Internet DRAFT - draft-ietf-pilc-link-design
Internet Engineering Task Force Phil Karn, editor
INTERNET DRAFT Qualcomm
Universitaet Bremen FB3 TZI
Godred (Gorry) Fairhurst
University of Aberdeen
Sun Microsystems Laboratories, Europe
File: draft-ietf-pilc-link-design-15.txt December, 2003
Expires: June, 2004
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-
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
The list of Internet-Draft Shadow Directories can be accessed at
This document provides advice to the designers of digital
communication equipment, link-layer protocols and packet-switched
local networks (collectively referred to as subnetworks) who wish to
support the Internet protocols but who may be unfamiliar with the
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
This document would not have been possible without the contributions
of a great number of people in the Performance Implications of Link
Characteristics Working Group. In particular, the following people
provided major contributions of text, editing and advice to this
document: Mark Allman provided the final editing to complete this
document. Carsten Bormann provided text on robust header
compression. Gorry Fairhurst provided text on broadcast and
multicast issues and many valuable comments on the entire document.
Aaron Falk provided text on bandwidth on demand. Dan Grossman
provided text on security considerations as well as on many facets of
the document. Reiner Ludwig provided thorough document review and
text on TCP vs. Link-Layer Retransmission. Jamshid Mahdavi provided
text on TCP performance calculations. Saverio Mascolo provided
feedback on the document. Gabriel Montenegro provided feedback on
the document. Marie-Jose Montpetit provided text on bandwidth on
demand. Joe Touch provided text on multicast and broadcast. and
Lloyd Wood provided many valuable comments on drafts of the document.
Table of Contents
1 Introduction and Overview
2 Maximum Transmission Units (MTUs) and IP Fragmentation
2.1 Choosing the MTU in Slow Networks
3 Framing on Connection-Oriented Subnetworks
4 Connection-Oriented Subnetworks
5 Broadcasting and Discovery
7 Bandwidth on Demand (BoD) Subnets
8 Reliability and Error Control
8.1 TCP vs Link-Layer Retransmission
8.2 Recovery from Subnetwork Outages
8.3 CRCs, Checksums and Error Detection
8.4 How TCP Works
8.5 TCP Performance Characteristics
8.5.1 The Formulae
8.5.3 Analysis of Link-Layer Effects on TCP Performance
9 Quality-of-Service (QoS) considerations
10 Fairness vs Performance
11 Delay Characteristics
12 Bandwidth Asymmetries
13 Buffering, flow & congestion control
15 Packet Reordering
18 Security Considerations
1 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. A subnetwork
refers to any network operating immediately below the IP layer to
connect two or more systems using IP (i.e., end hosts or routers).
In its simplest form, this may be a direct connection between the IP
systems (e.g., using a length of cable or over a wireless medium).
This document defines a subnetwork as a layer 2 network, which is a
network that does not rely upon the services of IP routers to forward
packets between parts of the subnetwork. Although, IP routers may
bridge frames at layer 2 between parts of a subnetwork. Sometimes, it
is convenient to aggregate a group of such subnetworks into a single
logical subnetwork. IP routing protocols (e.g., OSPF, IS-IS, and PIM)
can be configured to support this aggregation, but typically present
a layer-3 subnetwork rather than a layer-2 subnetwork. This may also
result in a specific packet passing several times over the same
layer-2 subnetwork via an intermediate layer-3 gateway (router).
Because that aggregation requires layer-3 components, issues thereof
are beyond the scope of this document.
However, 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 packets comprise 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
Optimizing a subnetwork for IP involves three complementary
1. Providing functionality sufficient to carry IP.
2. Eliminating unnecessary functions that increase cost or
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. That is to say, subnetwork designers make many more
errors of commission than errors of omission. However, certain
enhancements to Internet features, such as multicasting and quality-
of-service, benefit significantly from support given by 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]. That
paper argued that many functions can be implemented properly *only*
on an end-to-end basis, i.e., at the highest 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.
Such 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 original architecture of the Internet was influenced by the end-
to-end principle, and in our view it has been part of the reason for
the Internet's success.
