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TCP Maintenance and Minor Extensions                             F. Gont
(tcpm)                                                           UK CPNI
Internet-Draft                                            March 13, 2012
Intended status: Informational
Expires: September 14, 2012


 Survey of Security Hardening Methods for Transmission Control Protocol
                         (TCP) Implementations
                  draft-ietf-tcpm-tcp-security-03.txt

Abstract

   This document surveys methods to harden Transmission Control Protocol
   (TCP) implementations.  It provides an overview of known attacks and
   refers to the corresponding solutions in the TCP standards.

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 September 14, 2012.

Copyright Notice

   Copyright (c) 2012 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
   described in the Simplified BSD License.



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Table of Contents

   1.  Preface  . . . . . . . . . . . . . . . . . . . . . . . . . . .  5
     1.1.  Introduction . . . . . . . . . . . . . . . . . . . . . . .  5
     1.2.  Scope of this document . . . . . . . . . . . . . . . . . .  6
     1.3.  Organization of this document  . . . . . . . . . . . . . .  7
   2.  The Transmission Control Protocol  . . . . . . . . . . . . . .  7
   3.  TCP header fields  . . . . . . . . . . . . . . . . . . . . . .  8
     3.1.  Source Port and Destination Port . . . . . . . . . . . . .  8
     3.2.  Sequence number  . . . . . . . . . . . . . . . . . . . . .  9
     3.3.  Acknowledgement Number . . . . . . . . . . . . . . . . . . 10
     3.4.  Data Offset  . . . . . . . . . . . . . . . . . . . . . . . 10
     3.5.  Control bits . . . . . . . . . . . . . . . . . . . . . . . 10
       3.5.1.  Reserved (four bits) . . . . . . . . . . . . . . . . . 10
       3.5.2.  CWR (Congestion Window Reduced)  . . . . . . . . . . . 11
       3.5.3.  ECE (ECN-Echo) . . . . . . . . . . . . . . . . . . . . 11
       3.5.4.  URG  . . . . . . . . . . . . . . . . . . . . . . . . . 11
       3.5.5.  ACK  . . . . . . . . . . . . . . . . . . . . . . . . . 12
       3.5.6.  PSH  . . . . . . . . . . . . . . . . . . . . . . . . . 12
       3.5.7.  RST  . . . . . . . . . . . . . . . . . . . . . . . . . 12
       3.5.8.  SYN  . . . . . . . . . . . . . . . . . . . . . . . . . 12
       3.5.9.  FIN  . . . . . . . . . . . . . . . . . . . . . . . . . 12
     3.6.  Window . . . . . . . . . . . . . . . . . . . . . . . . . . 13
       3.6.1.  Security implications arising from closed windows  . . 14
     3.7.  Checksum . . . . . . . . . . . . . . . . . . . . . . . . . 14
     3.8.  Urgent pointer . . . . . . . . . . . . . . . . . . . . . . 16
     3.9.  Options  . . . . . . . . . . . . . . . . . . . . . . . . . 16
     3.10. Padding  . . . . . . . . . . . . . . . . . . . . . . . . . 19
     3.11. Data . . . . . . . . . . . . . . . . . . . . . . . . . . . 19
   4.  Common TCP Options . . . . . . . . . . . . . . . . . . . . . . 19
     4.1.  End of Option List (Kind = 0)  . . . . . . . . . . . . . . 19
     4.2.  No Operation (Kind = 1)  . . . . . . . . . . . . . . . . . 19
     4.3.  Maximum Segment Size (Kind = 2)  . . . . . . . . . . . . . 19
     4.4.  Selective Acknowledgement Option . . . . . . . . . . . . . 20
       4.4.1.  SACK-permitted Option (Kind = 4) . . . . . . . . . . . 20
       4.4.2.  SACK Option (Kind = 5) . . . . . . . . . . . . . . . . 20
     4.5.  MD5 Option (Kind=19) . . . . . . . . . . . . . . . . . . . 21
     4.6.  Window scale option (Kind = 3) . . . . . . . . . . . . . . 21
     4.7.  Timestamps option (Kind = 8) . . . . . . . . . . . . . . . 22
       4.7.1.  Generation of timestamps . . . . . . . . . . . . . . . 22
       4.7.2.  Vulnerabilities  . . . . . . . . . . . . . . . . . . . 22
   5.  Connection-establishment mechanism . . . . . . . . . . . . . . 24
     5.1.  SYN flood  . . . . . . . . . . . . . . . . . . . . . . . . 24
     5.2.  Connection forgery . . . . . . . . . . . . . . . . . . . . 28
     5.3.  Connection-flooding attack . . . . . . . . . . . . . . . . 29
       5.3.1.  Vulnerability  . . . . . . . . . . . . . . . . . . . . 29
       5.3.2.  Countermeasures  . . . . . . . . . . . . . . . . . . . 30
     5.4.  Firewall-bypassing techniques  . . . . . . . . . . . . . . 32



