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12 13 14 15 RFC 3819
Internet Engineering Task Force Phil Karn, editor
INTERNET DRAFT Carsten Bormann
Gorry Fairhurst
Aaron Falk
Dan Grossman
Reiner Ludwig
Jamshid Mahdavi
Saverio Mascolo
Marie-Jose Montpetit
Gabriel Montenegro
Joe Touch
Lloyd Wood
File: draft-ietf-pilc-link-design-06.txt July, 2001
Expires: January, 2002
Advice for Internet Subnetwork Designers
Status of this Memo
This document is an Internet-Draft and is in full conformance with
all provisions of Section 10 of RFC2026.
Internet-Drafts are working documents of the Internet Engineering
Task Force (IETF), its areas, and its working groups. Note that
other groups may also distribute working documents as Internet-
Drafts.
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."
The list of current Internet-Drafts can be accessed at
http://www.ietf.org/ietf/1id-abstracts.txt
The list of Internet-Draft Shadow Directories can be accessed at
http://www.ietf.org/shadow.html.
Abstract
This document provides advice to the designers of digital
communication equipment, link layer protocols and packet switched
subnetworks (collectively referred to as subnetworks) who wish to
support the Internet protocols but who may be unfamiliar with
Internet architecture and the implications of their design choices on
the performance and efficiency of the Internet.
This document represents a consensus of the members of the IETF
Performance Implications of Link Characteristics (PILC) working
group.
Introduction and Overview
IP, the Internet Protocol [RFC791] is the core protocol of the
Internet. IP defines a simple "connectionless" packet-switched
network. The success of the Internet is largely attributed to IP's
simplicity, the "end-to-end principle" [SRC81] on which the Internet
is based, and the resulting ease of carrying IP on a wide variety of
subnetworks not necessarily designed with IP in mind.
But while many subnetworks carry IP, they do not necessarily do so
with maximum efficiency, minimum complexity or minimum cost. Nor do
they implement certain features to efficiently support newer Internet
features of increasing importance, such as multicasting or quality of
service.
With the explosive growth of the Internet, IP is an increasingly
large fraction of the traffic carried by the world's
telecommunications networks. It therefore makes sense to optimize
both existing and new subnetwork technologies for IP as much as
possible.
Optimizing a subnetwork for IP involves three complementary
considerations:
1. Providing functionality sufficient to carry IP.
2. Eliminating unnecessary functions that increase cost or
complexity.
3. Choosing subnetwork parameters that maximize the performance of
the Internet protocols.
Because IP is so simple, consideration 2 is more of an issue than
consideration 1. I.e., subnetwork designers make many more errors of
commission than errors of omission. But certain enhanced Internet
features, such as multicasting and quality-of-service, rely on
support from the underlying subnetworks beyond that necessary to
carry "traditional" unicast, best-effort IP.
A major consideration in the efficient design of any layered
communication network is the appropriate layer(s) in which to
implement a given function. This issue was first addressed in the
seminal paper "End-to-End Arguments in System Design" [SRC81]. This
paper argued that many functions can be implemented properly *only*
on an end-to-end basis, i.e., at the higher protocol layers, outside
the subnetwork. These functions include ensuring the reliable
delivery of data and the use of cryptography to provide
confidentiality and message integrity.
These functions cannot be provided solely by the concatenation of
hop-by-hop services, so duplicating these functions at the lower
protocol layers (i.e., within the subnetwork) can be needlessly
redundant or even harmful to cost and performance.
However, partial duplication of functionality in a lower layer can
*sometimes* be justified by performance, security or availability
considerations. Examples include link layer retransmission to improve
the performance of an unusually lossy channel, e.g., mobile radio;
link level encryption intended to thwart traffic analysis; and
redundant transmission links to improve availability, increase
throughput, or to guarantee performance for certain classes of
traffic. Duplication of protocol function should be done only with
an understanding of system level implications, including possible
interactions with higher-layer mechanisms.
The architecture of the Internet was heavily influenced by the end-
to-end principle, and in our view it was crucial to the Internet's
success.
The remainder of this document discusses the various subnetwork
design issues that the authors consider relevant to efficient IP
support.
Maximum Transmission Units (MTUs) and IP Fragmentation
IPv4 packets (datagrams) vary in size from 20 bytes (the size of the
IP header alone) to a maximum of 65535 bytes. Subnetworks need not
support maximum-sized (64KB) IP packets, as IP provides a scheme that
breaks packets that are too large for a given subnetwork into
fragments that travel as independent packets and are reassembled at
the destination. The maximum packet size supported by a subnetwork is
known as its Maximum Transmission Unit (MTU).
Subnetworks may, but are not required to indicate the length of each
packet they carry. One example is Ethernet with the widely used DIX
[DIX] (not IEEE 802.3 [IEEE8023]) header, which lacks a length field
to indicate the true data length when the packet is padded to the 60
byte minimum. This is not a problem for uncompressed IP because it
carries its own length field.
If optional header compression [RFC1144] [RFC2507] [RFC2508]
[RFC3095] is used, however, it is required that the link framing
indicate frame length as it is needed for the reconstruction of the
original header.
In IP version 4 (the version now in wide use), fragmentation can
occur at either the sending host or in an intermediate router, and
fragments can be further fragmented at subsequent routers if
necessary.
In IP version 6, fragmentation can occur only at the sending host; it
cannot occur in a router.
Both IPv4 and IPv6 provide a "path MTU discovery" procedure [RFC1191]
[RFC1435] [RFC1981] that allows the sending host to avoid
fragmentation by discovering the minimum MTU along a given path and
reducing its packet sizes accordingly. This procedure is optional in
IPv4 but mandatory in IPv6 where there is no router fragmentation.
Path MTU discovery is widely deployed, but it sometimes encounters
problems. Some routers fail to generate the ICMP messages that convey
path MTU information to the sender, and sometimes the ICMP messages
are blocked by overly restrictive firewalls. The result can be a
"Path MTU Black Hole" [RFC2923] [RFC1435].
The Path MTU Discovery procedure, the persistence of path MTU black
holes, and the deletion of router fragmentation in IPv6 reflects a
consensus of the Internet technical community that IP fragmentation
is best avoided. This requires that subnetworks support MTUs that are
"reasonably" large. The smallest MTU permitted in IPv4 by [RFC791] is
68 bytes, but such a small value would clearly be inefficient.
Because IPv6 omits router fragmentation, [RFC 2460] specifies a
larger minimum MTU of 1280 bytes. Any subnetwork with an internal
packet payload smaller than 1280 bytes MUST implement an internal
fragmentation/reassembly mechanism if it is to support IPv6.
If a subnetwork cannot directly support a "reasonable" MTU with
native framing mechanisms, it should internally fragment. That is, it
should transparently break IP packets into internal data elements and
reassemble them at the other end of the subnetwork.
This leaves the question of what is a "reasonable" MTU. Ethernet (10
and 100 Mb/s) has a MTU of 1500 bytes, and because of its ubiquity
few Internet paths have MTUs larger than this value. This severely
limits the utility of larger MTUs provided by other subnetworks. But
larger MTUs are increasingly desirable on high speed subnetworks to
reduce the per-packet processing overhead in host computers, and
implementers are encouraged to provide them even though they may not
be usable when Ethernet is also in the path.
Various "tunneling" schemes, such as IP Security [RFC2406] treat IP
as a subnetwork for IP. Since tunneling adds header overhead, it can
trigger fragmentation even when the same physical subnetworks (e.g.,
Ethernet) are used on both sides of the IP router. Tunneling has made
it more difficult to avoid IP fragmentation and has increased the
incidence of path MTU black holes. Larger subnetwork MTUs may help
alleviate this problem.
Choosing the MTU in Slow Networks
In slow networks, the largest possible packet may take a considerable
time to send. Interactive response time should not exceed the well-
known human factors limit of 100 to 200 ms. This includes all sources
of delay: electromagnetic propagation delay, queuing delay, and the
store-and-forward time, i.e,. the time to transmit a packet at link
speed.
At low link speeds, store-and-forward delays can dominate total end-
to-end delay, and these are in turn directly influenced by the
maximum transmission unit (MTU). Even when an interactive packet is
given a higher queuing priority, it may have to wait for a large bulk
transfer packet to finish transmission. This worst-case wait can be
set by an appropriate choice of MTU.
For example, if the MTU is set to 1500 bytes, then a MTU-sized packet
will take about 8 milliseconds to send on a T1 (1.536 Mb/s) link.
But if the link speed is 19.2kb/s, then the transmission time becomes
625 ms -- well above our 100-200ms limit. A 256-byte MTU would lower
this delay to a little over 100 ms. However, care should be taken not
to lower the MTU excessively, as this will increase header overhead
and trigger frequent IP fragmentation (if Path MTU discovery is not
in use).
One way to limit delay for interactive traffic without imposing a
small MTU is to give priority to this traffic and to preempt (abort)
the transmission of a lower priority packet when a higher priority
packet arrives in the queue. However, the link resources used to
send the aborted packet are lost, and overall throughput will
decrease.
Another way is to implement a link-level multiplexing scheme that
allows several packets to be in progress simultaneously, with
transmission priority given to segments of higher priority IP
packets. For links using the Point-To-Point Protocol (PPP) [RFC1661],
multi-class multilink [RFC2686] [RFC2687] [RFC2689] provides such a
facility.
ATM (asynchronous transfer mode) is another example of this
technique. However, ATM is generally used on high speed links where
the store-and-forward delays are already minimal, and it introduces
significant (~9%) additional overhead due to the addition of 5-byte
frame headers to each 48-byte ATM cell.
A third example of link-layer multiplexing is the Data Over Cable
Service Interface Specifications. [DOCSIS1] [DOCSIS2] [DOCSIS3]
DOCSIS version 1.1 introduces an internal fragmentation and
reassembly procedure so that real-time voice packets can avoid undue
delay when sharing the relatively slow upstream channel of a cable
modem with large data packets.
To summarize, there is a fundamental tradeoff between efficiency and
latency in the design of a subnetwork, and the designer should keep
this in mind.
Framing on Connection-Oriented Subnetworks
IP requires that subnetworks mark the beginning and end of each
variable-length, asynchronous IP packet. Some examples of links and
subnetworks that do not provide this as an intrinsic feature include:
1. leased lines carrying a synchronous bit stream;
2. ISDN B-channels carrying a synchronous octet stream;
3. dialup telephone modems carrying an asynchronous octet stream;
and
4. Asynchronous Transfer Mode (ATM) networks carrying an asynchronous
stream of fixed-sized "cells".
