change is implemented exactly as described in
      [Ramaiah et al, 2008], the potential of interoperability problems
      is introduced, as a heuristic widely incorporated in many TCP
      implementations is disabled.

      In a number of scenarios a socket pair may need to be reused while
      the corresponding four-tuple is still in the TIME-WAIT state in a
      remote TCP peer.  For example, a client accessing some service on
      a host may try to create a new incarnation of a previous
      connection, while the corresponding four-tuple is still in the
      TIME-WAIT state at the remote TCP peer (the server).  This may
      happen if the ephemeral port numbers are being reused too quickly,
      either because of a bad policy of selection of ephemeral ports, or
      simply because of a high connection rate to the corresponding
      service.  In such scenarios, the establishment of new connections
      that reuse a four-tuple that is in the TIME-WAIT state would fail.
      In order to avoid this problem, RFC 1122 [Braden, 1989] states (in
      Section 4.2.2.13) that when a connection request is received with
      a four-tuple that is in the TIME-WAIT state, the connection
      request could be accepted if the sequence number of the incoming
      SYN segment is greater than the last sequence number seen on the
      previous incarnation of the connection (for that direction of the
      data transfer).



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      This requirement aims at avoiding the sequence number space of the
      new and old incarnations of the connection to overlap, thus
      avoiding old segments from the previous incarnation of the
      connection to be accepted as valid by the new connection.

      The requirement in [Ramaiah et al, 2008] to disregard SYN segments
      received for connections in any of the synchronized states forbids
      the implementation of the heuristic described above.  As a result,
      we argue that the processing of SYN segments proposed in [Ramaiah
      et al, 2008] should apply only for connections in any of the
      synchronized states other than the TIME-WAIT state.

11.1.3.  Security/Compartment

   If the security/compartment field of an incoming TCP segment does not
   match the value recorded in the corresponding TCB, TCP SHOULD NOT
   abort the connection, but simply discard the corresponding packet.
   Additionally, this whole event SHOULD be logged as a security
   violation.

   DISCUSSION:

      Section 3.9 (page 71) of RFC 793 [Postel, 1981c] states that if
      the IP security/compartment of an incoming segment does not
      exactly match the security/compartment in the TCB, a RST segment
      should be sent, and the connection should be aborted.

      A discussion of the IP security options relevant to this section
      can be found in Section 3.13.2.12, Section 3.13.2.13, and Section
      3.13.2.14 of [CPNI, 2008].

      This certainly provides another attack vector for performing
      connection-reset attacks, as an attacker could forge TCP segments
      with a security/compartment that is different from that recorded
      in the corresponding TCB and, as a result, the attacked connection
      would be reset.

      It is interesting to note that for connections in the ESTABLISHED
      state, this check is performed after validating the TCP Sequence
      Number and checking the RST bit, but before validating the
      Acknowledgement field.  Therefore, even if the stricter validation
      of the Acknowledgement field (described in Section 3.4) was
      implemented, it would not help to mitigate this attack vector.

      This attack vector can be easily mitigated by relaxing the
      reaction to TCP segments with "incorrect" security/compartment
      values as specified in this section.




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11.1.4.  Precedence

   If the Precedence field of an incomming TCP segment does not match
   the value recorded in the corresponding TCB, TCP MUST NOT abort the
   connection, and MUST instead continue processing the segment as
   specified by RFC 793.

   DISCUSSION:

      Section 3.9 (page 71) of RFC 793 [Postel, 1981c] states that if
      the IP Precedence of an incoming segment does not exactly match
      the Precedence recorded in the TCB, a RST segment should be sent,
      and the connection should be aborted.

      This certainly provides another attack vector for performing
      connection-reset attacks, as an attacker could forge TCP segments
      with a IP Precedence that is different from that recorded in the
      corresponding TCB and, as a result, the attacked connection would
      be reset.

      It is interesting to note that for connections in the ESTABLISHED
      state, this check is performed after validating the TCP Sequence
      Number and checking the RST bit, but before validating the
      Acknowledgement field.  Therefore, even if the stricter validation
      of the Acknowledgement field (described in Section 3.4) were
      implemented, it would not help to mitigate this attack vector.

      This attack vector can be easily mitigated by relaxing the
      reaction to TCP segments with "incorrect" IP Precedence values.
      That is, even if the Precedence field does not match the value
      recorded in the corresponding TCB, TCP should not abort the
      connection, and should instead continue processing the segment as
      specified by RFC 793.

      It is interesting to note that resetting a connection due to a
      change in the Precedence value might have a negative impact on
      interoperability.  For example, the packets that correspond to the
      connection could temporarily take a different internet path, in
      which some middle-box could re-mark the Precedence field (due to
      administration policies at the network to be transited).  In such
      a scenario, an implementation following the advice in RFC 793
      would abort the connection, when the connection would have
      probably survived.

      While the IPv4 Type of Service field (and hence the Precedence
      field) has been redefined by the Differentiated Services (DS)
      field specified in RFC 2474 [Nichols et al, 1998], RFC 793
      [Postel, 1981c] was never formally updated in this respect.  We



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      note that both legacy systems that have not been upgraded to
      implement the differentiated services architecture described in
      RFC 2475 [Blake et al, 1998] and current implementations that have
      extrapolated the discussion of the Precedence field to the
      Differentiated Services field may still be vulnerable to the
      connection reset vector discussed in this section.

11.1.5.  Illegal options

   TCP MUST silently drop those TCP segments that contain TCP options
   with illegal option lengths.

   DISCUSSION:

      Section 4.2.2.5 of RFC 1122 [Braden, 1989] discusses the
      processing of TCP options.  It states that TCP must be able to
      receive a TCP option in any segment, and must ignore without error
      any option it does not implement.  Additionally, it states that
      TCP should be prepared to handle an illegal option length (e.g.,
      zero) without crashing, and suggests handling such illegal options
      by resetting the corresponding connection and logging the reason.
      However, this suggested behavior could be exploited to perform
      connection-reset attacks.  Therefore, as discussed in Section 3.10
      of this document, we advise TCP implementations to silently drop
      those TCP segments that contain illegal option lengths.

11.2.  Blind data-injection attacks

   An attacker could try to inject data in the stream of data being
   transferred on the connection.  As with the other attacks described
   in Section 11 of this document, in order to perform a blind data
   injection attack the attacker would need to know or guess the four-
   tuple that identifies the TCP connection to be attacked.
   Additionally, he should be able to guess a valid ("in window") TCP
   Sequence Number, and a valid Acknowledgement Number.

   As discussed in Section 3.4 of this document, [Ramaiah et al, 2008]
   proposes to enforce a more strict check on the Acknowledgement Number
   of incoming segments than that specified in RFC 793 [Postel, 1981c].

   Implementation of the proposed check requires more packets on the
   side of the attacker to successfully perform a blind data-injection
   attack.  However, it should be noted that applications concerned with
   any of the attacks discussed in Section 11 of this document should
   make use of proper authentication techniques, such as those specified
   for IPsec in RFC 4301 [Kent and Seo, 2005].