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   6.  Connection-termination mechanism . . . . . . . . . . . . . . . 32
     6.1.  FIN-WAIT-2 flooding attack . . . . . . . . . . . . . . . . 32
       6.1.1.  Vulnerability  . . . . . . . . . . . . . . . . . . . . 32
       6.1.2.  Countermeasures  . . . . . . . . . . . . . . . . . . . 33
   7.  Buffer management  . . . . . . . . . . . . . . . . . . . . . . 35
     7.1.  TCP retransmission buffer  . . . . . . . . . . . . . . . . 36
       7.1.1.  Vulnerability  . . . . . . . . . . . . . . . . . . . . 36
       7.1.2.  Countermeasures  . . . . . . . . . . . . . . . . . . . 37
     7.2.  TCP segment reassembly buffer  . . . . . . . . . . . . . . 40
     7.3.  Automatic buffer tuning mechanisms . . . . . . . . . . . . 42
       7.3.1.  Automatic send-buffer tuning mechanisms  . . . . . . . 43
       7.3.2.  Automatic receive-buffer tuning mechanism  . . . . . . 45
   8.  TCP segment reassembly algorithm . . . . . . . . . . . . . . . 47
     8.1.  Problems that arise from ambiguity in the reassembly
           process  . . . . . . . . . . . . . . . . . . . . . . . . . 47
   9.  TCP Congestion Control . . . . . . . . . . . . . . . . . . . . 48
     9.1.  Congestion control with misbehaving receivers  . . . . . . 48
       9.1.1.  ACK division . . . . . . . . . . . . . . . . . . . . . 48
       9.1.2.  DupACK forgery . . . . . . . . . . . . . . . . . . . . 49
       9.1.3.  Optimistic ACKing  . . . . . . . . . . . . . . . . . . 49
     9.2.  Blind DupACK triggering attacks against TCP  . . . . . . . 50
       9.2.1.  Blind throughput-reduction attack  . . . . . . . . . . 52
       9.2.2.  Blind flooding attack  . . . . . . . . . . . . . . . . 53
       9.2.3.  Difficulty in performing the attacks . . . . . . . . . 53
       9.2.4.  Modifications to TCP's loss recovery algorithms  . . . 54
       9.2.5.  Countermeasures  . . . . . . . . . . . . . . . . . . . 55
     9.3.  TCP Explicit Congestion Notification (ECN) . . . . . . . . 55
   10. TCP API  . . . . . . . . . . . . . . . . . . . . . . . . . . . 56
     10.1. Passive opens and binding sockets  . . . . . . . . . . . . 56
     10.2. Active opens and binding sockets . . . . . . . . . . . . . 57
   11. Blind in-window attacks  . . . . . . . . . . . . . . . . . . . 59
     11.1. Blind TCP-based connection-reset attacks . . . . . . . . . 59
       11.1.1. RST flag . . . . . . . . . . . . . . . . . . . . . . . 60
       11.1.2. SYN flag . . . . . . . . . . . . . . . . . . . . . . . 60
       11.1.3. Security/Compartment . . . . . . . . . . . . . . . . . 60
       11.1.4. Precedence . . . . . . . . . . . . . . . . . . . . . . 61
       11.1.5. Illegal options  . . . . . . . . . . . . . . . . . . . 61
     11.2. Blind data-injection attacks . . . . . . . . . . . . . . . 61
   12. Information leaking  . . . . . . . . . . . . . . . . . . . . . 62
     12.1. Remote Operating System detection via TCP/IP stack
           fingerprinting . . . . . . . . . . . . . . . . . . . . . . 62
       12.1.1. FIN probe  . . . . . . . . . . . . . . . . . . . . . . 63
       12.1.2. Bogus flag test  . . . . . . . . . . . . . . . . . . . 63
       12.1.3. TCP ISN sampling . . . . . . . . . . . . . . . . . . . 63
       12.1.4. TCP initial window . . . . . . . . . . . . . . . . . . 63
       12.1.5. RST sampling . . . . . . . . . . . . . . . . . . . . . 64
       12.1.6. TCP options  . . . . . . . . . . . . . . . . . . . . . 65
       12.1.7. Retransmission Timeout (RTO) sampling  . . . . . . . . 65