The Internet community has defined packet framing methods for all
these subnetworks. The Point-To-Point Protocol (PPP) [RFC1661] is
applicable to bit synchronous, octet synchronous and octet
asynchronous links (i.e., examples 1-3 above). ATM has its own
framing methods described in [RFC2684] [RFC2364].
At high speeds, a subnetwork should provide a framed interface
capable of carrying asynchronous, variable-length IP datagrams. The
maximum packet size supported by this interface is discussed above in
the MTU/Fragmentation section. The subnetwork may implement this
facility in any convenient manner.
IP packet boundaries need not coincide with any framing or
synchronization mechanisms internal to the subnetwork. When the
subnetwork implements variable sized data units, the most
straightforward approach is to place exactly one IP packet into each
subnetwork data unit (SNDU), and to rely on the subnetwork's existing
ability to delimit SNDUs to also delimit IP packets. A good example
is Ethernet. But some subnetworks have SNDUs of one or more fixed
sizes, as dictated by switching, forward error correction and/or
interleaving considerations. Examples of such subnetworks include
ATM, with a single cell size of 48 bytes plus a 5-byte header, and
IS-95 digital cellular, with two "rate sets" of four fixed frame
sizes each that may be selected on 20 millisecond boundaries.
Because IP packets are variable length, they may not necessarily fit
into an integer multiple of fixed-sized SNDUs. An "adaptation layer"
is needed to convert IP packets into SNDUs while marking the boundary
between each IP packet in some manner.
There are several approaches to the problem. The first is to encode
each IP packet into one or more SNDUs, with no SNDU containing pieces
of more than one IP packet, and padding out the last SNDU of the
packet as needed. Bits in a control header added to each SNDU
indicate where it belongs in the IP packet. If the subnetwork
provides in-order, at-most-once delivery, the header can be as simple
as a pair of bits to indicate whether the SNDU is the first and/or
the last in the IP packet. Or only the last SNDU of the packet could
be marked, as this would implicitly mark the next SNDU as the first
in a new IP packet. The AAL5 (ATM Adaption Layer 5) scheme used with
ATM is an example of this approach, though it adds other features,
including a payload length field and a payload CRC.
In AAL5, the 1-bit per segment flag, carried in the ATM header,
indicates the end of a packet. The packet control information
(trailer) is located at the end of the segment. Placing the trailer
in a fixed position may simplify hardware reassembly.
Another framing technique is to insert per segment overhead to
indicate the presence of a segment option. When present, the option
carries a pointer to the end of the packet. This differs from AAL5
in that it permits another packet to follow within the same segment.
MPEG-2 [EN301] [ISO13181] supports this style of fragmentation, and
may utilize either padding (limiting each transport stream packet to
carry only part of one packet), or to allow a second packet to start
(no padding).
A third approach is to insert a special flag sequence into the data
stream between each IP packet, and to pack the resulting data stream
into SNDUs without regard to SNDU boundaries. The flag sequence can
also pad unused space at the end of an SNDU. If the special flag
appears in the user data, it is escaped to an alternate sequence
(usually larger than a flag) to avoid being misinterpreted as a flag.
The HDLC-based framing schemes used in PPP are all examples of this
approach.
All three adaptation schemes introduce overhead; how much depends on
the distribution of IP packet sizes, the size(s) of the SNDUs, and in
the HDLC-like approaches, the content of the IP packet (since flags
occurring in the packet must be escaped, which expands them). The
designer must also weigh implementation complexity in the choice and
design of an adaptation layer.
Connection-Oriented Subnetworks
IP has no notion of a "connection"; it is a purely connectionless
protocol. When a connection is required by an application, it is
usually provided by TCP, the Transmission Control Protocol, running
atop IP on an end-to-end basis.
Connection-oriented subnetworks can be (and are) widely used to carry
IP, but often with considerable complexity. Subnetworks with a few
nodes can simply open a permanent connection between each pair of
nodes. This is frequently done with ATM. But the number of
connections increases as the square of the number of nodes, so this
is clearly impractical for large subnetworks. A "shim" layer between
IP and the subnetwork is therefore required to manage connections.
This is one of the most common functions of a Subnetwork Dependent
Convergence Function (SNDCF) sublayer between IP and a subnetwork.
SNDCFs typically open subnetwork connections as needed when an IP
packet is queued for transmission and close them after an idle
timeout. There is no relation between subnetwork connections and any
connections that may exist at higher layers (e.g., TCP).
Because Internet traffic is typically bursty and transaction-
oriented, it is often difficult to pick an optimal idle timeout. If
the timeout is too short, subnetwork connections are opened and
closed rapidly, possibly over-stressing its call management system
(especially if was designed for voice traffic holding times). If the
timeout is too long, subnetwork connections are idle much of the
time, wasting any resources dedicated to them by the subnetwork.
The ideal subnetwork for IP is connectionless. Connection-oriented
networks that dedicate minimal resources to each connection (e.g.,
ATM) are a distant second, and connection-oriented networks that
dedicate a fixed amount of capacity to each connection (e.g., the
PSTN, including ISDN) are the least efficient. If such subnetworks
must be used to carry IP, their call-processing systems should be
capable of rapid call set-up and tear-down.
Bandwidth on Demand (BoD) Subnets
Wireless networks, including both satellite and terrestrial, may use
Bandwidth on Demand (BoD). Bandwidth on demand, which is implemented
at the link layer by Demand Assignment Multiple Access (DAMA) in TDMA
systems, is currently one proposed mechanism to efficiently share
limited spectrum resources amongst a large number of users.
The design parameters for BoD are similar to those in connection
oriented subnetworks, however the implementations may be very
different. In BoD, the user typically requests access to the shared
channel for some duration. Access may be allocated for a period of
time at a specific rate, a certain number of packets, or until the
user releases the channel. Access may be coordinated through a
central management entity or with a distributed algorithm amongst the
users. The resource shared may be a terrestrial wireless hop, a
satellite uplink, or an end-to-end satellite channel.
Long delay BoD subnets pose problems similar to connection oriented
networks in anticipating traffic. While connection oriented subnets
hold idle channels open expecting new data to arrive, BoD subnets
request channel access based on buffer occupancy (or expected buffer
occupancy) on the sending port. Poor performance will likely result
if the sender does not anticipate additional traffic arriving at that
port during the time it takes to grant a transmission request. It is
recommended that the algorithm have the capability to extend a hold
on the channel for data that has arrived after the original request
was generated (this may done by piggybacking new requests on user
data).
There are a wide variety of BoD protocols available and there has
been relatively little comprehensive research on the interactions
between the BoD mechanisms and Internet protocol performance. A
tradeoff exists balancing the time a user can be allowed to hold a
channel to drain port buffers with the additional imposed latency on
other users who are forced to wait to get access to the channel. It
is desirable to design mechanisms that constrain the BoD imposed
latency variation. This will be helpful in preventing spurious
timeouts from TCP.
Reliability and Error Control
In the Internet architecture, the ultimate responsibility for error
recovery is at the end points. The Internet may occasionally drop,
corrupt, duplicate or reorder packets, and the transport protocol
(e.g., TCP) or application (e.g., if UDP is used) must recover from
these errors on an end-to-end basis. Error recovery in the
subnetwork is therefore justified only to the extent that it can
enhance overall performance. It is important to recognize that a
subnetwork can go too far in attempting to provide error recovery
services in the Internet environment. Subnet reliability should be
"lightweight", i.e., it only has to be "good enough", *not* perfect.
In this section we discuss how to analyze characteristics of a
subnetwork to determine what is "good enough". The discussion below
focuses on TCP, which is the most widely used transport protocol in
the Internet. It is widely believed (and is a stated goal within the
IETF) that non-TCP transport protocols should attempt to be "TCP-
friendly" and have many of the same performance characteristics.
Thus, the discussion below should be applicable even to portions of
the Internet where TCP may not be the predominant protocol.
TCP vs Link Layer Retransmission
Error recovery involves the generation and transmission of redundant
information computed from user data. Depending on how much redundant
information is sent and how it is generated, the receiver can use it
to reliably detect transmission errors; correct up to some maximum
number of transmission errors; or both. The general approach is known
as Error Control Coding, or ECC.
For largely historical reasons, the use of ECC to detect transmission
errors so that retransmissions (hopefully without errors) can be
requested is widely known as "ARQ" (Automatic Repeat Request). ARQ
has been used for many decades in computer networking protocols.
When enough ECC information is available to permit the receiver to
correct transmission errors without a retransmission, the approach is
known as Forward Error Correction (FEC). Due to the greater
complexity of the required ECC and the need to tailor its design to
the characteristics of a specific modem and channel, FEC has
traditionally been implemented in special purpose hardware integral
to a modem. This effectively makes it part of the physical layer.
Unlike ARQ, FEC was seldom used for telecommunications outside of
deep space links until the 1990s. It is now nearly universal in
telephone, cable and DSL modems, digital satellite links and digital
mobile telephones. FEC is also heavily used in optical and magnetic
storage where "retransmissions" are not possible.
Some systems use hybrid combinations of ARQ layered atop FEC; V.90
dialup modems with V.42bis error control are one example. Most errors
are corrected by the trellis (FEC) code within the V.90 modem, and
most that remain are detected and corrected by the ARQ mechanisms in
V.42bis.
Work is now underway to apply FEC above the physical layer, primarily
in connection with reliable multicasting [RFC3048] where conventional
ARQ mechanisms are inefficient or difficult to implement. But in this
discussion we will assume that if FEC is present, it is implemented
within the physical layer.
Depending on the layer where it is implemented, error control can
operate on an end-to-end basis or over a shorter span such as a
single link. TCP is the most important example of an end-to-end
protocol that uses an ARQ strategy.
Many link layer protocols use ARQ, usually some flavor of HDLC
[ISO3309]. Examples include the X.25 link layer, the AX.25 protocol
used in amateur packet radio, 802.11 wireless LANs, and the reliable
link layer specified in IEEE 802.2.
As explained in the introduction, only end-to-end error recovery can
ensure a reliable service to the application. But some subnetworks
(e.g., many wireless links) also require link layer error recovery as
a performance enhancement. For example, many cellular links have
small physical frame sizes (< 100 bytes) and relatively high frame
loss rates. Relying entirely on end-to-end error recovery clearly
yields a performance degradation, as retransmissions across the end-
to-end path take much longer to be received than when link-local
retransmissions are used. Thus, link-layer error recovery can often
increase end-to-end performance. As a result, link-layer and end-to-
end recovery often co-exist; this can lead to the possibility of
inefficient interactions between the two layers of ARQ protocols.