The remainder of this document discusses the various subnetwork
design issues that the authors consider relevant to efficient IP
2 Maximum Transmission Units (MTUs) and IP Fragmentation
IPv4 packets (datagrams) vary in size from 20 bytes (the size of the
IPv4 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 IP 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
[DIX82] (not IEEE 802.3 [IEEE8023]) header, which lacks a length
field to indicate the true data length when the packet is padded to a
minimum of 60 bytes. This is not a problem for uncompressed IP
because each IP packet 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 because it is needed for the reconstruction of
the original header.
In IP version 4 (the version now in widespread use), fragmentation
can occur at either the sending host or in an intermediate router,
and fragments can be further fragmented at subsequent routers if
In IP version 6 [RFC 2460], fragmentation can occur only at the
sending host; it cannot occur in a router (called "router
fragmentation" in this document).
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 and IPv6.
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 router
fragmentation is best avoided. This requires that subnetworks support
MTUs that are "reasonably" large. The smallest MTU permitted in IPv4
by [RFC791] is 576 bytes, but such a small value would clearly be
inefficient. Because IPv6 omits fragmentation by routers, [RFC 2460]
specifies a larger minimum MTU of 1280 bytes. Any subnetwork with an
internal packet payload smaller than 1280 bytes must implement a
mechanism that performs fragmentation/reassembly of IP packets
to/from subnetwork frames 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 the ubiquity of
Ethernet few Internet paths have MTUs larger than this value. This
severely limits the utility of larger MTUs provided by other
subnetworks. Meanwhile 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 GRE [RFC2784] or IP Security in
tunnel mode [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 host performing IPsec encapsulation. Tunneling has made
it more difficult to avoid router fragmentation and has increased the
incidence of path MTU black holes [RFC2401], [RFC2923]. Larger
subnetwork MTUs may help to alleviate this problem.
2.1 Choosing the MTU in Slow Networks
In slow networks, the largest possible packet may take a considerable
time to send. This is known as channelisation or serialisation
delay. Total end-to-end 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, serialisation 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) size. 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 router fragmentation (if Path MTU discovery is
not in use). This is likely the case with multicast.
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
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
ATM (asynchronous transfer mode), where SNDUs are fragmented and
interleaved across smaller 53-byte ATM cells, 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 cell overhead to each 48-byte ATM cell.
A third example is Data-Over-Cable Service Interface Specifications
(DOCSIS) with typical upstream bandwidths of 2.56 Mb/s or 5.12 Mb/s.
To reduce the impact of a 1500-byte MTU in DOCSIS 1.0 [DOCSIS1], a
data link layer fragmentation mechanism is specified in DOCSIS 1.1
[DOCSIS2]. To accommodate the installed base, DOCSIS 1.1 must be
backward compatible with DOCSIS 1.0 cable modems, which generally do
not support fragmentation. Under the co-existence of DOCSIS 1.0 and
DOCSIS 1.1, the unfragmented large data packets from DOCSIS 1.0 cable
modems may affect the quality of service for voice packets from
DOCSIS 1.1 cable modems. In this case, it has been shown in [DOCSIS3]
that use of bandwidth allocation algorithms can mitigate this effect.
To summarize, there is a fundamental tradeoff between efficiency and
latency in the design of a subnetwork, and the designer should keep
this tradeoff in mind.
3 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;
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], which
uses a variant of HDLC, is applicable to bit synchronous, octet
synchronous and octet asynchronous links (i.e., examples 1-3 above).
PPP is one prefered framing method for IP, since a large number of
systems interoperate with PPP. 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. However, 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
Because IP packets are of 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 the data segment 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. Alternatively for subnetworks
that do not reorder the fragments of A SNDU, only the last SNDU of
the packet could be marked, as this would implicitly indicate 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
In AAL5, the ATM User-User Indication, which is encoded in the
Payload Type field of an ATM cell, indicates the end cell of a
packet. The packet trailer is located at the end of the SNDU and
contains the packet length and a CRC.