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12.  Information leaking

12.1.  Remote Operating System detection via TCP/IP stack fingerprinting

   Clearly, remote Operating System (OS) detection is a useful tool for
   attackers.  Tools such as nmap [Fyodor, 2006b] can usually detect the
   operating system type and version of a remote system with an
   amazingly accurate precision.  This information can in turn be used
   by attackers to tailor their exploits to the identified operating
   system type and version.

   Evasion of OS fingerprinting can prove to be a very difficult task.
   Most systems make use of a variety of protocols, each of which have a
   large number of parameters that can be set to arbitrary values.
   Thus, information on the operating system may be obtained from a
   number of sources ranging from application banners to more obscure
   parameters such as TCP's retransmission timer.

   Nmap [Fyodor, 2006b] is probably the most popular tool for remote OS
   detection via active TCP/IP stack fingerprinting. p0f [Zalewski,
   2006a], on the other hand, is a tool for performing remote OS
   detection via passive TCP/IP stack fingerprinting.  SinFP [SinFP,
   2006] can perform both active and passive fingerprinting.  Finally,
   TBIT [TBIT, 2001] is a TCP fingerprinting tool that aims at
   characterizing the behavior of a remote TCP peer based on active
   probes, and which has been widely used in the research community.

   TBIT [TBIT, 2001] implements a number of tests not present in other
   tools, such as characterizing the behavior of a TCP peer with respect
   to TCP congestion control.

   [Fyodor, 1998] and [Fyodor, 2006a] are classic papers on the subject.
   [Miller, 2006] and [Smith and Grundl, 2002] provide an introduction
   to passive TCP/IP stack fingerprinting.  [Smart et al, 2000] and
   [Beck, 2001] discuss some techniques for evading OS detection through
   TCP/IP stack fingerprinting.

   The following subsections discuss TCP-based techniques for remote OS
   detection via and, where possible, propose ways to mitigate them.

12.1.1.  FIN probe

   TCP MUST silently drop TCP any segments received for a connection in
   the LISTEN state that do not have the SYN, RST, or ACK flags set.  In
   the rest of the cases, the processing rules in RFC 793 MUST be
   applied.

   DISCUSSION:



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      The attacker sends a FIN (or any packet without the SYN or the ACK
      flags set) to an open port.  RFC 793 [Postel, 1981c] leaves the
      reaction to such segments unspecified.  As a result, some
      implementations silently drop the received segment, while others
      respond with a RST.

12.1.2.  Bogus flag test

   TCP MUST ignore any flags not supported, and MUST NOT reflect them if
   a TCP segment is sent in response to the one just received.

   DISCUSSION:

      The attacker sends a TCP segment setting at least one bit of the
      Reserved field.  Some implementations ignore this field, while
      others reset the corresponding connection or reflect the field in
      the TCP segment sent in response.

12.1.3.  TCP ISN sampling

   The attacker samples a number of Initial Sequence Numbers by sending
   a number of connection requests.  Many TCP implementations differ on
   the ISN generator they implement, thus allowing the correlation of
   ISN generation algorithm to the operating system type and version.

   This document advises implementing an ISN generator that follows the
   behavior described in RFC 1948 [Bellovin, 1996].  However, it should
   be noted that even if all TCP implementations generated their ISNs as
   proposed in RFC 1948, there is still a number of implementation
   details that are left unspecified, which would allow remote OS
   fingerprinting by means of ISN sampling.  For example, the time-
   dependent parameter of the hash could have a different frequency in
   different TCP implementations.

12.1.4.  TCP initial window

   Many TCP implementations differ on the initial TCP window they use.
   There are a number of factors that should be considered when
   selecting the TCP window to be used for a given system.  A number of
   implementations that use static windows (i.e., no automatic buffer
   tuning mechanisms are implemented) default to a window of around 32
   KB, which seems sensible for the general case.  On the other hand, a
   window of 4 KB seems to be common practice for connections servicing
   critical applications such as BGP.  It is clear that the window size
   is a tradeoff among a number of considerations.  Section 3.7
   discusses some of the considerations that should be made when
   selecting the window size for a TCP connection.




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   If automatic tuning mechanisms are implemented, we suggest the
   initial window to be at least 4 * RMSS segments.  We note that a
   remote OS fingerprinting tool could still sample the advertised TCP
   window, trying to correlate the advertised window with the potential
   automatic buffer tuning algorithm and Operating System.

12.1.5.  RST sampling

   If an RST must be sent in response to an incoming segment, then if
   the ACK bit of an incoming TCP segment is off, a Sequence Number of
   zero MUST be used in the RST segment sent in response.  That is,

                 <SEQ=0><ACK=SEG.SEQ+SEG.LEN><CTL=RST, ACK>

   It should be noted that the SEG.LEN value used for the
   Acknowledgement Number MUST be incremented once for each flag set in
   the original segment that makes use of a byte of the sequence number
   space.  That is, if only one of the SYN or FIN flags were set in the
   received segment, the Acknowledgement Number of the response should
   be set to SEG.SEQ+SEG.LEN+1.  If both the SYN and FIN flags were set
   in the received segment, the Acknowledgement Number should be set to
   SEG.SEQ+SEG.LEN+2.

   We also RECOMMEND that TCP sets ACK bit (and the Acknowledgement
   Number) in all outgoing RST segments, as it allows for additional
   validation checks to be enforced at the system receiving the segment.

   DISCUSSION:

      [Fyodor, 1998] reports that many implementations differ in the
      Acknowledgement Number they use in response to segments received
      for connections in the CLOSED state.  In particular, these
      implementations differ in the way they construct the RST segment
      that is sent in response to those TCP segments received for
      connections in the CLOSED state.

      RFC 793 [Postel, 1981c] describes (in pages 36-37) how RST
      segments are to be generated.  According to this RFC, the ACK bit
      (and the Acknowledgment Number) is set in a RST only if the
      incoming segment that elicited the RST did not have the ACK bit
      set (and thus the Sequence Number of the outgoing RST segment must
      be set to zero).  However, we recommend TCP implementations to set
      the ACK bit (and the Acknowledgement Number) in all outgoing RST
      segments, as it allows for additional validation checks to be
      enforced at the system receiving the segment.






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12.1.6.  TCP options

   Different implementations differ in the TCP options they enable by
   default.  Additionally, they differ in the actual contents of the
   options, and in the order in which the options are included in a TCP
   segment.  There is currently no recommendation on the order in which
   to include TCP options in TCP segments.

12.1.7.  Retransmission Timeout (RTO) sampling

   TCP uses a retransmission timer for retransmitting data in the
   absence of any feedback from the remote data receiver.  The duration
   of this timer is referred to as "retransmission timeout" (RTO).  RFC
   2988 [Paxson and Allman, 2000] specifies the algorithm for computing
   the TCP retransmission timeout (RTO).

   The algorithm allows the use of clocks of different granularities, to
   accommodate the different granularities used by the existing
   implementations.  Thus, the difference in the resulting RTO can be
   used for remote OS fingerprinting.  [Veysset et al, 2002] describes
   how to perform remote OS fingerprinting by sampling and analyzing the
   RTO of the target system.  However, this fingerprinting technique has
   at least the following drawbacks:

   o  It is usually much slower than other fingerprinting techniques, as
      it may require considerable time to sample the RTO of a given
      target.

   o  It is less reliable than other fingerprinting techniques, as
      latency and packet loss can lead to bogus results.