This inter-layer "competition" might lead to the following wasteful
situation. When the link layer retransmits a packet, the link latency
momentarily increases. Since TCP bases its retransmission timeout on
prior measurements of end-to-end latency, including that of the link
in question, this sudden increase in latency may trigger an
unnecessary retransmission by TCP of a packet that the link layer is
still retransmitting. Such spurious end-to-end retransmissions
generate unnecessary load and reduce end-to-end throughput. One may
even have multiple copies of the same packet in the same link queue
at the same time. In general, one could say the competing error
recovery is caused by an inner control loop (link layer error
recovery) reacting to the same signal as an outer control loop (end-
to-end error recovery) without any coordination between the loops.
Note that this is solely an efficiency issue; TCP continues to
provide reliable end-to-end delivery over such links.
This raises the question of how persistent a link layer sender should
be in performing retransmission. We define the link layer (LL) ARQ
persistency as the maximum time that a particular link will spend
trying to transfer a packet before it can be discarded. This
deliberately simplified definition says nothing about maximum number
of retransmissions, retransmission strategies, queue sizes, queuing
disciplines, transmission delays, or the like. The reason we use the
term LL ARQ persistency instead of a term such as 'maximum link layer
packet holding time' is that the definition closely relates to link
layer error recovery. For example, on links that implement
straightforward error recovery strategies, LL ARQ persistency will
often correspond to a maximum number of retransmissions permitted per
link layer frame [ARQ-DRAFT].
For link layers that do not or cannot differentiate between flows
(e.g., due to network layer encryption), the LL ARQ persistency
should be small. This avoids any harmful effects or performance
degradation resulting from indiscriminate high persistence. A
detailed discussion of these issues is provided in [ARQ-DRAFT].
However, when a link layer can identify individual flows and apply
ARQ selectively [ARQ-DRAFT], then the link ARQ persistency should be
high for a flow using reliable unicast transport protocols (e.g.,
TCP) and must be low for all other flows. Setting the link ARQ
persistency larger than the largest link outage allows TCP to rapidly
restore transmission without the need to wait for a retransmission
time out. This generally improves TCP performance in the face of
transient outages. However, excessively high persistence may be
disadvantageous; a practical upper limit of 30-60 seconds may be
desirable. Implementation of such schemes remains a research issue.
(See also Section "Recovery from Subnetwork Outages").
Many subnetwork designers have opportunities to reduce the
probability of packet loss, e.g., with FEC, ARQ and interleaving, at
the cost of increased delay. TCP performance improves with decreasing
loss but worsens with increasing end-to-end delay, so it is important
to find the proper balance through analysis and simulation.
Recovery from Subnetwork Outages
Some types of subnetworks, particularly mobile radio, are subject to
frequent temporary outages. For example, an active cellular data user
may drive or walk into an area (such as a tunnel) that is out of
range of any base station. No packets will be successfully delivered
until the user returns to an area with coverage.
The Internet protocols currently provide no standard way for a
subnetwork to explicitly notify an upper layer protocol (e.g., TCP)
that it is experiencing an outage rather than severe congestion.
Under these circumstances TCP will, after each unsuccessful
retransmission, wait even longer before trying again; this is its
"exponential back-off" algorithm. And TCP will not discover that the
subnetwork outage has ended until its next retransmission attempt. If
TCP has backed off, this may take some time. This can lead to
extremely poor TCP performance over such subnetworks.
It is therefore highly desirable that a subnetwork subject to outages
not silently discard packets during an outage. Ideally, it should
define an interface to the next higher layer (i.e., IP) that allows
it to refuse packets during an outage, and to automatically ask IP
for new packets when it is again able to deliver them. If it cannot
do this, then the subnetwork should hold onto at least some of the
packets it accepts during an outage and attempt to deliver them when
the subnetwork comes back up.
Note that it is *not* necessary to completely avoid dropping packets
during an outage. The purpose of holding onto a packet during an
outage, either in the subnetwork or at the IP layer, is so that its
eventual delivery will implicitly notify TCP that the subnetwork is
again operational. This is to enhance performance, not to ensure
reliability -- a task that as discussed earlier can only be done
properly on an end-to-end basis.
Only a single packet per TCP connection, including ACKs, need be held
in this way to cause the TCP sender to recover from the additional
losses once the flow resumes.
Because it would be a layering violation (and possibly a performance
hit) for IP or a subnetwork layer to look at TCP headers (which would
in any event be impossible if IPSEC [RFC2401] encryption is in use),
it would be reasonable for the IP or subnetwork layers to choose, as
a design parameter, some small number of packets that it will retain
during an outage.
CRCs, Checksums and Error Detection
The TCP, UDP and IPv4 protocols all use the same simple 16-bit 1's
complement checksum algorithm to detect corrupted packets. The IP
checksum protects only the IP header, while the TCP and UDP checksums
protect both the TCP/UDP header and any user data.
These checksums are not very strong from a coding theory standpoint.
But they are easy to compute in software, and various proposals to
replace them with stronger checksums have failed. Yet a study
[SP2000] has shown that the Internet corrupts one in 1,100 to 32,000
packets, and it is up to the end-to-end Internet checksum to detect
these errors.
Most packet corruption appears to be caused by bugs and errors in
host and router hardware and software. So even if every subnetwork
implemented strong error detection, the end-to-end use of TCP and UDP
checksums would still be necessary.
Most subnetworks implement error detection just above the physical
layer. Packets corrupted in transmission are detected and discarded
before delivery to the IP layer. A 16-bit cyclic redundancy check
(CRC) is usually the minimum, and this is known to be considerably
stronger against most kinds of errors than the 16-bit standard
Internet checksum. The Point-to-Point Protocol [RFC1662] requires
support of a 16-bit CRC, with a 32-bit CRC as an option. (Note that
PPP is often used in conjunction with a dialup modem, which provides
its own error control). Other subnetworks, including 802.3/Ethernet,
AAL5/ATM, FDDI, Token Ring and PPP over SONET/SDH all use a 32-bit
CRC that is considerably stronger. In addition, many subnetworks
(notably dialup modems, mobile radio and satellite channels) also
incorporate forward error correction, often in hardware.
Any new subnetwork designed to carry IP should therefore provide
error detection at least as strong as the 32-bit CRC specified in
[ISO3309]. While this will achieve a very low undetected packet
error rate, it will not (and need not) achieve a very low packet loss
rate as the Internet protocols are better suited to dealing with lost
packets than with corrupted packets.
For link layers that can differentiate between flows, it may be
appropriate to reduce the error detection level for certain flows
with large numbers of small packets, such as voice flows. As such
flows also benefit significantly from header compression, this should
only be combined with a header compression scheme that is robust
against residual bit errors [RFC3095].
Designers of complex subnetworks consisting of internal links and
packet switches should consider implementing error detection on an
edge-to-edge basis, i.e., at the interface to IP, either in addition
to or instead of error detection at the interface to each physical
link. This has the significant advantage of protecting against errors
introduced anywhere in the subnetwork, not just its transmission
links.
This is straightforward if the interface presented to IP by the
subnetwork already includes error detection, as with PPP or Ethernet.
If the subnetwork carries the PPP or Ethernet CRC without change
through the subnetwork, it will automatically provide the desired
edge-to-edge error detection. An existing example of such a
subnetwork is an Ethernet bridge, also known as a switched hub.
IP version 6 (IPv6) has no IP header checksum. The destination host
detects "important" errors in the IP header such as the delivery of
the packet to the wrong destination. This is done by including the IP
source and destination addresses in the computation of the checksum
in the TCP or UDP header, a practice already performed in IPv4.
Errors in other IPv6 header fields may go undetected; this was
considered a reasonable price to pay for a considerable reduction in
the processing required by each router. If desired, additional
protection may be obtained for the IPv6 header by the use of the
authentication and packet integrity services of the IP Security
(IPSEC) protocol.
How TCP Works
One of TCP's functions is end-host based congestion control for the
Internet. This is a critical part of the overall stability of the
Internet, so it is important that link layer designers understand
TCP's congestion control algorithms.
TCP assumes that, at the most abstract level, the network consists of
links and queues. Queues provide output-buffering on links that are
momentarily oversubscribed. They smooth instantaneous traffic bursts
to fit the link bandwidth.
When demand exceeds link capacity long enough to fill the queue,
packets must be dropped. The traditional action of dropping the most
recent packet ("tail dropping") is no longer recommended [RED93], but
it is still widely practiced.
TCP uses sequence numbering and acknowledgments (ACKs) on an end-to-
end basis to provide reliable, sequenced, once-only delivery. TCP
ACKs are cumulative, i.e., each implicitly ACKs every segment
received so far. If a packet is lost, the cumulative ACK will cease
to advance.
Since the most common cause of packet loss is congestion, TCP treats
packet loss as an Internet congestion indicator. This happens
automatically, and the subnetwork need not know anything about IP or
TCP. It simply drops packets whenever it must, though some packet-
dropping strategies (e.g., RED) are more fair than others.
TCP recovers from packet losses in two different ways. The most
important is the retransmission timeout. If an ACK fails to arrive
after a certain period of time, TCP retransmits the oldest unacked
packet. Taking this as a hint that the network is congested, TCP
waits for the retransmission to be ACKed before it continues, and it
gradually increases the number of packets in flight as long as a
timeout does not occur again.
A retransmission timeout can impose a significant performance
penalty, as the sender is idle during the timeout interval and
restarts with a congestion window of 1 following the timeout. To
allow faster recovery from the occasional lost packet in a bulk
transfer, an alternate scheme known as "fast recovery" was introduced
[RFC2581] [RFC2582] [RFC2914] [TCPF98].
Fast recovery relies on the fact that when a single packet is lost in
a bulk transfer, the receiver continues to return ACKs to subsequent
data packets that do not actually acknowledge any data. These are
known as "duplicate acknowledgments" or "dupacks". The sending TCP
can use dupacks as a hint that a packet has been lost and retransmit
it without waiting for a timeout. Dupacks effectively constitute a
negative acknowledgment (NAK) for the packet sequence number in the
acknowledgment field. TCP waits until a certain number of dupacks
(currently 3) are seen prior to assuming a loss has occurred; this
helps avoid an unnecessary retransmission during out-of-sequence
delivery. Recent proposals have been made to lower the dupack
threshold to 2.
A new technique called "Explicit Congestion Notification" (ECN)
allows routers to directly signal congestion to hosts without
dropping packets. This is done by setting a bit in the IP header.
Since this is currently an optional behavior (and, longer term, there
will always be the possibility of congestion in portions of the
network which don't support ECN), the lack of an ECN bit MUST NEVER
be interpreted as a lack of congestion. Thus, for the foreseeable
future, TCP MUST interpret a lost packet as a signal of congestion.