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] [ISO13818] 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
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. This may have
implications when frames are lost. 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 flag-
like sequences occurring in the packet must be escaped, which expands
them). The designer must also weigh implementation complexity and
performance in the choice and design of an adaptation layer.
4 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 [RFC793], 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. However, 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 the subnetwork 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
Purely connectionless subnets (such as Ethernet), which have no state
and dynamically share resources, are optimal to supporting best-
effort IP, which is stateless and dynamically shares resources.
Connection-oriented packet networks (such as ATM and Frame Relay),
which have state and dynamically share resources, are less optimal,
since best effort IP does not benefit from the overhead of creating
and maintaining state. Connection-oriented circuit switched networks
(including the PSTN and ISDN) both have state and statically allocate
resources for a call, and thus not only require state creation and
maintenance overhead, but also do not benefit from the efficiencies
of statistical multiplexing sharing of capacity inherent in IP.
In any event, if an SNDCF that opens and closes subnet connections is
used to support IP, care should be taken to make sure that call
processing in the subnet can keep up with relatively short holding
5 Broadcasting and Discovery
Subnetworks fall into two categories: point-to-point and shared. A
point-to-point subnet has exactly two endpoint components (hosts or
routers); a shared link has more than two, using either an inherent
broadcast medium (e.g., Ethernet, radio) or that are on a switching
layer hidden from the network layer (e.g., switched Ethernet, Myrinet
[MYR95], ATM). Switched subnetworks handle broadcast by copying
broadcast packets to give to each interface that supports one, or
more, systems (hosts or routers) a copy of each packet.
Several Internet protocols for IPv4 make use of broadcast
capabilities, including link-layer address lookup (ARP), auto-
configuration (RARP, BOOTP, DHCP), and routing (RIP).
A lack of broadcast capability can impede the performance of these
protocols, or render them inoperable (e.g. DHCP). 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. Shared links
should support native, link-layer subnet broadcast.
A corresponding set of IPv6 protocols uses multicasting (see next
section) instead of broadcast to provide similar functions with
improved scaling in large networks.
The Internet model includes "multicasting", where IP packets are sent
to all the members of a multicast group [RFC1112] [RFC2236].
Multicast is an option in IPv4, but a standard feature of IPv6. IPv4
multicast is currently used by multimedia, teleconferencing, gaming,
and file distribution (web, peer-to-peer sharing) applications, as
well as by some key network and host protocols (e.g., RIPv2, OSPF,
NTP). IPv6 additionally relies on multicast for network
configuration (DHCP-like autoconfiguration) and link-layer address
discovery [RFC2461] (replacing ARP). In the case of IPv6 this can
allow autoconfiguration and address discovery to span across routers,
whereas the IPv4 broadcast-based services cannot without ad-hoc
router support [RFC1812].
Multicast enabled IP routers organize each multicast group into a
spanning tree, and route multicast packets by making a copy of each
multicast packet and forwards the copies to 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 subnet need send only one copy of the packet,
no matter how many members of the multicast group are connected to
the segment. Without native multicast support, routers and switches
on shared links would need to use broadcast with software filters,
such that every multicast packet sent incurs software overhead for
every node on the subnetwork, even if a node is not a member of the
multicast group. Alternately, the router would transmit a separate
copy to every member of the multicast group on the segment, as is
done on multicast-incapable switched subnets.
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. On
these networks, multicast is basically broadcast on the medium, with
Layer-2 receiver filters.
Subnet interfaces also need to be designed to accept packets
addressed to some number of multicast addresses in addition to the
unicast packets specifically addressed to them. How many multicast
addresses need to be supported by a host depends on the requirements
of the associated host; at least several dozen will meet most current
On low-speed networks the multicast 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 where the IP layer receives 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/router software then has to discard the unwanted
packets that pass the Layer-2 multicast address filter [RFC1112].