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     12.2. System uptime detection  . . . . . . . . . . . . . . . . . 66
   13. Covert channels  . . . . . . . . . . . . . . . . . . . . . . . 66
   14. TCP Port scanning  . . . . . . . . . . . . . . . . . . . . . . 66
     14.1. Traditional connect() scan . . . . . . . . . . . . . . . . 67
     14.2. SYN scan . . . . . . . . . . . . . . . . . . . . . . . . . 67
     14.3. FIN, NULL, and XMAS scans  . . . . . . . . . . . . . . . . 68
     14.4. Maimon scan  . . . . . . . . . . . . . . . . . . . . . . . 69
     14.5. Window scan  . . . . . . . . . . . . . . . . . . . . . . . 69
     14.6. ACK scan . . . . . . . . . . . . . . . . . . . . . . . . . 70
   15. Processing of ICMP error messages by TCP . . . . . . . . . . . 70
   16. TCP interaction with the Internet Protocol (IP)  . . . . . . . 70
     16.1. TCP-based traceroute . . . . . . . . . . . . . . . . . . . 71
     16.2. Blind TCP data injection through fragmented IP traffic . . 71
     16.3. Broadcast and multicast IP addresses . . . . . . . . . . . 73
   17. Security Considerations  . . . . . . . . . . . . . . . . . . . 73
   18. Acknowledgements . . . . . . . . . . . . . . . . . . . . . . . 73
   19. References (to be translated to xml) . . . . . . . . . . . . . 74
   20. References . . . . . . . . . . . . . . . . . . . . . . . . . . 84
     20.1. Normative References . . . . . . . . . . . . . . . . . . . 84
     20.2. Informative References . . . . . . . . . . . . . . . . . . 84
   Appendix A.  TODO list . . . . . . . . . . . . . . . . . . . . . . 85
   Appendix B.  Change log (to be removed by the RFC Editor
                before publication of this document as an RFC)  . . . 85
     B.1.  Changes from draft-ietf-tcpm-tcp-security-02 . . . . . . . 85
     B.2.  Changes from draft-ietf-tcpm-tcp-security-01 . . . . . . . 86
   Author's Address . . . . . . . . . . . . . . . . . . . . . . . . . 86

























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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.

   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 described.  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 is 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 machanisms described in the following documents were selected for
   assessment as part of this work:

   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)



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   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)

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.

   The core TCP specification, RFC 793 [RFC0793], dates back to 1981 and



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   standardizes the basic mechanisms and policies of TCP.  RFC 1122
   [RFC1122] provides clarifications and errata for the original
   specification.  RFC 2581 [RFC5681] 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 [RFC0793] defines the syntax of a TCP segment, along with the
   semantics of each of the header fields.

   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.

   Most active attacks against ongoing TCP connections require the
   attacker to guess or know the four-tuple that identifies the
   connection.  As a result, randomization of the TCP ephemeral ports
   provides a (partial) mitigation against off-path attacks.  [RFC6056]
   provides guidance in this area.

   Some implementations have been known to crash when a TCP segment in



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   which the source end-point (IP Source Address, TCP Source Port) is
   the same as the destination end-point (IP Destination Address, TCP
   Destination Port). [draft-gont-tcpm-tcp-mirrored-endpoints-00.txt]
   describes this issue in detail and provides advice in this area.

   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.

      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.

      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

   Predictable sequence numbers allow 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 [Morris1985],
   and its exploitation was widely publicized about 10 years later
   [Shimomura1995].