The TCP "congestion avoidance" [RFC2581] algorithm maintains a
congestion window (cwnd) controlling the amount of data TCP may have
in flight at any moment. Reducing cwnd reduces the overall bandwidth
obtained by the connection; similarly, raising cwnd increases the
performance, up to the limit of the available bandwidth.
TCP probes for available network bandwidth by setting cwnd at one
packet and then increasing it by one packet for each ACK returned
from the receiver. This is TCP's "slow start" mechanism. When a
packet loss is detected (or congestion is signaled by other
mechanisms), cwnd is reset to one and the slow start process is
repeated until cwnd reaches one half of its previous setting before
the reset. Cwnd continues to increase past this point, but at a much
slower rate than before. If no further losses occur, cwnd will
ultimately reach the window size advertised by the receiver.
This is an "Additive Increase, Multiplicative Decrease" (AIMD)
algorithm. The steep decrease of cwnd in response to congestion
provides for network stability; the AIMD algorithm also provides for
fairness between long running TCP connections sharing the same path.
TCP Performance Characteristics
Caveat
Here we present the current "state-of-the-art" understanding of TCP
performance. This analysis attempts to characterize the performance
of TCP connections over links of varying characteristics.
Link designers may wish to use the techniques in this section to
predict what performance TCP/IP may achieve over a new link layer
design. Such analysis is encouraged. Because this is relatively new
analysis, and the theory is based on single stream TCP connections
under "ideal" conditions, it should be recognized that the results of
such analysis may be different than actual performance in the
Internet. That being said, we have done the best we can to provide
information which will help designers get an accurate picture of the
capabilities and limitations of TCP under various conditions.
The Formulae
The performance of TCP's AIMD Congestion Avoidance algorithm has been
extensively analyzed. The current best formula for the performance
of the specific algorithms used by Reno TCP is given by Padhye, et al
[PFTK98]. This formula is:
MSS
BW = --------------------------------------------------------
RTT*sqrt(1.33*p) + RTO*p*[1+32*p^2]*min[1,3*sqrt(.75*p)]
where
BW is the maximum throughput achievable
MSS is the segment size being used by the connection
RTT is the end-to-end round trip time of the TCP connection
RTO is the packet timeout (based on RTT)
p is the packet loss rate for the path
(i.e. .01 if there is 1% packet loss)
Note that the speed of the links making up the Internet path does not
explicitly appear in this formula. Attempting to send faster than the
slowest link in the path causes the queue to grow at the transmitter
driving the bottleneck. This increases the RTT, which in turn reduces
the achievable throughput.
This is currently considered to be the best approximate formula for
Reno TCP performance. A further simplification to this formula is
generally made by assuming that RTO is approximately 5*RTT.
TCP is constantly being improved. A simpler formula, which gives an
upper bound on the performance of any AIMD algorithm which is likely
to be implemented in TCP in the future, was derived by Ott, et al
[MSMO97][OKM96].
MSS 1
BW = C --- -------
RTT sqrt(p)
where C is 0.93.
Assumptions
Both formulae assume that the TCP Receiver Window is not limiting the
performance of the connection. Because receiver window is entirely
determined by end-hosts, we assume that hosts will maximize the
announced receiver window to maximize their network performance.
Both of these formulae allow BW to become infinite if there is no
loss. This is because an Internet path will drop packets at
bottleneck queues if the load is too high. Thus, a completely
lossless TCP/IP network can never occur (unless the network is being
underutilized).
The RTT used is the average, including queuing delays.
The formulae are for a single TCP connection. If a path carries many
TCP connections, each will follow the formulae above independently.
The formulae assume long running TCP connections. For connections
which are extremely short (<10 packets) and don't lose any packets,
performance is driven by the TCP slow start algorithm. For
connections of medium length, where on average only a few segments
are lost, single connection performance will actually be slightly
better than given by the formulae above.
The difference between the simple and complex formulae above is that
the complex formula includes the effects of TCP retransmission
timeouts. For very low levels of packet loss (significantly less
than 1%), timeouts are unlikely to occur, and the formulae lead to
very similar results. At higher packet losses (1% and above), the
complex formula gives a more accurate estimate of performance (which
will always be significantly lower than the result from the simple
formula).
Note that these formulae break down as p approaches 100%.
Analysis of Link Layer Effects on TCP Performance
Consider the following example:
A designer invents a new wireless link layer which, on average, loses
1% of IP packets. The link layer supports packets of up to 1040
bytes, and has a one-way delay of 20 msec.
If this link layer were used in the Internet, on a path that
otherwise had a round trip of of 80 msec, you could compute an upper
bound on the performance as follows:
For MSS, use 1000 bytes to exclude the 40 bytes of TCP/IP headers.
For RTT, use 120 msec (80 msec for the Internet part, plus 20 msec
each way for the new wireless link).
For p, use .01. For C, assume 1.
The simple formula gives:
BW = (1000 * 8 bits) / (.120 sec * sqrt(.01)) = 666 kbit/sec
The more complex formula gives:
BW = 402.9 kbit/sec
If this were a 2 Mb/s wireless LAN, the designers might be somewhat
disappointed.
Some observations on performance:
1. We have assumed that the packet losses on the link layer are
interpreted as congestion by TCP. This is a "fact of life" that must
be accepted.
2. The equations for TCP performance are all expressed in terms of
packet loss, but many link-layer designers think in terms of bit-
error rate. *If* channel bit errors are independent, then the
probability of a packet being corrupted is:
p = 1 - ([1 - BER]^[PACKET_SIZE*8])
Here we assume PACKET_SIZE is in bytes. It includes the user data and
all headers (TCP,IP and subnetwork). If the inequality
BER * [PACKET_SIZE*8] << 1
holds, the packet loss probability p can be approximated by:
p = BER * [PACKET_SIZE*8]
These equations can be used to apply BER to the performance equations
above.
Note that PACKET_SIZE can vary from one packet to the next. Small
packets (such as TCP acks) generally have a smaller probability of
packet error than, say, a TCP packet carrying one MSS (maximum
segment size) of user data. A flow of small TCP acks can be expected
to be slightly more reliable than a stream of larger TCP data
segments.
It bears repeating that the above analysis assumes that bit errors
are statistically independent. Because this is not true for many real
links, our computation of p is actually an upper bound, not the exact
probability of packet loss.
There are many reasons why bit errors are not independent on real
links. Many radio links are affected by propagation fading or by
interference that lasts over many bit times.
Also, links with Forward Error Correction (FEC) generally have very
non-uniform bit error distributions that depend on the type of FEC,
but in general the uncorrected errors tend to occur in bursts even
when channel symbol errors are independent. In all such cases our
computation of p from BER can only place an upper limit on the packet
loss rate.
If the distribution of error distributions under the FEC scheme is
known, one could apply the same type of analysis as above, using the
correct distribution function for the BER. It is more likely in
these FEC cases, however, that empirical methods will need to be used
to determine the actual packet loss rate.
3. Note that the packet size plays an important role. If the
subnetwork loss characteristics are such that large packets have the
same probability of loss as smaller packets, then larger packets will
yield improved performance.
4. We have chosen a specific RTT that might occur on a wide-area
Internet path within the USA. It is important to recognize that RTTs
vary considerably in the Internet.
For example, RTTs are typically less than 10 msec in a wired LAN
environment. International connections may have RTTs of 200 msec or
more. Modems and other low-capacity links can add considerable delay
due to their long packet transmission times.
Links over geostationary repeater satellites have one-way speed-of-
light delays of around 250ms: 125ms up to the satellite and 125ms
down. The RTT of an end-to-end TCP connection that includes such a
link can be expected to be greater than 250ms.
Queues on heavily congested links may back up, increasing RTTs.
Finally, VPNs and other forms of encryption and tunneling can add
significant end-to-end delay to network connections.
Quality-of-Service (QoS) considerations
It is generally recognized that specific service guarantees are
needed to support real-time multimedia, toll quality telephony and
other performance critical applications. The provision of such
Quality of Service guarantees in the Internet is an active area of
research and standardization. The IETF has not converged on a single
service model, set of services or mechanisms that will offer useful
guarantees to applications and be scalable to the Internet. Indeed,
the IETF does not have a single definition of Quality of Service.
[RFC2990] represents the present understanding of the challenges in
architecting QoS for the Internet.
There are presently two architectural approaches to providing
mechanisms for QoS support in the Internet.
IP Integrated Services (Intserv) [RFC1633] provides fine-grained
service guarantees to individual flows. Flows are identified by a
flow specification (flowspec), which creates a stateful association
between individual packets by matching fields in the packet header.
Bandwidth is reserved for the flow, and appropriate traffic
conditioning and scheduling is installed in routers along the path.
The ReSerVation Protocol (RSVP) [RFC2205, RFC2210] usually, but not
necessarily, is used to install flows. Intserv defines two services,
in addition to the Default (best effort) service.
-- Guaranteed Service (GS) [RFC 2212] offers hard upper bounds on
delay to flows that conform to a traffic specification (TSpec). It
uses a fluid flow model to relate the TSpec and reserved bandwidth
(RSpec) to variable delay. Non-conforming packets are forwarded on a
best-effort basis.
-- Controlled Load Service (CLS) [RFC2211] offers delay and packet
loss equivalent to that of an unloaded network to flows that conform
to a TSpec, but no hard bounds. Non-conforming packets are forwarded
on a best-effort basis.
Intserv requires installation of state information in every
participating router. Absent this state in every router along the
path, performance guarantees cannot be made. This, along with RSVP
processing and the need for usage-based accounting, is believed to
have scalability problems, particularly in the core of the Internet
[RFC2208].
IP Differentiated Services (Diffserv) [RFC2475] provides a "toolkit"
offering coarse-grained controls to aggregates of flows. Diffserv in
itself does NOT provide QoS guarantees, but can be used to construct
services with QoS guarantees across a Diffserv domain. It attempts
to address the scaling issues associated with Intserv by requiring
state awareness only at the edge of a Diffserv domain. At the edge,
packets are classified into flows, and the flows are conditioned
(marked, policed or shaped) to a traffic conditioning specification
(TCS). A Diffserv Codepoint (DSCP), identifying a per-hop behavior
(PHB), is set in each packet header. The DSCP is carried in the DS-
field, subsuming six bits of the former TOS byte of the IP header
[RFC2474]. The PHB denotes the forwarding behavior to be applied to
the packet in each node in the Diffserv domain. Although there is a
"recommended" DSCP associated with each PHB, the mappings from DSCPs
to PHBs are defined by the DS-domain. In fact, there can be several
DSCPs associated with the same PHB. Diffserv presently defines three
PHBs.