   While in principle it would be possible to defeat this fingerprinting
   technique (e.g., by obfuscating the granularity of the clock used for
   computing the RTO), we consider that a more important step to defeat
   remote OS detection is for implementations to address the more
   effective fingerprinting techniques described in Sections 12.1.1
   through 12.1.7 of this document.

12.2.  System uptime detection

   The "uptime" of a system may prove to be valuable information to an
   attacker.  For example, it might reveal the last time a security
   patch was applied.  Information about system uptime is usually leaked
   by TCP header fields or options that are (or may be) time-dependent,
   and are usually initialized to zero when the system is bootstrapped.
   As a result, if the attacker knows the frequency with which the
   corresponding parameter or header field is incremented, and is able
   to sample the current value of that parameter or header field, the



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   system uptime will be easily obtained.  Two fields that can
   potentially reveal the system uptime is the Sequence Number field of
   a SYN or SYN/ACK segment (i.e., when it contains an ISN) and the
   TSval field of the timestamp option.  Section 3.3.1 of this document
   discusses the generation of TCP Initial Sequence Numbers.  Section
   4.7.1 of this document discusses the generation of TCP timestamps.


13.  Covert channels

   As virtually every communications protocol, TCP can be exploited to
   establish covert channels.  While an exhaustive discussion of covert
   channels is out of the scope of this document, for completeness of
   the document we simply note that it is possible for a (probably
   malicious) user to establish a covert channel by means of TCP, such
   that data can be surreptitiously passed to a remote system, probably
   unnoticed by a monitoring system, and with the possibility of
   concealing the location of the source system.

   In most cases, covert channels based on manipulation of TCP fields
   can be eliminated by protocol scrubbers and other middle-boxes.  On
   the other hand, "timing channels" may prove to be more difficult to
   eliminate.

   [Rowland, 1996] contains a discussion of covert channels in the
   TCP/IP protocol suite, with some TCP-based examples.  [Giffin et al,
   2002] describes the use of TCP timestamps for the establishment of
   covert channels.  [Zander, 2008] contains an extensive bibliography
   of papers on covert channels, and a list of freely-available tools
   that implement covert channels with the TCP/IP protocol suite.


14.  TCP Port scanning

   TCP port scanning aims at identifying TCP port numbers on which there
   is a process listening for incoming connections.  That is, it aims at
   identifying TCPs at the target system that are in the LISTEN state.
   The following subsections describe different TCP port scanning
   techniques that have been implemented in freely-available tools.
   These subsections focus only on those port scanning techniques that
   exploit features of TCP itself, and not of other communication
   protocols.

   For example, the following subsections do not discuss the
   exploitation of application protocols (such as FTP) or the
   exploitation of features of underlying protocols (such as the IP
   Identification field) for port-scanning purposes.




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14.1.  Traditional connect() scan

   The most trivial scanning technique consists in trying to perform the
   TCP three-way handshake with each of the port numbers at the target
   system (e.g. by issuing a call to the connect() function of the
   Sockets API).  The three-way handshake will complete for port numbers
   that are "open", but will fail for those port numbers that are
   "closed".

   As this port-scanning technique can be implemented by issuing a call
   to the connect() function of the Sockets API that normal applications
   use, it does not require the attacker to have superuser privileges.
   The downside of this port-scanning technique is that it is less
   efficient than other scanning methods (e.g., the "SYN scan" described
   in Section 14.2), and that it can be easily logged by the target
   system.

14.2.  SYN scan

   The SYN scan was introduced as a "stealth" port-scanning technique.
   It aims at avoiding the target system from logging the port scan by
   not completing the TCP three-way handshake.  When a SYN/ACK segment
   is received in response to the initial SYN segment, the system
   performing the port scan will respond with an RST segment, thus
   preventing the three-way handshake from completing.  While this port-
   scanning technique is harder to detect and log than the traditional
   connect() scan described in Section 14.1, most current NIDS (Network
   Intrusion Detection Systems) can detect and log it.

   SYN scans are sometimes mistakenly reported as "SYN flood" attacks by
   NIDS, though.

   The main advantage of this port scanning technique is that it is much
   more efficient than the traditional connect() scan.

   In order to implement this port-scanning technique, port-scanning
   tools usually bypass the TCP API, and forge the SYN segments they
   send (e.g., by using raw sockets).  This typically requires the
   attacker to have superuser privileges to be able to run the port-
   scanning tool.

14.3.  FIN, NULL, and XMAS scans

   TCP SHOULD respond with an RST when a TCP segment is received for a
   connection in the LISTEN state, and the incoming segment has neither
   the SYN bit nor the RST bit set.

   DISCUSSION:



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      RFC 793 [Postel, 1981c] states, in page 65, that an incoming
      segment that does not have the RST bit set and that is received
      for a connection in the fictional state CLOSED causes an RST to be
      sent in response.  Pages 65-66 of RFC 793 describes the processing
      of incoming segments for connections in the state LISTEN, and
      implicitly states that an incoming segment that does not have the
      ACK bit set (and is not a SYN or an RST) should be silently
      dropped.

      As a result, an attacker can exploit this situation to perform a
      port scan by sending TCP segments that do not have the ACK bit set
      to the target system.  When a port is "open" (i.e., there is a TCP
      in the LISTEN state on the corresponding port), the target system
      will respond with an RST segment.  On the other hand, if the port
      is "closed" (i.e., there is a TCP in the fictional state CLOSED)
      the attacker will not get any response from the target system.

      Since the only requirement for exploiting this port scanning
      vector is that the probe segments must not have the ACK bit set,
      there are a number of different TCP control-bits combinations that
      can be used for the probe segments.

      When the probe segment sent to the target system is a TCP segment
      that has only the FIN bit set, the scanning technique is usually
      referred to as a "FIN scan".  When the probe packet is a TCP
      segment that does not have any of the control bits set, the
      scanning technique is usually known as a "NULL scan".  Finally,
      when the probe packet sent to the target system has only the FIN,
      PSH, and the URG bits set, the port-scanning technique is known as
      a "XMAS scan".

      It should be clear that while the aforementioned control-bits
      combinations are the most popular ones, other combinations could
      be used to exploit this port-scanning vector.  For example, the
      CWR, ECE, and/or any of the Reserved bits could be set in the
      probe segments.

      The advantage of this port-scanning technique is that in can
      bypass some stateless firewalls.  However, the downside is that a
      number of implementations do not comply strictly with RFC 793
      [Postel, 1981c], and thus always respond to the probe segments
      with an RST, regardless of whether the port is open or closed.

      This port-scanning vector can be easily defeated as rby responding
      with an RST when a TCP segment is received for a connection in the
      LISTEN state, and the incoming segment has neither the SYN bit nor
      the RST bit set.




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14.4.  Maimon scan

   If a TCP that is in the CLOSED or LISTEN states receives a TCP
   segment with both the FIN and ACK bits set, it MUST respond with a
   RST.

   DISCUSSION:

      This port scanning technique was introduced in [Maimon, 1996] with
      the name "StealthScan" (method #1), and was later incorporated
      into the nmap tool [Fyodor, 2006b] as the "Maimon scan".