   In order to mitigate this vulnerabilities, some implementations set
   the TCP ISN to a PRNG.  However, this has been known to cause
   interoperability problems.  [RFC6528] provides advice in this area.

   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 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".  [Gont and
   Srisuresh, 2008] provides a detailed discussion of the security
   implications of NATs and of the possible mitigations for this and
   other issues.






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3.3.  Acknowledgement Number

   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) [RFC5961] has proposed to
   enforce a more strict check on the Acknowledgement Number of segments
   that have the ACK bit set.  See for more details.

   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.

   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

   [draft-gont-tcpm-tcp-sanity-checks-00.txt] specifies a number of
   sanity checks that should be performed on the Data Offset field.

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 [RFC1025] [RFC1379].  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)

   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



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   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)

   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)

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

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

3.5.4.  URG

   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.




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3.5.5.  ACK

   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

   [draft-gont-tcpm-tcp-push-semantics-00.txt] describes a number of
   security issues that may arise as a result of the PUSH semantics, and
   proposes a number of ways to mitigate these issues.

3.5.7.  RST

   The RST bit is used to request the abortion (abnormal close) of a TCP
   connection.  RFC 793 [RFC0793] 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], [RFC6429] proposed
   stricter validity checks.  Please see [RFC6429] for additional
   details.

   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

   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

   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



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   (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

   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.

   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.

   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.




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   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.

3.6.1.  Security implications arising from closed windows

   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").

      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

   While in principle there should not be security implications arising
   from the Checksum field, due to non-RFC-compliant implementations,



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   the Checksum can be exploited to detect firewalls, evade network
   intrusion detection systems (NIDS), and/or perform Denial of Service
   attacks.

   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.






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3.8.  Urgent pointer

   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.

   [draft-gont-tcpm-tcp-sanity-checks-00.txt] describes a number of
   sanity checks to be enforced on TCP segments regarding urgent
   indications.  [RFC6093] deprecates the use of urgent indications in
   new applications.

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".

   [draft-gont-tcpm-tcp-sanity-checks-00.txt] describes a number of
   sanity checks that should be performed on TCP options.

   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.


4.  Common TCP Options

4.1.  End of Option List (Kind = 0)

   This option indicates the "End of Options".  As noted in
   [draft-gont-tcpm-tcp-sanity-checks-00.txt], some implementations pad
   the end of options with "No Operation" options rather than including
   an "End of Options List" option.

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 MSS option has been employed for performing DoS attacks, by
   advertising very small MSS values thus greatly increasing the packet-



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   rate used by the victim system.
   [draft-gont-tcpm-tcp-sanity-checks-00.txt] describes this issue, and
   proposes sanity checks to mitigate it.

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)

   [draft-gont-tcpm-tcp-sanity-checks-00.txt] to be performed on this
   option.

4.4.2.  SACK Option (Kind = 5)

   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.

   [draft-gont-tcpm-tcp-sanity-checks-00.txt] describes sanity checks to



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   be performed on this option such that this and other possible issues
   are mitigated.

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.

   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.

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).

   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



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   this document.  Additionally, large windows can exacerbate the impact
   of resource exhaustion attacks such as those described in Section 7
   of this document.

   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).

4.7.1.  Generation of timestamps

   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.

   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.
   [I-D.gont-timestamps-generation] discusses the generation of TCP
   timestamps in detail.

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



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   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.

   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



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   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.

5.1.  SYN flood

   TCP uses a mechanism known as the "three-way handshake" for the
   establishment of a connection between two TCP peers.  RFC 793
   [RFC0793] 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 [RFC0793].

   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].

   Given that the attacker does not need to complete the three-way



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   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 [RFC0793].

   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.

   [Borman, 1997] and RFC 4987 [Eddy, 2007] contain a general discussion



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   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:

   o  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).

   o  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.

   o  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.

   o  Fourthly, in those scenarios in which establishment of new
      connections is blocked by simply dropping segments with the SYN



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      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 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



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   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:

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

   o  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
   the use of IP source routing.  By forging the Source Address of the
   IP packets that encapsulate the TCP segments of a connection, and



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   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

   NOTE: THIS SECTION IS BEING EDITED.  RFC2119-LANGUAGE IS BEING
   REMOVED.