The class selector PHB [RFC2474] replaces the IP precedence field of
the former TOS byte. It offers relative forwarding priorities.
The Expedited Forwarding (EF) PHB [RFC2598] guarantees that packets
will have a well-defined minimum departure rate which, if not
exceeded, ensures that the associated queues are short or empty. EF
is intended to support services that offer tightly bounded loss,
delay and delay jitter.
The Assured Forwarding (AF) PHB group [RFC2597] offers different
levels of forwarding assurances for packets belonging to an
aggregated flow. Each AF group is independently allocated forwarding
resources. Packets are marked with one of three drop precedences;
those with the highest drop precedence are dropped with lower
probability than those marked with the lowest drop precedence. DSCPs
are recommended for four independent AF groups, although a DS domain
can have more or fewer AF groups.
Ongoing work in the IETF is addressing ways to support Intserv with
Diffserv. There is some belief (e.g. as expressed in [RFC 2990])
that such an approach will allow individual flows to receive service
guarantees and scale to the global Internet.
The QoS guarantees that can be offered by the IP layer are a product
of two factors:
-- the concatenation of the QoS guarantees offered by the subnets
along the path of a flow. This implies that a subnet may wish to
offer multiple services (with different QoS guarantees) to the IP
layer, which can then determine which flows use which subnet service.
Or, to put it another way, forwarding behavior in the subnet needs to
be 'clued' by the forwarding behavior (service or PHB) at the IP
layer, and
-- the operation of a set of cooperating mechanisms, such as
bandwidth reservation and admission control, policy management,
traffic classification, traffic conditioning (marking, policing
and/or shaping), selective discard, queuing and scheduling. Note
that support for QoS in subnets may require similar mechanisms,
especially when these subnets are general topology subnets (e.g.,
ATM, frame relay or MPLS) or shared media subnets.
Many subnetwork designers face inherent tradeoffs between delay,
throughput, reliability and cost. Other subnetworks have parameters
that manage bandwidth, internal connection state, and the like.
Therefore, the following subnetwork capabilities may be desirable,
although some might be trivial or moot if the subnet is a simple
point-to-point link.
- The subnetwork should have the ability to reserve bandwidth for a
connection or flow and schedule packets accordingly.
- Bandwidth reservations should be based on a one- or two- token
bucket model, depending on whether the service is intended to support
constant rate or bursty traffic.
- If a connection or flow does not use its reserved bandwidth at a
given time, the unused bandwidth should be available for other flows.
- Packets in excess of a connection or flow's agreed rate should be
forwarded as best effort or discarded, depending on the service
offered by the subnet to the IP layer.
- If a subnet contains error control mechanisms (retransmission
and/or FEC), it should be possible for the IP layer to influence the
inherent tradeoff between uncorrected errors, packet losses and
delay. These capabilities at the subnet/IP layer service boundary
correspond to to selection of more or less error control and/or to
selection of particular error control mechanisms within the
subnetwork.
- The subnet layer should know, and be able to inform the IP layer,
how much fixed delay and delay jitter it offers for a flow or
connection. If the Intserv model is used, the delay jitter component
may best be expressed in terms of the TSpec/RSpec model described in
[RFC2212].
- Support of the Diffserv class selectors [RFC2474] suggests that the
subnet might consider mechanisms that support priorities.
Fairness vs Performance
Subnetwork designers should be aware of the tradeoffs between
fairness and efficiency inherent in many transmission scheduling
algorithms. For example, many local area networks use contention
protocols to resolve access to a shared transmission channel. These
protocols represent overhead. Limiting the amount of data that a
station may transmit per contention cycle helps assure each station
of timely access to the channel, but it also increases contention
overhead per unit of data sent.
In some mobile radio networks, capacity is limited by interference,
which in turn depends on average transmitter power. Some receivers
may require considerably more transmitter power (generating more
interference and consuming more channel capacity) than others.
In each case, the scheduling algorithm designer must balance
competing objectives: providing a fair share of capacity to each
station while maximizing the total capacity of the network.
Delay Characteristics
TCP bases its retransmission timeout (RTO) on measurements of the
round trip delay experienced by previous packets. This allows TCP to
adapt automatically to the very wide range of delays found on the
Internet. The recommended algorithms are described in [RFC2988].
These algorithms model the delay along an Internet path as a
normally-distributed random variable with slowly varying mean and
standard deviation. TCP estimates these two parameters by
exponentially smoothing individual delay measurements, and it sets
the RTO to the estimated mean delay plus some fixed number of
standard deviations. (The algorithm actually uses mean deviation as
an approximation to standard deviation, as it is easier to compute.)
The goal is to compute a RTO that is small enough to detect and
recover from packet losses while minimizing unnecessary ("spurious")
retransmissions when packets are unexpectedly delayed but not lost.
Although these goals conflict, the algorithm works well when the
delay variance along the Internet path is low, or the packet loss
rate is low.
If the path delay variance is high, TCP sets a RTO that is much
larger than the mean of the measured delays. But if the packet loss
rate is low, the large RTO is of little consequence, as timeouts
occur only rarely. Conversely, if the path delay variance is low,
then TCP recovers quickly from lost packets; again, the algorithm
works well.
But when delay variance and the packet loss rate are both high, these
algorithms perform poorly, especially when the mean delay is also
high.
Because TCP uses returning acknowledgments as a "clock" to time the
transmission of additional data, excessively high delays (even if the
delay variance is low) also affects TCP's ability to fully utilize a
high speed transmission pipe. It also slows down the recovery of lost
packets even when delay variance is small.
Subnetwork designers should therefore minimize all three parameters
(delay, delay variance and packet loss) as much as possible.
In many subnetworks, these parameters are inherently in conflict.
For example, on a mobile radio channel the subnetwork designer can
use retransmission (ARQ) and/or forward error correction (FEC) to
trade off delay, delay variance and packet loss in an effort to
improve TCP performance. For example, while ARQ increases delay
variance, FEC does not. However, FEC (especially when combined with
interleaving) often increases mean delay even on good channels where
ARQ would not increase either the delay or the delay variance.
The tradeoffs among these error control mechanisms and their
interactions with TCP can be quite complex, and they are the subject
of much ongoing research. We therefore recommend that subnetwork
designers provide as much flexibility as possible in the
implementation of these mechanisms, and to provide access to them as
discussed above in the section on Quality of Service.
Bandwidth Asymmetries
Some subnetworks may provide asymmetric bandwidth and the Internet
protocol suite will generally still work fine. However, there is a
case when such a scenario reduces TCP performance. Since TCP data
segments are 'clocked' out by returning acknowledgments TCP senders
are limited by the rate at which ACKs can be returned [BPK98].
Therefore, when the ratio of the bandwidth of the subnetwork carrying
the data to the bandwidth of the subnetwork carrying the
acknowledgments is too large, the slow return of of the ACKs directly
impacts performance. Since ACKs are generally smaller than data
segments, TCP can tolerate some asymmetry, but as a general rule
designers of subnetworks should avoid large differences in the
incoming and outgoing bandwidth.
One way to cope with asymmetric subnetworks is to increase the size
of the data segments as much as possible. This allows more data to
be sent per ACK, mitigating the slow flow of ACKs. Using the delayed
acknowledgment mechanism [Bra89] that reduces the number of ACKs
transmitted by the receiver by roughly half can also improve
performance by reducing the congestion on the ACK channel. These
mechanisms should be employed in asymmetric networks.
Several researchers have introduced strategies for coping with
bandwidth asymmetry. These mechanisms generally attempt to reduce
the number of ACKs being transmitted over the low bandwidth channel
by limiting the ACK frequency or filtering out ACKs at an
intermediate router [BPK98]. While these solutions mitigate the
performance problems caused by asymmetric subnetworks, they do have
some cost. Therefore, as suggested above, bandwidth asymmetry should
be minimized in subnetwork designs.
Buffering, flow & congestion control
Many subnets include multiple links with varying traffic demands and
possibly different transmission speeds. At each link there must be a
queuing system, including buffering, scheduling and a capability to
discard excess subnet packets. These queues may also be part of a
subnet flow control or congestion control scheme.
For the purpose of this discussion, we talk about packets without
regard to whether they refer to a complete IP datagram or a
subnetwork packet. At each queue, a packet experiences a delay that
depends on competing traffic and the scheduling discipline, and is
subjected to a local discarding policy.
Some subnets may have flow or congestion control mechanisms in
addition to packet dropping. Such mechanisms can operate on
components in the subnet layer, such as schedulers, shapers or
discarders, and can affect the operation of IP forwarders at the
edges of the subnet. However, with the exception of RFC2481
explicit congestion notification (discussed below), IP has no way to
pass explicit congestion or flow control signals to TCP.
TCP traffic, especially aggregated TCP traffic, is bursty. As a
result, instantaneous queue depths can vary dramatically, even in
nominally stable networks. For optimal performance, packets should
be dropped in a controlled fashion, not just when buffer space is
unavailable. How much buffer space should be supplied is still a
matter of debate, but as a rule of thumb, each node should have
enough buffering to hold one bandwidth*delay product's worth of data
for each TCP connection sharing the link.
This is often difficult to estimate, since it depends on parameters
beyond the subnetwork's control or knowledge. Internet nodes
generally do not implement admission control policies, and cannot
limit the number of TCP connections that use them. In general, it is
wise to err in favor of too much buffering rather than too little.
It may also be useful for subnets to incorporate mechanisms that
measure propagation delays to assist in buffer sizing calculations.
There is a rough consensus in the research community that active
queue management is important to improving fairness, link utilization
and throughput [RFC2309]. Although there are questions and concerns
about the effectiveness of active queue management (e.g., see
[MBDL99]), it is widely considered an improvement over tail-drop
discard policies.
One form of active queue management is the Random Early Detection
(RED) algorithm [RED93], actually a family of related algorithms. In
one version of RED, an exponentially weighted moving average of the
queue depth is maintained. When this average queue depth is between
a maximum threshold max_th, and a minimum threshold min_th, packets
are dropped with a probability which is proportional to the amount by
which the average queue depth exceeds min_th. When this average
queue depth is equal to max_th, the drop probability is equal to a
configurable parameter max_p. When this average queue depth is
greater than max_th, packets are always dropped. Numerous variants
on RED appear in the literature, and there are other active queue
management algorithms which claim various advantages over RED [MG01].