There does not need to be a one-to-one mapping between a layer 2
multicast address and an IP multicast address. An address overlap may
significantly degrade the filtering capability of a receiver's
hardware multicast address filter. A subnetwork supporting only
broadcast should use this service for multicast and must rely on
Switched subnetworks must also provide a mechanism for copying
multicast packets to ensure the packets reach at least all members of
a multicast group. One option is to "flood" multicast packets, in
the same manner as broadcast. This can lead to unnecessary
transmissions on some subnetwork links (notably non-multicast-aware
ethernet switches). Some subnetworks therefore allow multicast filter
tables to control which links receive packets belonging to a specific
group. To configure this automatically requires access to layer 3
group membership information (e.g., IGMP). Various implementation
options currently exist to provide a subnet node with a list of
multicast addresses to port/interface mappings [MBONED-GAP]. These
employ a range of approaches, including signaling from end hosts
(e.g. IEEE 802 GARP/GMRP [802.1p]), signaling from switches (e.g.
CGMP [CGMP] and RGMP [RFC3488]), interception and proxy of IP group
membership packets (e.g. IGMP/MLD Proxy [MAGMA-PROXY]), and enabling
Layer 2 devices to snoop/inspect/peek into forwarded Layer 3 protocol
headers (e.g. IGMP, MLD, PIM) so that they may infer L3 multicast
group membership. These approaches differ in their complexity,
flexibility and ability to support new protocols.
7 Bandwidth on Demand (BoD) Subnets
Some subnets allow a number of subnet nodes to share a channel
efficiently by assigning transmission opportunities dynamically.
Transmission opportunities are requested by a subnet node when it has
packets to send. The subnet schedules and grants transmission
opportunities sufficient to allow the transmitting subnet node to
send one or more packets (or packet fragments). We call these
subnets Bandwidth on Demand (BoD) subnets. Examples of BoD subnets
include Demand Assignment Multiple Access (DAMA) satellite and
terrestrial wireless networks, IEEE 802.11 point coordination
function (PCF) mode, and DOCSIS. A connection-oriented network (like
the PSTN, ATM or Frame Relay) reserves resources on a much longer
timescale, and is therefore not a BoD subnet in our taxonomy.
The design parameters for BoD are similar to those in connection
oriented subnetworks, although the implementations may vary
significantly. 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, for 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. Examples of the resource that may be shared
include a terrestrial wireless hop, a cable modem uplink, a satellite
uplink, and an end-to-end satellite channel.
Long-delay BoD subnets pose problems similar to connection-oriented
networks in anticipating traffic. While connection-oriented subnets
that expect new data to arrive hold idle channels open, 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
There is a wide variety of BoD protocols available. However, there
has been relatively little comprehensive research on the interactions
between BoD mechanisms and Internet protocol performance. Research
on some specific mechanisms is available (e.g., [AR02]). One item
that has been studied is TCP's retransmission timer [KY02]. BoD
systems can cause spurious timeouts when adjusting from a relatively
high data rate to a relatively low data rate. In this case, TCP's
transmitted data takes longer to get through the network than
predicted by the TCP sender's computed retransmission timeout and
therefore the TCP sender is prone to resending a segment prematurely.
8 Reliability and Error Control
In the Internet architecture, the ultimate responsibility for error
recovery is at the end points [SRC81]. 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 as the
transport protocol) must recover from these errors on an end-to-end
basis. Error recovery in the subnetwork is therefore justifiable
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.
8.1 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.
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).
When enough ECC information is available to permit the receiver to
correct some 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
Unlike ARQ, FEC was seldom used for telecommunications outside of
space links prior to 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 (in the upstream direction) with V.42 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.42.
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. However, 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.
Only end-to-end error recovery can ensure a reliable service to the
application (see Section 8). However, some subnetworks (e.g., many
wireless links) also require link-layer error recovery as a
performance enhancement [RFC3366]. 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
layer 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 (parts of) a packet, the
link latency momentarily increases. Since TCP bases its
retransmission timeout on prior measurements of total 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. As a result, the link layer 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 [RFC3366]. 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
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 [RFC3366].
However, when a link layer can identify individual flows and apply
ARQ selectively [LKJK02], 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 for expected TCP traffic on its end-to-end paths
across the subnet balance through analysis and simulation.