      This port scanning technique employs TCP segments that have both
      the FIN and ACK bits sets as the probe segments.  While according
      to RFC 793 [Postel, 1981c] these segments should elicit an RST
      regardless of whether the corresponding port is open or closed, a
      programming flaw found in a number of TCP implementations has
      caused some systems to silently drop the probe segment if the
      corresponding port was open (i.e., there was a TCP in the LISTEN
      state), and respond with an RST only if the port was closed.

      Therefore, an RST would indicate that the scanned port is closed,
      while the absence of a response from the target system would
      indicate that the scanned port is open.

      While this bug has not been found in current implementations of
      TCP, it might still be present in some legacy systems.

14.5.  Window scan

   When sending an RST segment, TCP SHOULD set the Window field to zero.

   DISCUSSION:

      This port-scanning technique employs ACK segments as the probe
      packets.  ACK segments will elicit an RST from the target system
      regardless of whether the corresponding TCP port is open or
      closed.  However, as described in [Maimon, 1996], some systems set
      the Window field of the RST segments with different values
      depending on whether the corresponding TCP port is open or closed.
      These systems set the Window field of their RST segments to zero
      when the corresponding TCP port is closed, and set the Window
      field to a non-zero value when the corresponding TCP port is open.

      As a result, an attacker could exploit this situation for
      performing a port scan by sending ACK segments to the target
      system, and examining the Window field of the RST segments that
      his probe segments elicit.



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      In order to defeat this port-scanning technique, we recommend TCP
      implementations to set the Window field to zero in all the RST
      segments they send.  Most popular implementations of TCP already
      implement this policy.

14.6.  ACK scan

   The so-called "ACK scan" is not really a port-scanning technique
   (i.e., it does not aim at determining whether a specific port is open
   or closed), but rather aims at determining whether some intermediate
   system is filtering TCP segments sent to that specific port number.

   The probe packet is a TCP segment with the ACK bit set which,
   according to RFC 793 [Postel, 1981c] should elicit an RST from the
   target system regardless of whether the corresponding TCP port is
   open or closed.  If no response is received from the target system,
   it is assumed that some intermediate system is filtering the probe
   packets sent to the target system.

   It should be noted that this "port scanning" techniques exploits
   basic TCP processing rules, and therefore cannot be defeated at an
   end-system.


15.  Processing of ICMP error messages by TCP

   TCP SHOULD silently ignore received ICMP Source Quench messages.

   TCP SHOULD process ICMP "hard errors" as "soft errors" when they are
   received for connections that are in any of he synchronized states.

   TCP SHOULD process ICMP "fragmentation needed and DF bit set" and
   ICMPv6 "Packet Too Big" error messages as described in [RFC5927].

   DISCUSSION:

      [RFC5927] analyzes a number of vulnerabilities based on crafted
      ICMP messages, along with possible counter-measures.


16.  TCP interaction with the Internet Protocol (IP)

16.1.  TCP-based traceroute

   The traceroute tool is used to identify the intermediate systems the
   local system and the destination system.  It is usually implemented
   by sending "probe" packets with increasing IP Time to Live values
   (starting from 0), without maintaining any state with the final



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   destination.

   Some traceroute implementations use ICMP "echo request" messages as
   the probe packets, while others use UDP packets or TCP SYN segments.

   In some cases, the state-less nature of the traceroute tool may
   prevent it from working correctly across stateful devices such as
   Network Address Translators (NATs) or firewalls.

   In order to by-pass this limitation, an attacker could establish a
   TCP connection with the destination system, and start sending TCP
   segments on that connection with increasing IP Time to Live values
   (starting from 0) [Zalewski, 2007] [Zalewski, 2008].  Provided ICMP
   error messages are not blocked by any intermediate system, an
   attacker could exploit this technique to map the network topology
   behind the aforementioned stateful devices in scenarios in which he
   could not have achieved this goal using the traditional traceroute
   tool.

   NATs [Srisuresh and Egevang, 2001] and other middle-boxes could
   defeat this network-mapping technique by overwriting the Time to Live
   of the packets they forward to the internal network.  For example,
   they could overwrite the Time to Live of all packets being forwarded
   to an internal network with a value such as 128.  We strongly
   recommend against overwriting the IP Time to Live field with the
   value 255 or other similar large values, as this could allow an
   attacker to bypass the protection provided by the Generalized TTL
   Security Mechanism (GTSM) described in RFC 5087 [Gill et al, 2007].

   [Gont and Srisuresh, 2008] discusses the security implications of
   NATs, and proposes mitigations for this and other issues.

16.2.  Blind TCP data injection through fragmented IP traffic

   As discussed in Section 11.2, TCP data injection attacks usually
   require an attacker to guess or know a number of parameters related
   with the target TCP connection, such as the connection-id {Source
   Address, Source Port, Destination Address, Destination Port}, the TCP
   Sequence Number, and the TCP Acknowledgement Number.  Provided these
   values are obfuscated as recommended in this document, the chances of
   an off-path attacker of successfully performing a data injection
   attack against a TCP connection are fairly low for many of the most
   common scenarios.

   As discussed in this document, randomization of the values contained
   in different TCP header fields is not a replacement for cryptographic
   methods for protecting a TCP connection, such as IPsec (specified in
   RFC 4301 [Kent and Seo, 2005]).



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   However, [Zalewski, 2003b] describes a possible vector for performing
   a TCP data injection attack that does not require the attacker to
   guess or know the aforementioned TCP connection parameters, and could
   therefore be successfully exploited in some scenarios with less
   effort than that required to exploit the more traditional data-
   injection attack vectors.

   The attack vector works as follows.  When one system is transferring
   information to a remote peer by means of TCP, and the resulting
   packet gets fragmented, the first fragment will usually contain the
   entire TCP header which, together with the IP header, includes all
   the connection parameters that an attacker would need to guess or
   know to successfully perform a data injection attack against TCP.  If
   an attacker were able to forge all the fragments other than the first
   one, his forged fragments could be reassembled together with the
   legitimate first fragment, and thus he would be relieved from the
   hard task of guessing or knowing connection parameters such as the
   TCP Sequence Number and the TCP Acknowledgement Number.

   In order to successfully exploit this attack vector, the attacker
   should be able to guess or know both of the IP addresses involved in
   the target TCP connection, the IP Identification value used for the
   specific packet he is targeting, and the TCP Checksum of that target
   packet.  While it would seem that these values are hard to guess, in
   some specific scenarios, and with some security-unwise implementation
   approaches for the TCP and IP protocols, these values may be feasible
   to guess or know.  For example, if the sending system uses
   predictable IP Identification values, the attacker could simply
   perform a brute force attack, trying each of the possible
   combinations for the TCP Checksum field.  In more specific scenarios,
   the attacker could have more detailed knowledge about the data being
   transferred over the target TCP connection, which might allow him to
   predict the TCP Checksum of the target packet.  For example, if both
   of the involved TCP peers used predictable values for the TCP
   Sequence Number and for the IP Identification fields, and the
   attacker knew the data being transferred over the target TCP
   connection, he could be able to carefully forge the IP payload of his
   IP fragments so that the checksum of the reassembled TCP segment
   matched the Checksum included in the TCP header of the first (and
   legitimate) IP fragment.