With an active queue management algorithm, dropped packets become a
feedback signal to trigger more appropriate congestion behavior by
the TCPs in the end hosts. Randomization of dropping tends to break
up the observed tendency of TCP windows belonging to different TCP
connections to become synchronized by correlated drops, and it also
imposes a degree of fairness on those connections that properly
implement TCP congestion avoidance. Another important property of
active queue management algorithms is that they attempt to keep
average queue depths short while accommodating large short term
bursts.
Since TCP neither knows nor cares whether congestive packet loss
occurs at the IP layer or in a subnet, it may be advisable for
subnets that perform queuing and discarding to consider implementing
some form of active queue management. This is especially true if
large aggregates of TCP connections are likely to share the same
queue. However, active queue management may be less effective in the
case of many queues carrying smaller aggregates of TCP connections,
e.g., in an ATM switch that implements per-VC queuing.
Note that the performance of active queue management algorithms is
highly sensitive to settings of configurable parameters, and also to
factors such as RTT [MBB00][FB00].
Some subnets, most notably ATM, perform segmentation and reassembly
at the subnetwork edges. Care should be taken here in designing
discard policies. If the subnet discards a fragment of an IP packet,
then the remaining fragments become an unproductive load on the
subnet that can markedly degrade end-to-end performance [RF95].
Subnetworks should therefore attempt to discard these extra fragments
whenever one of them must be discarded. If the IP packet has already
been partially forwarded when discarding becomes necessary, then
every remaining fragment except the one marking the end of the IP
packet should also be discarded. For ATM subnets, this specifically
means using Early Packet Discard and Partial Packet Discard [ATMFTM].
Some subnets might include flow control mechanisms that effectively
require that the rate of traffic flows be shaped as they enter the
subnet. One example of such a subnet mechanism is in the ATM
Available Bit rate (ABR) service category [ATMFTM]. Such flow
control mechanisms have the effect of making the subnet nearly
lossless by pushing congestion into the IP routers at the edges of
the subnet. In such a case, adequate buffering and discard policies
are needed in these routers to deal with a subnet that appears to
have varying bandwidth. Whether there is benefit in this kind of
flow control is controversial; there are numerous simulation and
analytical studies that go both ways. It appears that some of the
issues that lead to such different results include sensitivity to ABR
parameters, use of binary rather than explicit rate feedback, use (or
not) of per-VC queuing, and the specific ATM switch algorithms
selected for the study. Anecdotally, some large networks have used
IP over ABR to carry TCP traffic, have claimed it to be successful,
but have published no results.
Another possible approach to flow control in the subnet would be to
work with TCP Explicit Congestion Notification (ECN) semantics
[RFC2481] [RFB01]. Routers at the edges of the subnet, rather than
shaping, would set the ECN bit in those IP packets that are received
in subnet packets that have an ECN indication. Nodes in the subnet
would need to implement an active queue management protocol that
marks subnet packets instead of dropping them.
ECN is currently a proposed standard, and it is not yet widely
deployed.
Compression
Application data compression is a function that can usually be
omitted in the subnetwork. The endpoints typically have more CPU and
memory resources to run a compression algorithm and a better
understanding of what is being compressed. End-to-end compression
benefits every network element in the path, while subnetwork-layer
compression, by definition, benefits only a single subnetwork.
Data presented to the subnetwork layer may already be in compressed
format (e.g., a JPEG file), compressed at the application layer
(e.g., the optional "gzip", "compress", and "deflate" compression in
HTTP/1.1 [RFC2616]), or compressed at the IP layer (the IP Payload
Compression Protocol [RFC2393] supports DEFLATE [RFC2394] and LZS
[RFC2395]). In any of these cases, compression in the subnetwork is
of no benefit.
The subnetwork may also process data that has been encrypted by the
application (OpenPGP [RFC2440] or S/MIME [RFCs-2630-2634]), just
above TCP (SSL, TLS [RFC2246]), or just above IP (IPSEC ESP
[RFC2406]). Ciphers generate random-looking bit streams lacking any
patterns that can be exploited by a compression algorithm.
However, much data is still transmitted uncompressed over the
Internet, so subnetwork compression may be beneficial. Any
subnetwork compression algorithm must not expand uncompressible data,
e.g., data that has already been compressed or encrypted.
We make a stronger recommendation that subnetworks operating at low
speed or with small MTUs compress IP and transport-level headers (TCP
and UDP) using several header compression schemes developed within
the IETF. An uncompressed 40-byte TCP/IP header takes about 33
milliseconds to send at 9600 bps. "VJ" TCP/IP header compression
[RFC1144] compresses most headers to 3-5 bytes, reducing transmission
time to several milliseconds. This is especially beneficial for
small, latency-sensitive packets in interactive sessions.
Similarly, RTP compression schemes such as CRTP [RFC2508] and ROHC
[RFC3095] compress most IP/UDP/RTP headers to one to four bytes. The
resulting savings are especially significant when audio packets are
kept small to minimize store-and-forward latency.
Designers should consider the effect of the subnetwork error rate on
the performance of header compression. TCP ordinarily recovers from
lost packets by retransmitting only those packets that were actually
lost; packets arriving correctly after a packet loss are kept on a
resequencing queue and do not need to be retransmitted. In VJ TCP/IP
[RFC1144] header compression, however, the receiver cannot explicitly
notify a sender about data corruption and subsequent loss of
synchronization between compressor and decompressor. It relies
instead on TCP retransmission to re-synchronize the decompressor.
After a packet is lost, the decompressor must discard every
subsequent packet, even if the subnetwork makes no further errors,
until the sending TCP retransmits to re-synchronize the decompressor.
This effect can substantially magnify the effect of subnetwork packet
losses if the sending TCP window is large, as it will often be on a
path with a large bandwidth*delay product.
Alternative header compression schemes such as those described in
[RFC2507] include an explicit request for retransmission of an
uncompressed packet to allow decompressor resynchronization without
waiting for a TCP retransmission. However, these schemes are not yet
in widespread use.
Both TCP header compression schemes do not compress widely used TCP
options such as selective acknowledgements (SACK). Both fail to
compress TCP traffic that makes use of explicit congestion
notification (ECN). Work is under way in the IETF ROHC WG to address
these shortcomings in a ROHC header compression scheme for TCP.
The subnetwork error rate also is important for RTP header
compression. CRTP uses delta encoding, so a packet loss on the link
causes uncertainty about the subsequent packets, which often must be
discarded until the decompressor has notified the compressor and the
compressor has sent re-synchronizing information. This typically
takes slightly more than a round-trip time. For links that combine
significant error rates with latencies that require multiple packets
to be in flight at a time, this leads to significant error
propagation, i.e. subsequent losses caused by an initial loss.
For links that are both high-latency (multiple packets in flight from
a typical RTP stream) and error-prone, RTP ROHC provides a more
robust way of RTP header compression, at a cost of higher complexity
at the compressor and decompressor. Within a talk spurt, only
extended losses of (depending on the mode chosen) 12 to 64 packets
typically cause error propagation.
Packet Reordering
The Internet architecture does not guarantee that packets will arrive
in the same order in which they were originally transmitted, and
transport protocols like TCP must take this into account.
But reordering does come at a cost with TCP as it is currently
defined. Because TCP returns a cumulative acknowledgment (ACK)
indicating the last in-order segment that has arrived, out-of-order
segments cause a TCP receiver to transmit a duplicate acknowledgment.
When the TCP sender notices three duplicate acknowledgments it
assumes that a segment was dropped by the network and uses the fast
retransmit algorithm [Jac90] [APS99] to resend the segment. In
addition, the congestion window is reduced by half, effectively
halving TCP's sending rate. If a subnetwork badly re-orders segments
such that three duplicate ACKs are generated, the TCP sender
needlessly reduces the congestion window and performance suffers.
Packet reordering does frequently occur in parts of the Internet, and
it seems to be difficult or impossible to eliminate [BPS99]. For
this reason, research has begun into improving TCP's behavior in the
face of packet reordering.
[BPS99] cites reasons why it may even be undesirable to eliminate
reordering. There are situations where average packet latency can be
reduced, link efficiency can be increased, and/or reliability can be
improved if reordering is permitted. Examples include certain high
speed switches within the Internet backbone and the parallel links
used over many Internet paths for load splitting and redundancy.
This suggests that subnetwork implementers should try to avoid packet
reordering whenever possible, but not if doing so compromises
efficiency, impairs reliability or increases average packet delay.
Note that every header compression scheme currently standardized for
the Internet requires in-order packet delivery on the link between
compressor and decompressor.
Mobility
Internet users are increasingly mobile. Not only are many Internet
nodes laptop computers, but pocket organizers and mobile embedded
systems are also becoming nodes on the Internet. These nodes may
connect to many different access points on the Internet over time,
and they expect this to be largely transparent to their activities.
Except when they are not connected to the Internet at all, and for
performance differences when they are connected, they expect that
everything will "just work" regardless of their current Internet
attachment point or local subnetwork technology.
Changing a host's Internet attachment point involves one or more of
the following steps.
First, if use of the local subnetwork is restricted, the user's
credentials must be verified and access granted. There are many ways
to do this. A trivial example would be an "Internet cafe" that grants
physical access to the subnetwork for a fee. Subnetworks may
implement technical access controls of their own; one example is IEEE
802.11 Wireless Equivalent Privacy [IEEE80211]. And it is common
practice for both cellular telephone and Internet service providers
(ISPs) to agree to serve each others users; RADIUS [RFC2865] is the
standard means for ISPs to exchange authorization information.
Second, the host may have to be reconfigured with IP parameters
appropriate for the local subnetwork. This usually includes setting
an IP address, default router, and domain name system (DNS) servers.
On multiple-access networks, the Dynamic Host Configuration Protocol
(DHCP) [RFC2131] is almost universally used for this purpose. On PPP
links, these functions are performed by the IP Control Protocol
(IPCP) [RFC1332].
Third, traffic destined for the mobile host must be routed to its
current location. This function is the most common meaning of the
term "Internet mobility".
Internet mobility can be provided at any of several layers in the
Internet protocol stack, and there is ongoing debate as to which are
the most appropriate and efficient. Mobility is already an feature of
certain application layer protocols; the Post Office Protocol (POP)
[RFC1939] and the Internet Message Access Protocol (IMAP) [RFC2060]
were created specifically to provide mobility in the receipt of
electronic mail.
Mobility can also be provided at the IP layer [RFC2002]. This
mechanism provides greater transparency, viz., IP addresses that
remain fixed as the nodes move, but at the cost of potentially
significant network overhead and increased delay because of the non-
optimum network routing and tunneling involved.