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

   [draft-gont-tcpm-tcp-sanity-checks-00.txt] discusses how packets with
   both the SYN and RST bits set have been employed in the wild to
   bypass firewall rules, and provides advices in this area.


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.




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   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 [RFC0793] 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
   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



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      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).

   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



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   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
   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







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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
      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



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   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

   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.




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   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.

   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



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   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

   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:






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   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 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.

   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



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   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 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



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   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 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

   NOTE: THIS SECTION IS BEING EDITED.  PLEASE DISREGARD THE RFC2119-
   LANGUAGE RECOMMENDATIONS.



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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
      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



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      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,
      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.





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      *  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)

   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






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   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.

      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.





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      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

   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 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 [RFC0793], if a TCP segment
   arrives containing some data bytes that have already been received,



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   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

   NOTE: THIS SECTION IS BEING EDITED.

   TCP implements two algorithms, "slow start" and "congestion
   avoidance", for controlling the rate at which data is transmitted on
   a TCP connection [RFC5681].

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

   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,



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   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

   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



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   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
   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 5681 [RFC5681], 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



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   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 [RFC0793] describes the processing of 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



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   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
   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.



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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
   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)

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.

   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.




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   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
   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.

   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



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   Section 9.2.2 is performed.

   SACK-based loss recovery

   [I-D.ietf-tcpm-3517bis] 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 [RFC5681].  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
   that implement SACK-based loss recovery will not be vulnerable to the
   blind flooding attack described in Section 9.2.2.  Additionally, as
   [I-D.ietf-tcpm-3517bis] requires 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 will not be vulnerable to the blind throughput-
   reduction attack described in Section 9.2.1.

9.2.5.  Countermeasures

   [draft-gont-tcpm-limiting-aow-segments-00.txt] proposes to rate-limit
   the reaction to out-of-window segments.  This would mitigate the
   attacks described earlier in this section.

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



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   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]).

   RFC 3168 [RFC3168] provides a very throrough security assessment of
   ECN.  Among the possible mitigations, it 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

   NOTE: THIS SECTION IS BEING EDITED.

   Section 3.8 of RFC 793 [RFC0793] 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:




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      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:

      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 [RFC0793], 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.



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   DISCUSSION:

      As discussed in Section 10.1, the "OPEN" command specified in
      Section 3.8 of RFC 793 [RFC0793] 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

   NOTE: THIS SECTION IS BEING EDITED.

   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

   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 [RFC0793], 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.

   [RFC5961] 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.

      We note that 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

   Section 3.9 (page 71) of RFC 793 [RFC0793] 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.  This could be leveraged to perform a blind connection-
   reset attack.  [RFC5961] proposes a change in TCP's state diagram to
   mitigate this attack vector.

11.1.3.  Security/Compartment

   Section 3.9 (page 71) of RFC 793 [RFC0793] 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.  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.



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   [draft-gont-tcpm-tcp-seccomp-prec-00.txt] aims to update RFC 793 such
   that this issue is eliminated.

11.1.4.  Precedence

   Section 3.9 (page 71) of RFC 793 [RFC0793] states that if the IP
   precedence of an incoming segment does not exactly match the
   precedence 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 precedence that is different from that
   recorded in the corresponding TCB and, as a result, the attacked
   connection would be reset.

   [draft-gont-tcpm-tcp-seccomp-prec-00.txt] aims to update RFC 793 such
   that this issue is eliminated.

11.1.5.  Illegal options

   Section 4.2.2.5 of RFC 1122 [RFC1122] discusses the processing of TCP
   options.  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.

   [draft-gont-tcpm-tcp-illegal-option-lengths-00] aims at formally
   updating RFC 1122, such that this issue is eliminated.

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 [RFC0793].

   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



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   for IPsec in RFC 4301 [Kent and Seo, 2005].


12.  Information leaking

   NOTE: THIS SECTION IS BEING EDITED.

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.







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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:

      The attacker sends a FIN (or any packet without the SYN or the ACK
      flags set) to an open port.  RFC 793 [RFC0793] 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



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   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.

   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 [RFC0793] 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



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      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.

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.







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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
   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

   NOTE: THIS SECTION IS BEING EDITED.