   As discussed in Section 4.1 of [CPNI, 2008], IP fragmentation
   provides a vector for performing a variety of attacks against an IP
   implementation.  Therefore, we discourage the reliance on IP
   fragmentation by end-systems, and recommend the implementation of
   mechanisms for the discovery of the Path-MTU, such as that described
   in Section 15.7.3 of this document and/or that described in RFC 4821
   [Mathis and Heffner, 2007].  We nevertheless recommend randomization



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   of the IP Identification field as described in Section 3.5.2 of
   [CPNI, 2008].  While randomization of the IP Identification field
   does not eliminate this attack vector, it does require more work on
   the side of the attacker to successfully exploit it.

16.3.  Broadcast and multicast IP addresses

   TCP connection state is maintained between only two endpoints at a
   time.  As a result, broadcast and multicast IP addresses should not
   be allowed for the establishment of TCP connections.  Section 4.3 of
   [CPNI, 2008] provides advice about which specific IP address blocks
   should not be allowed for connection-oriented protocols such as TCP.


17.  Security Considerations

   This document provides a thorough security assessment of the
   Transmission Control Protocol (TCP), identifies a number of
   vulnerabilities, and specifies possible counter-measures.
   Additionally, it provides implementation guidance such that the
   resilience of TCP implementations is improved.


18.  Acknowledgements

   The author would like to thank (in alphabetical order) David Borman,
   Wesley Eddy, and Alfred Hoenes, for providing valuable feedback on
   earlier versions of thi document.

   This document is heavily based on the document "Security Assessment
   of the Transmission Control Protocol (TCP)" [CPNI, 2009] written by
   Fernando Gont on behalf of CPNI (Centre for the Protection of
   National Infrastructure).

   The author would like to thank (in alphabetical order) Randall
   Atkinson, Guillermo Gont, Alfred Hoenes, Jamshid Mahdavi, Stanislav
   Shalunov, Michael Welzl, Dan Wing, Andrew Yourtchenko, Michal
   Zalewski, and Christos Zoulas, for providing valuable feedback on
   earlier versions of the UK CPNI document.

   Additionally, the author would like to thank (in alphabetical order)
   Mark Allman, David Black, Ethan Blanton, David Borman, James Chacon,
   John Heffner, Jerrold Leichter, Jamshid Mahdavi, Keith Scott, Bill
   Squier, and David White, who generously answered a number of
   questions that araised while the aforementioned document was being
   written.

   Finally, the author would like to thank CPNI (formely NISCC) for



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   their continued support.


19.  References

   Abley, J., Savola, P., Neville-Neil, G. 2007.  Deprecation of Type 0
   Routing Headers in IPv6.  RFC 5095.

   Allman, M. 2003.  TCP Congestion Control with Appropriate Byte
   Counting (ABC).  RFC 3465.

   Allman, M. 2008.  Comments On Selecting Ephemeral Ports.  Available
   at: http://www.icir.org/mallman/share/ports-dec08.pdf

   Allman, M., Paxson, V., Stevens, W. 1999.  TCP Congestion Control.
   RFC 2581.

   Allman, M., Balakrishnan, H., Floyd, S. 2001.  Enhancing TCP's Loss
   Recovery Using Limited Transmit.  RFC 3042.

   Allman, M., Floyd, S., and C. Partridge. 2002.  Increasing TCP's
   Initial Window.  RFC 3390.

   Baker, F. 1995.  Requirements for IP Version 4 Routers.  RFC 1812.

   Baker, F., Savola, P. 2004.  Ingress Filtering for Multihomed
   Networks.  RFC 3704.

   Barisani, A. 2006.  FTester - Firewall and IDS testing tool.
   Available at: http://dev.inversepath.com/trac/ftester

   Beck, R. 2001.  Passive-Aggressive Resistance: OS Fingerprint
   Evasion.  Linux Journal.

   Bellovin, S. M. 1989.  Security Problems in the TCP/IP Protocol
   Suite.  Computer Communication Review, Vol. 19, No. 2, pp. 32-48.

   Bellovin, S. M. 1996.  Defending Against Sequence Number Attacks.
   RFC 1948.

   Bellovin, S. M. 2006.  Towards a TCP Security Option.  IETF Internet-
   Draft (draft-bellovin-tcpsec-00.txt), work in progress.

   Bernstein, D. J. 1996.  SYN cookies.  Available at:
   http://cr.yp.to/syncookies.html

   Blake, S., Black, D., Carlson, M., Davies, E., Wang, Z., and Weiss,
   W., 1998.  An Architecture for Differentiated Services.  RFC 2475.



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   Blanton, E., Allman, M., Fall, K., Wang, L. 2003.  A Conservative
   Selective Acknowledgment (SACK)-based Loss Recovery Algorithm for
   TCP.  RFC 3517.

   Borman, D. 1997.  Post to the tcp-impl mailing-list.  Message-Id:
   <199706061526.KAA01535@frantic.BSDI.COM>.  Available at:
   http://www.kohala.com/start/borman.97jun06.txt

   Borman, D., Deering, S., Hinden, R. 1999.  IPv6 Jumbograms.  RFC
   2675.

   Braden, R. 1989.  Requirements for Internet Hosts -- Communication
   Layers.  RFC 1122.

   Braden, R. 1992.  Extending TCP for Transactions -- Concepts.  RFC
   1379.

   Braden, R. 1994.  T/TCP -- TCP Extensions for Transactions Functional
   Specification.  RFC 1644.

   CCSDS. 2006.  Consultative Committee for Space Data Systems (CCSDS)
   Recommendation Communications Protocol Specification (SCPS) --
   Transport Protocol (SCPS-TP).  Blue Book.  Issue 2.  Available at:
   http://public.ccsds.org/publications/archive/714x0b2.pdf

   CERT. 1996.  CERT Advisory CA-1996-21: TCP SYN Flooding and IP
   Spoofing Attacks.  Available at:
   http://www.cert.org/advisories/CA-1996-21.html

   CERT. 1997.  CERT Advisory CA-1997-28 IP Denial-of-Service Attacks.
   Available at: http://www.cert.org/advisories/CA-1997-28.html

   CERT. 2000.  CERT Advisory CA-2000-21: Denial-of-Service
   Vulnerabilities in TCP/IP Stacks.  Available at:
   http://www.cert.org/advisories/CA-2000-21.html

   CERT. 2001.  CERT Advisory CA-2001-09: Statistical Weaknesses in
   TCP/IP Initial Sequence Numbers.  Available at:
   http://www.cert.org/advisories/CA-2001-09.html

   CERT. 2003.  CERT Advisory CA-2003-13 Multiple Vulnerabilities in
   Snort Preprocessors.  Available at:
   http://www.cert.org/advisories/CA-2003-13.html

   Cisco. 2008a.  Cisco Security Appliance Command Reference, Version
   7.0.  Available at: http://www.cisco.com/en/US/docs/security/asa/
   asa70/command/reference/tz.html#wp1288756




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   Cisco. 2008b.  Cisco Security Appliance System Log Messages, Version
   8.0.  Available at: http://www.cisco.com/en/US/docs/security/asa/
   asa80/system/message/logmsgs.html#wp4773952

   Clark, D.D. 1982.  Fault isolation and recovery.  RFC 816.