8.2 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. Furthermore, 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, the
subnetwork 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 outage ends. When packets are discarded, IP
should be notified so that the appropriate ICMP messages can be sent.
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 -- reliability, as discussed earlier, can only be ensured
on an end-to-end basis.
Only a few packets 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 [RFC3366].
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 will be
retained during an outage.
8.3 CRCs, Checksums and Error Detection
The TCP [RFC793], UDP [RFC768], ICMP, and IPv4 [RFC791] protocols all
use the same simple 16-bit 1's complement checksum algorithm
[RFC1071] to detect corrupted packets. The IPv4 header checksum
protects only the IPv4 header, while the TCP, ICMP, and UDP checksums
provide end-to-end error detection for both the transport pseudo
header (including network and transport layer information) and the
transport payload data. Protection of the data is optional for
applications using UDP [RFC768] for IPv4, but is required for IPv6.
The Internet checksum is not very strong from a coding theory
standpoint, but it is easy to compute in software, and various
proposals to replace the Internet checksums with stronger checksums
have failed. However, it is known that undetected errors can and do
occur in packets received by end hosts [SP2000].
To reduce processing costs, 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 (pseudo header) 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 within the network; this was considered a reasonable price
to pay for a considerable reduction in the processing required by
each router, and it was assumed that subnetworks would use a strong
One way to provide additional protection for an IPv4 or IPv6 header
is by the authentication and packet integrity services of the IP
Security (IPsec) protocol [RFC2401]. However, this may not be a
choice available to the subnetwork designer.
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 for error detection. This is
significantly more robust against most patterns of errors than the
16-bit Internet checksum. However, not that the error detection
properties of a specific CRC code diminish with increasing frame
size. The Point-to-Point Protocol [RFC1662] requires support of a
16-bit CRC for each link frame, with a 32-bit CRC as an option. (PPP
is often used in conjunction with a dialup modem, which can 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. Many subnetworks can also use other mechanisms to enhance the
error detection capability of the link CRC (e.g., FEC in dialup
modems, mobile radio and satellite channels).
Any new subnetwork designed to carry IP should therefore provide
error detection for each IP packet that is at least as strong as the
32-bit CRC specified in [ISO3309]. While this will achieve a very
low undetected packet error rate due to transmission errors, 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 to dealing with corrupted packets [SRC81].
Packet corruption may be, and is, also caused by bugs in host and
router hardware and software. Even if every subnetwork implemented
strong error detection, it is still essential that end-to-end
checksums are used at the receiving end host [SP2000].
Designers of complex subnetworks consisting of internal links and
packet switches should consider implementing error detection on an
edge-to-edge basis to cover an entire SNDU (or IP packet). A CRC
would be generated at the entry point to the subnetwork and checked
at the exit endpoint. This may be used instead of, or in combination
with, error detection at the interface to each physical link. An
edge-to-edge check has the significant advantage of protecting
against errors introduced anywhere within the subnetwork, not just
within its transmission links. Examples of this approach include the
way in the Ethernet CRC-32 is handled by LAN bridges [802.1D]. ATM
AAL5 [ITU-I363] also uses an edge-to-edge CRC-32.
Some specific applications may be tolerant of residual errors in the
data they exchange, but removal of the link CRC may expose the
network to an undesirable increase in undetected errors in the IP and
transport headers. Applications may also require a high level of
error protection for control information exchanged by protocols
acting above the transport layer. One example is a voice codec which
is robust against bit errors in the speech samples. For such
mechanisms to work, the receiving application must be able to
tolerate receiving corrupted data. This also requires that an
application uses a mechanism to signal payload corruption is
permitted and to indicate the coverage (headers and data) that is
required to be protected by the subnetwork CRC. Currently there is
no Internet standard for supporting partial payload protection.
Receipt of corrupt data by arbitrary application protocols carries a
serious danger that a subnet delivers data with errors which remain
undetected by the application and hence corrupt the communicated data
8.4 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 [RFC2309,RFC2914], but it is still widely
TCP uses sequence numbering and acknowledgments (ACKs) on an end-to-
end basis to provide reliable, sequenced delivery. TCP ACKs are
cumulative, i.e., each implicitly ACKs every segment received so far.