   TCP port scanning aims at identifying TCP port numbers on which there
   is a process listening for incoming connections.  That is, it aims at



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   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.

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



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   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:

      RFC 793 [RFC0793] 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



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      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
      [RFC0793], 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.

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 [RFC0793] 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:





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      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.

      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 [RFC0793] 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

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


16.  TCP interaction with the Internet Protocol (IP)





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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
   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



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   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]).

   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.




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   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
   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, Alfred Hoenes, and Michael Scharf, 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)



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   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
   their continued support.


19.  References (to be translated to xml)

   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-



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   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.

   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




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   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

   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



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   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
   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.



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   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

   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.




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   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.

   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.

   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:



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   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
   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



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   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
   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.



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   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.
   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.



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   Welzl, M. 2008.  Internet congestion control: evolution and current
   open issues.  CAIA guest talk, Swinburne University, Melbourne,
   Australia.  Available at:
   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



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   1146.


20.  References

20.1.  Normative References

   [RFC0793]  Postel, J., "Transmission Control Protocol", STD 7,
              RFC 793, September 1981.

   [RFC1122]  Braden, R., "Requirements for Internet Hosts -
              Communication Layers", STD 3, RFC 1122, October 1989.

   [RFC3168]  Ramakrishnan, K., Floyd, S., and D. Black, "The Addition
              of Explicit Congestion Notification (ECN) to IP",
              RFC 3168, September 2001.

   [RFC5681]  Allman, M., Paxson, V., and E. Blanton, "TCP Congestion
              Control", RFC 5681, September 2009.

   [RFC5961]  Ramaiah, A., Stewart, R., and M. Dalal, "Improving TCP's
              Robustness to Blind In-Window Attacks", RFC 5961,
              August 2010.

   [RFC6056]  Larsen, M. and F. Gont, "Recommendations for Transport-
              Protocol Port Randomization", BCP 156, RFC 6056,
              January 2011.

   [RFC6093]  Gont, F. and A. Yourtchenko, "On the Implementation of the
              TCP Urgent Mechanism", RFC 6093, January 2011.

   [RFC6191]  Gont, F., "Reducing the TIME-WAIT State Using TCP
              Timestamps", BCP 159, RFC 6191, April 2011.

   [RFC6528]  Gont, F. and S. Bellovin, "Defending against Sequence
              Number Attacks", RFC 6528, February 2012.

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.

   [I-D.ietf-tcpm-3517bis]
              Blanton, E., Jarvinen, I., Wang, L., Allman, M., Kojo, M.,
              and Y. Nishida, "A Conservative Selective Acknowledgment
              (SACK)-based Loss Recovery Algorithm for TCP",



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              draft-ietf-tcpm-3517bis-01 (work in progress),
              January 2012.

   [Morris1985]
              Morris, R., "A Weakness in the 4.2BSD UNIX TCP/IP
              Software", CSTR 117, AT&T Bell Laboratories, Murray Hill,
              NJ, 1985.

   [RFC1025]  Postel, J., "TCP and IP bake off", RFC 1025,
              September 1987.

   [RFC1379]  Braden, B., "Extending TCP for Transactions -- Concepts",
              RFC 1379, November 1992.

   [RFC5927]  Gont, F., "ICMP Attacks against TCP", RFC 5927, July 2010.

   [RFC6429]  Bashyam, M., Jethanandani, M., and A. Ramaiah, "TCP Sender
              Clarification for Persist Condition", RFC 6429,
              December 2011.

   [Shimomura1995]
              Shimomura, T., "Technical details of the attack described
              by Markoff in NYT",
               http://www.gont.com.ar/docs/post-shimomura-usenet.txt,
              Message posted in USENET's comp.security.misc newsgroup,
              Message-ID: <3g5gkl$5j1@ariel.sdsc.edu>, 1995.


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-02






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   o  Lots of text has been removed out of the document.

   o  The documento track has been changed from BCP to Informational
      (RFC2119-language recommendations ahve been removed).

   o  Where necessary, stand-alone std tracks documents have been
      produced.

B.2.  Changes from draft-ietf-tcpm-tcp-security-01

   A Number of formatting issues still have to be fixed in this
   document.  Among others are:

   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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