   Clark, D.D. 1988.  The Design Philosophy of the DARPA Internet
   Protocols, Computer Communication Review, Vol. 18, No.4, pp. 106-114.

   Connolly, T., Amer, P., Conrad, P. 1994.  An Extension to TCP :
   Partial Order Service.  RFC 1693.

   Conta, A., Deering, S., Gupta, M. 2006.  Internet Control Message
   Protocol (ICMPv6) for the Internet Protocol Version 6 (IPv6)
   Specification.  RFC 4443.

   CORE. 2003.  Core Secure Technologies Advisory CORE-2003-0307: Snort
   TCP Stream Reassembly Integer Overflow Vulnerability.  Available at:
   http://www.coresecurity.com/common/showdoc.php?idx=313&idxseccion=10

   CPNI, 2008.  Security Assessment of the Internet Protocol.  Available
   at: http://www.cpni.gov.uk/Docs/InternetProtocol.pdf

   CPNI, 2009.  Security Assessment of the Transmission Control Protocol
   (TCP).  Available at:
   http://www.cpni.gov.uk/Docs/tn-03-09-security-assessment-TCP.pdf

   daemon9, route, and infinity. 1996.  IP-spoofing Demystified (Trust-
   Relationship Exploitation), Phrack Magazine, Volume Seven, Issue
   Forty-Eight, File 14 of 18.  Available at:
   http://www.phrack.org/archives/48/P48-14

   Deering, S., Hinden, R. 1998.  Internet Protocol, Version 6 (IPv6)
   Specification.  RFC 2460.

   Dharmapurikar, S., Paxson, V. 2005.  Robust TCP Stream Reassembly In
   the Presence of Adversaries.  Proceedings of the USENIX Security
   Symposium 2005.

   Duke, M., Braden, R., Eddy, W., Blanton, E. 2006.  A Roadmap for
   Transmission Control Protocol (TCP) Specification Documents.  RFC
   4614.

   Ed3f. 2002.  Firewall spotting and networks analisys with a broken
   CRC.  Phrack Magazine, Volume 0x0b, Issue 0x3c, Phile #0x0c of 0x10.
   Available at: http://www.phrack.org/phrack/60/p60-0x0c.txt

   Eddy, W. 2007.  TCP SYN Flooding Attacks and Common Mitigations.  RFC



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   4987.

   Fenner, B. 2006.  Experimental Values in IPv4, IPv6, ICMPv4, ICMPv6,
   UDP, and TCP Headers.  RFC 4727.

   Ferguson, P., and Senie, D. 2000.  Network Ingress Filtering:
   Defeating Denial of Service Attacks which employ IP Source Address
   Spoofing.  RFC 2827.

   Fielding, R., Gettys, J., Mogul, J., Frystyk, H., Masinter, L.,
   Leach, P., and Berners-Lee, T. 1999.  Hypertext Transfer Protocol --
   HTTP/1.1.  RFC 2616.

   Floyd, S., Mahdavi, J., Mathis, M., Podolsky, M. 2000.  An Extension
   to the Selective Acknowledgement (SACK) Option for TCP.  RFC 2883.

   Floyd, S., Henderson, T., Gurtov, A. 2004.  The NewReno Modification
   to TCP's Fast Recovery Algorithm.  RFC 3782.

   Floyd, S., Allman, M., Jain, A., Sarolahti, P. 2007.  Quick-Start for
   TCP and IP.  RFC 4782.

   Fyodor. 1998.  Remote OS Detection via TCP/IP Stack Fingerprinting.
   Phrack Magazine, Volume 8, Issue, 54.

   Fyodor. 2006a.  Remote OS Detection via TCP/IP Fingerprinting (2nd
   Generation).  Available at: http://insecure.org/nmap/osdetect/.

   Fyodor. 2006b.  Nmap - Free Security Scanner For Network Exploration
   and Audit.  Available at: http://www.insecure.org/nmap.

   Fyodor. 2008.  Nmap Reference Guide: Port Scanning Techniques.
   Available at: http://nmap.org/book/man-port-scanning-techniques.html

   GIAC. 2000.  Egress Filtering v 0.2.  Available at:
   http://www.sans.org/y2k/egress.htm

   Giffin, J., Greenstadt, R., Litwack, P., Tibbetts, R. 2002.  Covert
   Messaging through TCP Timestamps.  PET2002 (Workshop on Privacy
   Enhancing Technologies), San Francisco, CA, USA, April2002.
   Available at:
   http://web.mit.edu/greenie/Public/CovertMessaginginTCP.ps

   Gill, V., Heasley, J., Meyer, D., Savola, P, Pignataro, C. 2007.  The
   Generalized TTL Security Mechanism (GTSM).  RFC 5082.

   Gont, F. 2006.  Advanced ICMP packet filtering.  Available at:
   http://www.gont.com.ar/papers/icmp-filtering.html



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   Gont, F. 2008a.  ICMP attacks against TCP.  IETF Internet-Draft
   (draft-ietf-tcpm-icmp-attacks-04.txt), work in progress.

   Gont, F.. 2008b.  TCP's Reaction to Soft Errors.  IETF Internet-Draft
   (draft-ietf-tcpm-tcp-soft-errors-09.txt), work in progress.

   Gont, F. 2009.  On the generation of TCP timestamps.  IETF Internet-
   Draft (draft-gont-tcpm-tcp-timestamps-01.txt), work in progress.

   Gont, F., Srisuresh, P. 2008.  Security Implications of Network
   Address Translators (NATs).  IETF Internet-Draft
   (draft-gont-behave-nat-security-01.txt), work in progress.

   Gont, F., Yourtchenko, A. 2009.  On the implementation of TCP urgent
   data.  IETF Internet-Draft (draft-gont-tcpm-urgent-data-01.txt), work
   in progress.

   Heffernan, A. 1998.  Protection of BGP Sessions via the TCP MD5
   Signature Option.  RFC 2385.

   Heffner, J. 2002.  High Bandwidth TCP Queuing.  Senior Thesis.

   Hnes, A. 2007.  TCP options - tcp-parameters IANA registry.  Post to
   the tcpm wg mailing-list.  Available at:
   http://www.ietf.org/mail-archive/web/tcpm/current/msg03199.html

   IANA. 2007.  Transmission Control Protocol (TCP) Option Numbers.
   Avialable at: http://www.iana.org/assignments/tcp-parameters/

   IANA. 2008.  Port Numbers.  Available at:
   http://www.iana.org/assignments/port-numbers

   Jacobson, V. 1988.  Congestion Avoidance and Control.  Computer
   Communication Review, vol. 18, no. 4, pp. 314-329.  Available at:
   ftp://ftp.ee.lbl.gov/papers/congavoid.ps.Z

   Jacobson, V., Braden, R. 1988.  TCP Extensions for Long-Delay Paths.
   RFC 1072.

   Jacobson, V., Braden, R., Borman, D. 1992.  TCP Extensions for High
   Performance.  RFC 1323.