Some subnetworks may provide internal mobility, transparent to IP, as
a feature of their own internal routing mechanisms. To the extent
that these simplify routing at the IP layer, reduce the need for
mechanisms like Mobile IP, or exploit mechanisms unique to the
subnetwork, this is generally desirable. This is especially true when
the subnetwork covers a relatively small geographic area and the
users move rapidly between the attachment points within that area.
Examples of internal mobility schemes include Ethernet switching and
intra-system handoff in cellular telephony.
However, if the subnetwork is physically large and connects to other
parts of the Internet at multiple geographic points, care should be
taken to optimize the wide-area routing of packets between nodes on
the external Internet and nodes on the subnet. This is generally done
with "nearest exit" routing strategies. Because a given subnetwork
may be unaware of the actual physical location of a destination on
another subnetwork, it simply routes packets bound for the other
subnetwork to the nearest gateway between the two. This implies some
awareness of IP addressing and routing within the subnetwork. The
subnetwork may wish to use IP routing internally for wide area
routing and restrict subnetwork-specific routing to constrained
geographic areas where the effects of suboptimal routing are
minimized.
Multicasting
The Internet model includes "multicasting", where IP packets are sent
to all the members of a multicast group [RFC1112] [RFC2236]. IP
routers organize each multicast group into a spanning tree, and they
route multicast packets by making a copy for each output interface
that includes at least one downstream member of the multicast group.
Multicasting is considerably more efficient when a subnetwork
explicitly supports it. For example, a router relaying a multicast
packet onto an Ethernet segment need send only one copy, no matter
how many members of the multicast group are connected to the segment.
Without native Ethernet multicast support, the router would have to
transmit a separate copy of every multicast packet to every member of
the multicast group on the segment.
Subnetworks using shared channels (e.g., radio LANs, Ethernets, etc)
are especially suitable for native multicasting, and their designers
should make every effort to support it. This involves designating a
section of the subnetwork's own address space for multicasting and
designing receivers to accept packets addressed to some number of
multicast addresses in addition to the unicast packets specifically
addressed to them. How many multicast addresses are supported depends
on the requirements of the associated host or router; at least
several dozen will meet most current needs.
On low speed networks this address recognition function may be
readily implemented in host software, but on high speed networks it
should be implemented in subnetwork hardware. This hardware need not
be complete; for example, many Ethernet interfaces implement a
"hashing" function that passes all of the multicast (and unicast)
traffic to which the associated host subscribes, plus some small
fraction of multicast traffic to which the host does not subscribe.
Host software then only has to discard the relatively few unwanted
packets that make it past the hardware filter.
Broadcasting and Discovery
Link layers fall into two categories: point-to-point and shared. A
point-to-point link has exactly two endpoint components (hosts or
gateways); a shared link has more than two, either on an inherently
broadcast media (e.g., Ethernet, radio) or on a switching layer
hidden from the network layer (switched Ethernet, Myrinet, ATM).
Several Internet protocols which make use of link layer broadcast
capabilities, including link layer address lookup (ARP), auto-
configuration (RARP, BOOTP, DHCP), and routing (RIP). These protocols
require broadcast-capable links. Shared links SHOULD support native,
link layer subnet broadcast.
The lack of broadcast can impede the performance of these protocols,
or in some cases render them inoperable. ARP-like link address lookup
can be provided by a centralized database but at the expense of
potentially higher response latency and the need for nodes to have
explicit knowledge of the ARP server address.
Other protocols, such as DHCP, cannot function at all without a
subnetwork broadcast mechanism.
Routing
Many subnetworks provide their own internal routing mechanisms.
Since routing is the major function of the Internet layer, the
question naturally arises as to the proper division of function
between routing at the Internet layer and routing in the subnet.
In general, routing in a subnetwork and at IP is more complementary
than competitive. Routing algorithms often have difficulty scaling to
very large networks, and a division of labor between IP and a large
subnetwork can often make the routing problem more tractable for
both.
Some subnetworks have special features that allow the use of more
effective or responsive routing mechanisms that cannot be implemented
in IP because of its need for generality. One example is the self-
learning bridge algorithm widely used in Ethernet networks. Another
is the "handoff" mechanism in cellular telephone networks,
particularly the "soft handoff" scheme in IS-95 CDMA.
On the other hand, routing optimality can suffer when a subnetwork's
routing architecture hides internal structure that an IP router could
have used to make more efficient decisions. Such situations occur
most often when the subnetwork covers a large geographic area and
includes links of widely varying capacities, but presents itself to
IP as a single, fully-connected network with uniform metrics between
border nodes.
The subnetwork designer who decides to implement internal routing
should also consider whether a custom routing algorithm is warranted,
or if an existing Internet routing algorithm or protocol may suffice.
Routing algorithms and protocols can be notoriously subtle, complex
and difficult to implement correctly. Much work can be avoided if an
existing protocol or off-the-shelf product can be readily used.
Security Considerations
Security has become a high priority in the design and operation of
the Internet. The Internet is vast, and countless organizations and
individuals own and operate its various components. A consensus has
emerged for what might be called a "security placement principle": a
security mechanism is most effective when it is placed as close as
possible to, and under the direct control of the owner of, the asset
that it protects.
The most important conclusion that follows from this principle is
that end-to-end security (e.g., confidentiality, integrity and access
control) cannot be ensured with subnetwork security mechanisms. Not
only are end-to-end security mechanisms much more closely associated
with the end-user assets they protect, they are also much more
comprehensive. For example, end-to-end security mechanisms cover gaps
that can appear when otherwise good subnetwork mechanisms are
concatenated. This is an important application of the end-to-end
principle [SRC81].
Several security mechanisms that can be used end-to-end have already
been deployed in the Internet and are enjoying increasing use. The
most important are the Secure Sockets Layer (SSL) [SSL2] [SSL3] and
TLS [RFC2246] primarily used to protect web commerce; Pretty Good
Privacy (PGP) [RFC1991], primarily used to protect and authenticate
email and software distributions; the Secure Shell (SSH), used for
secure remote access and file transfer; and IPSEC [RFC2401], a
general purpose encryption and authentication mechanism that sits
just above IP and can be used by any IP application. (IPSEC can
actually be used either on an end-to-end basis or between security
gateways that do not include either or both end systems.)
Nonetheless, end-to-end security mechanisms are not used as widely as
might be desired. However, the group could not reach consensus on
whether subnetwork designers should be actively encouraged to
implement mechanisms to protect user data.
One point of view actively promotes subnetwork security mechanisms on
the principle that they can't hurt. The argument is that while
subnetwork security is admittedly inferior to end-to-end security,
many users are unwilling or unable to implement end-to-end security;
so subnetwork security is better than no security at all. This
viewpoint calls for subnetworks to implement mechanisms to achieve a
degree of security commensurate with a series of concatenated,
physically protected point-to-point links. This approach, termed
"wire-equivalent privacy" (WEP), is especially applicable to wireless
links, e.g., 802.11, and wired links where physical security is
impractical, e.g., cable TV networks.
Another view holds that subnetwork security mechanisms, especially
when weak or incorrectly implemented [BGW], may actually be
counterproductive. The argument is that subnetwork security
mechanisms can lull end users into a false sense of security,
diminish the incentive to deploy effective end-to-end mechanisms, and
encourage "risky" uses of the Internet that would not be made if
users understood the inherent limits of subnetwork security
mechanisms.
We therefore recommend that subnetwork vendors who choose to
implement security mechanisms to protect user data be as candid as
possible with the details of such security mechanisms and the
inherent limits of even the most secure mechanisms when implemented
in a subnetwork rather than on an end-to-end basis.
A much stronger case exists for another role for subnetwork security:
the protection of the subnetwork itself in keeping with the
"placement principle". Possible threats to subnetwork assets include
theft of service and denial of service; shared media subnets tend to
be especially vulnerable to such attacks. In some cases, mechanisms
that protect subnet assets can also improve (but NOT ensure) end-to-
end security.
Subnetwork designers must keep in mind that design and implementation
for security is difficult [Schneier2][Schneier3]. [Schneier1]
describes protocols and algorithms which are considered well
understood and believed to be sound.
Poor design process, subtle design errors and flawed implementation
can result in gaping vulnerabilities. In recent years, a number of
subnet standards have had problems exposed. The following are
examples of the mistakes that have been made:
1. Use of weak and untested algorithms [Crypto9912], [BGW]. For a
variety of reasons, algorithms were chosen which had subtle flaws
that made them vulnerable to a variety of attacks.
2. Use of 'security by obscurity' [Schneier4], [Crypto9912]. One
common mistake is to assume that keeping cryptographic algorithms
secret makes them more secure. This is intuitive, but wrong. Full
public disclosure early in the design process attracts peer review by
knowledgeable cryptographers. Exposure of flaws by this review far
outweighs any imagined benefit from forcing attackers to reverse
engineer security algorithms.
3. Inclusion of trapdoors [Schneier4], [Crypto9912]. Trapdoors are
flaws surreptitiously left in an algorithm to allow it to be broken.
This might be done to recover lost keys or to permit surreptitious
access by governmental agencies. Trapdoors can be discovered and
exploited by malicious attackers.
4. Sending passwords or other identifying information as clear text.
For many years, analog cellular telephones could be cloned and used
to steal service. The cloners merely eavesdropped on the
registration protocols that exchanged everything in clear text.
5. Keys which are common to all systems on a subnet [BGW].
6. Incorrect use of a sound mechanism. For example [BGW], one subnet
standard includes an initialization vector which is poorly designed
and poorly specified. A determined attacker can easily recover
multiple ciphertexts encrypted with the same key stream and perform
statistical attacks to decipher them.
7. Identifying information sent in clear text that can be resolved to
an individual, identifiable device. This creates a vulnerability to
attacks targeted to that device (or its owner).
8. Inability to renew and revoke shared secret information.
9. Insufficient key length [Blaze96].
10. Failure to address "man-in-the-middle" attacks, e.g., with mutual
authentication.
This list is by no means comprehensive. Design problems are
difficult to avoid, but expert review is generally invaluable in
avoiding problems.
In addition, well designed security protocols can be compromised by
implementation defects. Examples of such defects include use of
predictable pseudo-random numbers [RFC1750], vulnerability to buffer
overflow attacks due to unsafe use of certain I/O system calls
[WFBA2000], and inadvertent exposure of secret data.
References
References of the form RFCnnnn are Internet Request for Comments
(RFC) documents available online at www.rfc-editor.org.
[APS99] Mark Allman, Vern Paxson, W. Richard Stevens. "TCP
Congestion Control". April 1999. RFC 2581.
[ATMFTM] The ATM Forum, "Traffic Management Specification, Version
4.0", April 1996, document af-tm-0056.000 (www.atmforum.com).