   Jones, S. 2003.  Port 0 OS Fingerprinting.  Available at:
   http://www.gont.com.ar/docs/port-0-os-fingerprinting.txt

   Kent, S. and Seo, K. 2005.  Security Architecture for the Internet
   Protocol.  RFC 4301.




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   Klensin, J. 2008.  Simple Mail Transfer Protocol.  RFC 5321.

   Ko, Y., Ko, S., and Ko, M. 2001.  NIDS Evasion Method named SeolMa.
   Phrack Magazine, Volume 0x0b, Issue 0x39, phile #0x03 of 0x12.
   Available at: http://www.phrack.org/issues.html?issue=57&id=3#article

   Lahey, K. 2000.  TCP Problems with Path MTU Discovery.  RFC 2923.

   Larsen, M., Gont, F. 2008.  Port Randomization.  IETF Internet-Draft
   (draft-ietf-tsvwg-port-randomization-02), work in progress.

   Lemon, 2002.  Resisting SYN flood DoS attacks with a SYN cache.
   Proceedings of the BSDCon 2002 Conference, pp 89-98.

   Maimon, U. 1996.  Port Scanning without the SYN flag.  Phrack
   Magazine, Volume Seven, Issue Fourty-Nine, phile #0x0f of 0x10.
   Available at:
   http://www.phrack.org/issues.html?issue=49&id=15#article

   Mathis, M., Mahdavi, J., Floyd, S. Romanow, A. 1996.  TCP Selective
   Acknowledgment Options.  RFC 2018.

   Mathis, M., and Heffner, J. 2007.  Packetization Layer Path MTU
   Discovery.  RFC 4821.

   McCann, J., Deering, S., Mogul, J. 1996.  Path MTU Discovery for IP
   version 6.  RFC 1981.

   McKusick, M., Bostic, K., Karels, M., and J. Quarterman. 1996.  The
   Design and Implementation of the 4.4BSD Operating System.  Addison-
   Wesley.

   Meltman. 1997. new TCP/IP bug in win95.  Post to the bugtraq mailing-
   list.  Available at: http://insecure.org/sploits/land.ip.DOS.html

   Miller, T. 2006.  Passive OS Fingerprinting: Details and Techniques.
   Available at: http://www.ouah.org/incosfingerp.htm .

   Mogul, J., and Deering, S. 1990.  Path MTU Discovery.  RFC 1191.

   Morris, R. 1985.  A Weakness in the 4.2BSD Unix TCP/IP Software.
   Technical Report CSTR-117, AT&T Bell Laboratories.  Available at:
   http://pdos.csail.mit.edu/~rtm/papers/117.pdf .

   Myst. 1997.  Windows 95/NT DoS.  Post to the bugtraq mailing-list.
   Available at: http://seclists.org/bugtraq/1997/May/0039.html

   Nichols, K., Blake, S., Baker, F., and Black, D. 1998.  Definition of



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   the Differentiated Services Field (DS Field) in the IPv4 and IPv6
   Headers.  RFC 2474.

   NISCC. 2004.  NISCC Vulnerability Advisory 236929: Vulnerability
   Issues in TCP.  Available at:
   http://www.uniras.gov.uk/niscc/docs/re-20040420-00391.pdf

   NISCC. 2005.  NISCC Vulnerability Advisory 532967/NISCC/ICMP:
   Vulnerability Issues in ICMP packets with TCP payloads.  Available
   at: http://www.niscc.gov.uk/niscc/docs/re-20050412-00303.pdf

   NISCC. 2006.  NISCC Technical Note 01/2006: Egress and Ingress
   Filtering.  Available at:
   http://www.niscc.gov.uk/niscc/docs/re-20060420-00294.pdf?lang=en

   Ostermann, S. 2008. tcptrace tool.  Tool and documentation available
   at: http://www.tcptrace.org.

   Paxson, V., Allman, M. 2000.  Computing TCP's Retransmission Timer.
   RFC 2988.

   PCNWG. 2009.  Congestion and Pre-Congestion Notification (pcn)
   charter.  Available at:
   http://www.ietf.org/html.charters/pcn-charter.html

   PMTUDWG. 2007.  Path MTU Discovery (pmtud) charter.  Available at:
   http://www.ietf.org/html.charters/OLD/pmtud-charter.html

   Postel, J. 1981a.  Internet Protocol.  DARPA Internet Program.
   Protocol Specification.  RFC 791.

   Postel, J. 1981b.  Internet Control Message Protocol.  RFC 792.

   Postel, J. 1981c.  Transmission Control Protocol.  DARPA Internet
   Program.  Protocol Specification.  RFC 793.

   Postel, J. 1987.  TCP AND IP BAKE OFF.  RFC 1025.

   Ptacek, T. H., and Newsham, T. N. 1998.  Insertion, Evasion and
   Denial of Service: Eluding Network Intrusion Detection.  Secure
   Networks, Inc. Available at:
   http://www.aciri.org/vern/Ptacek-Newsham-Evasion-98.ps

   Ramaiah, A., Stewart, R., and Dalal, M. 2008.  Improving TCP's
   Robustness to Blind In-Window Attacks.  IETF Internet-Draft
   (draft-ietf-tcpm-tcpsecure-10.txt), work in progress.

   Ramakrishnan, K., Floyd, S., and Black, D. 2001.  The Addition of



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   Explicit Congestion Notification (ECN) to IP.  RFC 3168.

   Rekhter, Y., Li, T., Hares, S. 2006.  A Border Gateway Protocol 4
   (BGP-4).  RFC 4271.

   Rivest, R. 1992.  The MD5 Message-Digest Algorithm.  RFC 1321.

   Rowland, C. 1997.  Covert Channels in the TCP/IP Protocol Suite.
   First Monday Journal, Volume 2, Number 5.  Available at:
   http://www.firstmonday.org/issues/issue2_5/rowland/

   Savage, S., Cardwell, N., Wetherall, D., Anderson, T. 1999.  TCP
   Congestion Control with a Misbehaving Receiver.  ACM Computer
   Communication Review, 29(5), October 1999.

   Semke, J., Mahdavi, J., Mathis, M. 1998.  Automatic TCP Buffer
   Tuning.  ACM Computer Communication Review, Vol. 28, No. 4.

   Shalunov, S. 2000.  Netkill.  Available at:
   http://www.internet2.edu/~shalunov/netkill/netkill.html

   Shimomura, T. 1995.  Technical details of the attack described by
   Markoff in NYT.  Message posted in USENETs comp.security.misc
   newsgroup, Message-ID: <3g5gkl$5j1@ariel.sdsc.edu>.  Available at:
   http://www.gont.com.ar/docs/post-shimomura-usenet.txt.

   Silbersack, M. 2005.  Improving TCP/IP security through randomization
   without sacrificing interoperability.  EuroBSDCon 2005 Conference.

   SinFP. 2006.  Net::SinFP - a Perl module to do OS fingerprinting.
   Available at:
   http://www.gomor.org/cgi-bin/index.pl?mode=view;page=sinfp

   Smart, M., Malan, G., Jahanian, F. 2000.  Defeating TCP/IP Stack
   Fingerprinting.  Proceedings of the 9th USENIX Security Symposium,
   pp. 229-240.  Available at: http://www.usenix.org/publications/
   library/proceedings/sec2000/full_papers/smart/smart_html/index.html

   Smith, C., Grundl, P. 2002.  Know Your Enemy: Passive Fingerprinting.
   The Honeynet Project.