[BGW] Nikita Borisov, Ian Goldberg, David Wagner, "Security of the
WEP Algorithm" www.www.isaac.cs.berkeley.edu/isaac/wep-faq.html
[Blaze96] M. Blaze, W. Diffie, R. Rivest, B. Schneier, T. Shimomura,
E. Thompson, M. Weiner, "Minimal Key Lengths for Symmetric Ciphers to
Provide Adequate Commercial Security", available at
http://www.counterpane.com/keylength.html.
[BPK98] Hari Balakrishnan, Venkata Padmanabhan, Randy H. Katz. 'The
Effects of Asymmetry on TCP Performance." ACM Mobile Networks and
Applications (MONET), 1998.
[BPS99] "Packet Reordering is Not Pathological Network Behavior", Jon
C. R. Bennet, Craig Partridge, Nicholas Shectman, IEEE/ACM
Transactions on Networking, Vol 7, No. 6, December 1999.
[Crypto9912] Schneier, Bruce "European Cellular Encryption
Algorithms" Crypto-Gram (December 15, 1999) www.counterpane.com
[DIX] Digital Equipment Corp, Intel Corp, Xerox Corp, Ethernet Local
Area Network Specification Version 2.0, November 1982.
[DOCSIS1] Data-Over-Cable Service Interface Specifications, Radio
Frequency Interface Specification 1.0, SP-RFI-I05-991105, November
1999, Cable Television Laboratories, Inc.
[DOCSIS2] Data-Over-Cable Service Interface Specifications, Radio
Frequency Interface Specification 1.1, SP-RFIv1.1-I05-000714, July
2000, Cable Television Laboratories, Inc.
[DOCSIS3] W.S. Lai, "DOCSIS-Based Cable Networks: Impact of Large
Data Packets on Upstream Capacity", 14th ITC Specialists Seminar on
Access Networks and Systems, Barcelona, Spain, April 25-27, 2001.
[EN301] ETSI, (European Broadcasting Union), Digital Video
Broadcasting (DVB); DVB Specification for Data Broadcasting, 1997.
Draft ETSI Standard EN 301 192 v1.1.1 (August 1997).
[FB00] Firoiu V., and Borden M., "A Study of Active Queue Management
for Congestion Control" to appear in Infocom 2000
[IEEE8023] IEEE 802.3 CSMA/CD Access Method. Available from
http://standards.ieee.org/catalog/IEEE802.3.html.
[IEEE80211] IEEE 802.11 Wireless LAN standard. Available from
http://standards.ieee.org/catalog/IEEE802.11.html.
[ISO3309] ISO/IEC 3309:1991(E), "Information Technology -
Telecommunications and information exchange between systems - High-
level data link control (HDLC) procedures - Frame structure",
International Organization For Standardization, Fourth edition
1991-06-01.
[ISO13181] ISO/IEC, ISO/IEC 13181-1: Information Technology - Generic
coding of moving pictures and associated audio information, 1995,
International Organization for Standardization and International
Electrotechnical Commission.
[Jac90] Van Jacobson. Modified TCP Congestion Avoidance Algorithm.
Email to the end2end-interest mailing list, April 1990. URL:
ftp://ftp.ee.lbl.gov/email/vanj.90apr30.txt.
[LK00] R. Ludwig, R. H. Katz, "The Eifel Algorithm: Making TCP Robust
Against Spurious Retransmissions", ACM Computer Communication Review,
Vol. 30, No. 1, January 2000.
[LKJK01] R. Ludwig, A. Konrad, A. D. Joseph, R. H. Katz, "Optimizing
the End-to-End Performance of Reliable Flows over Wireless Links", To
appear in ACM/Baltzer Wireless Networks Journal (Special issue:
Selected papers from ACM/IEEE MOBICOM 99), available at
http://iceberg.cs.berkeley.edu/publications.html.
[MBB00] May, M., Bonald, T., and Bolot, J-C., "Analytic Evaluation of
RED Performance" to appear INFOCOM 2000
[MBDL99] May, M., Bolot, J., Diot, C., and Lyles, B., "Reasons not to
deploy RED", technical report, June 1999.
[MG01] S. Mascolo, A. Grieco, "Easy RED: Improving Fairness via
Simple Early Discard", to appear in IEEE Communication Letters.
Available as http://www-ictserv.poliba.it/mascolo/papers/Easy_red.pdf
[MSMO97] M. Mathis, J. Semke, J. Mahdavi, T. Ott, "The Macroscopic
Behavior of the TCP Congestion Avoidance Algorithm", Computer
Communication Review, volume 27, number 3, July 1997.
[OKM96] T. Ott, J.H.B. Kemperman, M. Mathis, "The Stationary Behavior
of Ideal TCP Congestion Avoidance".
ftp://ftp.bellcore.com/pub/tjo/TCPwindow.ps
[PFTK98] Padhye, J., Firoiu, V., Towsley, D., and Kurose, J.,
"Modeling TCP Throughput: a Simple Model and its Empirical
Validation", UMASS CMPSCI Tech Report TR98-008, Feb. 1998.
[RED93] S. Floyd, V. Jacobson, "Random Early Detection gateways for
Congestion Avoidance", IEEE/ACM Transactions in Networking, V.1 N.4,
August 1993, http://www.aciri.org/floyd/papers/red/red.html
[RF95] Romanow, A., and Floyd, S., "Dynamics of TCP Traffic over ATM
Networks". IEEE JSAC, V. 13 N. 4, May 1995, p. 633-641.
[RFB01] Ramakrishnan, K.K., Floyd, S., and Black, D. "The Addition
of Explicit Congestion Notification (ECN) to IP", Internet draft,
work in progress, June 2001, available as
http://search.ietf.org/internet-drafts/draft-ietf-tsvwg-ecn-04.txt.
This draft is intended to supersede RFC 2481, draft-ietf-tsvwg-tcp-
ecn-00.txt, draft-ietf-tsvwg-ecn-tunnels-00.txt, and draft-ietf-
ipsec-ecn-02.txt. The IESG approved this document to go to Proposed
Standard on June 12, 2001.
[RFC791] Jon Postel. "Internet Protocol". September 1981.
[RFC1144] Jacobson, V., "Compressing TCP/IP Headers for Low-Speed
Serial Links," RFC 1144, February 1990.
[RFC1191] J. Mogul, S. Deering. "Path MTU Discovery". November 1990.
[RFC1435] S. Knowles. "IESG Advice from Experience with Path MTU
Discovery". March 1993.
[RFC1577] M. Laubach. "Classical IP and ARP over ATM". January 1994.
[RFC1661] W. Simpson. "The Point-to-Point Protocol (PPP)". July 1994.
[RFC1750] D. Eastlake, S. Crocker, J. Schiller, "Randomness
Recommendations for Security", December 1994 (this will be obsoleted
soon... check before submitting to RFC editor)
[RFC1981] J. McCann, S. Deering, J. Mogul. "Path MTU Discovery for IP
version 6". August 1996.
[RFC2246] T. Dierks, C. Allen. "The TLS Protocol Version 1.0".
January 1999.
[RFC2364] G. Gross et al. "PPP Over AAL5". July 1998.
[RFC2393] A. Shacham et al. "IP Payload Compression Protocol
(IPComp)". December 1998.
[RFC2394] R. Pereira. "IP Payload Compression Using DEFLATE".
December 1998.
[RFC2395] R. Friend, R. Monsour. "IP Payload Compression Using LZS".
December 1998.
[RFC2440] J. Callas et al. "OpenPGP Message Format". November 1998.
[RFC2481] Ramakrishan, K. and Floyd S., "A Proposal to add Explicit
Congestion Notification (ECN) to IP" RFC2481 January 1999
[RFC2507] M. Degermark, B. Nordgren, S. Pink. "IP Header
Compression". February 1999.
[RFC2508] S. Casner, V. Jacobson. "Compressing IP/UDP/RTP Headers for
Low-Speed Serial Links". February 1999.
[RFC2581] M. Allman, V. Paxson, W. Stevens. "TCP Congestion Control".
April 1999.
[RFC2406] S. Kent, R. Atkinson. "P Encapsulating Security Payload
(ESP)". November 1998.
[RFC2616] R. Fielding et al. "Hypertext Transfer Protocol --
HTTP/1.1". June 1999.
[RFC2684] D. Grossman, J. Heinanen. "Multiprotocol Encapsulation over
ATM Adaptation Layer 5". September 1999.
[RFC2686] C. Bormann, "The Multi-Class Extension to Multi-Link PPP",
September 1999.
[RFC2687] C. Bormann, "PPP in a Real-time Oriented HDLC-like
Framing", September 1999.
[RFC2689] C. Bormann, "Providing Integrated Services over Low-bitrate
Links", September 1999.
[RFC3095] C. Bormann, ed., C. Burmeister, M. Degermark, H. Fukushima,
H. Hannu, L-E. Jonsson, R. Hakenberg, T. Koren, K. Le, Z. Liu, A.
Martensson, A. Miyazaki, K. Svanbro, T. Wiebke, T. Yoshimura, H.
Zheng, "RObust Header Compression (ROHC): Framework and four
profiles: RTP, UDP, ESP, and uncompressed", July 2001.
[Schneier1] Schneier, Bruce, Applied Cryptography: Protocols,
Algorithms and Source Code in C (John Wiley and Sons, October 1995).
[Schneier2] Schneier, Bruce, Secrets and Lies: Digital Security in a
Networked World (John Wiley & Sons, August 2000).
[Schneier3] Schneier, Bruce "Why Cryptography is Harder Than it
Looks", www.counterpane.com
[SP2000] "When the CRC and TCP Checksum Disagree", Jonathan Stone &
Craig Partridge, ACM CCR p309-321, September 2000,
http://www.acm.org/sigcomm/sigcomm2000/conf/paper/sigcomm2000-9-1.pdf.
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http://people.qualcomm.com/karn/library.html.
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Authors' Addresses:
Phil Karn, editor (karn@qualcomm.com)
Carsten Bormann (cabo@tzi.org)
Gorry Fairhurst (gorry@erg.abdn.ac.uk)
Aaron Falk (afalk@panamsat.com)
Dan Grossman (dan@dma.isg.mot.com)
Reiner Ludwig (Reiner.Ludwig@ericsson.com)
Jamshid Mahdavi (mahdavi@novell.com)
Saverio Mascolo (mascolo@poliba.it)
Gabriel Montenegro (Gabriel.Montenegro@eng.sun.com)
Marie-Jose Montpetit (marie@teledesic.com)
Joe Touch (touch@isi.edu)
Lloyd Wood (lwood@cisco.com)
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