   Spring, N., Wetherall, D., Ely, D. 2003.  Robust Explicit Congestion
   Notification (ECN) Signaling with Nonces.  RFC 3540.

   Srisuresh, P., Egevang, K. 2001.  Traditional IP Network Address
   Translator (Traditional NAT).  RFC 3022.

   Stevens, W. R. 1994.  TCP/IP Illustrated, Volume 1: The Protocols.



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   Addison-Wesley Professional Computing Series.

   TBIT. 2001.  TBIT, the TCP Behavior Inference Tool.  Available at:
   http://www.icir.org/tbit/

   Touch, J. 2007.  Defending TCP Against Spoofing Attacks.  RFC 4953.

   US-CERT. 2001.  US-CERT Vulnerability Note VU#498440: Multiple TCP/IP
   implementations may use statistically predictable initial sequence
   numbers.  Available at: http://www.kb.cert.org/vuls/id/498440

   US-CERT. 2003a.  US-CERT Vulnerability Note VU#26825: Cisco Secure
   PIX Firewall TCP Reset Vulnerability.  Available at:
   http://www.kb.cert.org/vuls/id/26825

   US-CERT. 2003b.  US-CERT Vulnerability Note VU#464113: TCP/IP
   implementations handle unusual flag combinations inconsistently.
   Available at: http://www.kb.cert.org/vuls/id/464113

   US-CERT. 2004a.  US-CERT Vulnerability Note VU#395670: FreeBSD fails
   to limit number of TCP segments held in reassembly queue.  Available
   at: http://www.kb.cert.org/vuls/id/395670

   US-CERT. 2005a.  US-CERT Vulnerability Note VU#102014: Optimistic TCP
   acknowledgements can cause denial of service.  Available at:
   http://www.kb.cert.org/vuls/id/102014

   US-CERT. 2005b.  US-CERT Vulnerability Note VU#396645: Microsoft
   Windows vulnerable to DoS via LAND attack.  Available at:
   http://www.kb.cert.org/vuls/id/396645

   US-CERT. 2005c.  US-CERT Vulnerability Note VU#637934: TCP does not
   adequately validate segments before updating timestamp value.
   Available at: http://www.kb.cert.org/vuls/id/637934

   US-CERT. 2005d.  US-CERT Vulnerability Note VU#853540: Cisco PIX
   fails to verify TCP checksum.  Available at:
   http://www.kb.cert.org/vuls/id/853540.

   Veysset, F., Courtay, O., Heen, O. 2002.  New Tool And Technique For
   Remote Operating System Fingerprinting.  Intranode Research Team.

   Watson, P. 2004.  Slipping in the Window: TCP Reset Attacks,
   CanSecWest 2004 Conference.

   Welzl, M. 2008.  Internet congestion control: evolution and current
   open issues.  CAIA guest talk, Swinburne University, Melbourne,
   Australia.  Available at:



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   http://www.welzl.at/research/publications/caia-jan08.pdf

   Wright, G. and W. Stevens. 1994.  TCP/IP Illustrated, Volume 2: The
   Implementation.  Addison-Wesley.

   Zalewski, M. 2001a.  Strange Attractors and TCP/IP Sequence Number
   Analysis.  Available at:
   http://lcamtuf.coredump.cx/oldtcp/tcpseq.html

   Zalewski, M. 2001b.  Delivering Signals for Fun and Profit.
   Available at: http://lcamtuf.coredump.cx/signals.txt

   Zalewski, M. 2002.  Strange Attractors and TCP/IP Sequence Number
   Analysis - One Year Later.  Available at:
   http://lcamtuf.coredump.cx/newtcp/

   Zalewski, M. 2003a.  Windows URG mystery solved!  Post to the bugtraq
   mailing-list.  Available at:
   http://lcamtuf.coredump.cx/p0f-help/p0f/doc/win-memleak.txt

   Zalewski, M. 2003b.  A new TCP/IP blind data injection technique?
   Post to the bugtraq mailing-list.  Available at:
   http://lcamtuf.coredump.cx/ipfrag.txt

   Zalewski, M. 2006a. p0f passive fingerprinting tool.  Available at:
   http://lcamtuf.coredump.cx/p0f.shtml

   Zalewski, M. 2006b. p0f - RST+ signatures.  Available at:
   http://lcamtuf.coredump.cx/p0f-help/p0f/p0fr.fp

   Zalewski, M. 2007. 0trace - traceroute on established connections.
   Post to the bugtraq mailing-list.  Available at:
   http://seclists.org/bugtraq/2007/Jan/0176.html

   Zalewski, M. 2008.  Museum of broken packets.  Available at:
   http://lcamtuf.coredump.cx/mobp/

   Zander, S. 2008.  Covert Channels in Computer Networks.  Available
   at: http://caia.swin.edu.au/cv/szander/cc/index.html

   Zquete, A. 2002.  Improving the functionality of SYN cookies. 6th
   IFIP Communications and Multimedia Security Conference (CMS 2002).
   Available at: http://www.ieeta.pt/~avz/pubs/CMS02.html

   Zweig, J., Partridge, C. 1990.  TCP Alternate Checksum Options.  RFC
   1146.





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20.  References

20.1.  Normative References

   [I-D.ietf-tcpm-tcp-timestamps]
              Gont, F., "Reducing the TIME-WAIT state using TCP
              timestamps", draft-ietf-tcpm-tcp-timestamps-03 (work in
              progress), December 2010.

   [I-D.ietf-tsvwg-port-randomization]
              Larsen, M. and F. Gont, "Transport Protocol Port
              Randomization Recommendations",
              draft-ietf-tsvwg-port-randomization-09 (work in progress),
              August 2010.

   [RFC6093]  Gont, F. and A. Yourtchenko, "On the Implementation of the
              TCP Urgent Mechanism", RFC 6093, January 2011.

20.2.  Informative References

   [I-D.gont-timestamps-generation]
              Gont, F. and A. Oppermann, "On the generation of TCP
              timestamps", draft-gont-timestamps-generation-00 (work in
              progress), June 2010.

   [RFC5927]  Gont, F., "ICMP Attacks against TCP", RFC 5927, July 2010.


Appendix A.  TODO list

   A Number of formatting issues still have to be fixed in this
   document.  Among others are:

   o  The ASCII-art corresponding to some figures are still missing.  We
      still have to convert the nice JPGs of the UK CPNI document into
      ugly ASCII-art.

   o  The references have not yet been converted to xml, but are
      hardcoded, instead.  That's why they may not look as expected


Appendix B.  Change log (to be removed by the RFC Editor before
             publication of this document as an RFC)

B.1.  Changes from draft-ietf-tcpm-tcp-security-01

   A Number of formatting issues still have to be fixed in this
   document.  Among others are:



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   o  The whole document was reformatted with RFC 1122 style.


Author's Address

   Fernando Gont
   UK Centre for the Protection of National Infrastructure

   Email: fernando@gont.com.ar
   URI:   http://www.cpni.gov.uk









































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