If a packet with an unexpected sequence number is received, the ACK
field in the packets returned by the receiver will cease to advance.
Using an optional enhancement, TCP can send selective acknowledgments
(SACKs) [RFC2018] to indicate which segments have arrived at the
Since the most common cause of packet loss is congestion, TCP treats
packet loss as a potential indication of Internet congestion along
the path between TCP endhosts. This happens automatically, and the
subnetwork need not know anything about IP or TCP. A subnetwork node
simply drops packets whenever it must, though some packet-dropping
strategies (e.g., RED) are more fair to competing flows than others.
TCP recovers from packet losses in two different ways. The most
important mechanism 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 newly-received
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-
A new technique called "Explicit Congestion Notification" (ECN)
[RFC3168] allows routers to directly signal congestion to hosts
without dropping packets. This is done by setting a bit in the IP
header. Since ECN support is likely to remain optional, 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 capacity.
TCP probes for available network capacity by initially setting cwnd
to one or two packets and then increasing cwnd 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.
8.5 TCP Performance Characteristics
Here we present a 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 a 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 differ from 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.
8.5.1 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 (i.e., the TCP specified
in [RFC2581]) is given by Padhye, et al [PFTK98]. This formula is:
BW = --------------------------------------------------------
RTT*sqrt(1.33*p) + RTO*p*[1+32*p^2]*min[1,3*sqrt(.75*p)]
BW is the maximum TCP throughout achievable by an
individual TCP flow
MSS is the TCP 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
BW = C --- -------
where C is 0.93.
Both formulae assume that the TCP Receiver Window is not limiting the
performance of the connection. Because the 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. However, 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 arithmetic 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
that 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
Note that these formulae break down as p approaches 100%.
8.5.3 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 80 msec, you could compute an upper
bound on the performance as follows:
For MSS, use 1000 bytes to exclude the 40 bytes of minimum IPv4 and
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
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
2. The equations for TCP performance are all expressed in terms of
packet loss, but many subnetwork designers think in terms of bit-
error ratio. *If* channel bit errors are independent, then the
probability of a packet being corrupted is:
p = 1 - ([1 - BER]^[FRAME_SIZE*8])
Here we assume FRAME_SIZE is in bytes and "^" represents
exponentiation. It includes the user data and all headers (TCP,IP and
subnetwork). (Note: this analysis assumes the subnetwork does not
perform ARQ or transparent fragmentation [RFC3366].) If the
BER * [FRAME_SIZE*8] << 1
holds, the packet loss probability p can be approximated by:
p = BER * [FRAME_SIZE*8]
These equations can be used to apply BER to the performance equations
Note that FRAME_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
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
If the distribution of errors 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 are needed 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 a
variety of RTT values are experienced in the Internet.
For example, RTTs are typically less than 10 msec in a wired LAN
environment when communicating with a local host. 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 (serialisation) times.
Links over geostationary repeater satellites have one-way speed-of-
light delays of around 250ms: a minimum of 125ms propagation delay 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
Queues on heavily-congested links may back up, increasing RTTs.
Finally, virtual private networks (VPNs) and other forms of
encryption and tunneling can add significant end-to-end delay to
9 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 single mechanism 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 a current 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.
Capacity is reserved for the flow, and appropriate traffic
conditioning and scheduling is installed in routers along the path.
The ReSerVation Protocol (RSVP) [RFC2205, RFC2210] is usually, but
need not necessarily be, used to install the flow QoS state. Intserv
defines two services, in addition to the Default (best effort)
-- 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. Performance guarantees cannot be made unless
this state is present in every router along the path. 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. Diffserv
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 Type-of-
Service (ToS) byte [RFC791] 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 assurance for each aggregated flow of packets.
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. 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 dedicated
- 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
- 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 tradeoffs 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 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.
10 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
subnet node may transmit per contention cycle helps assure timely
access to the channel for each subnet node, 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
subnet node while maximizing the total capacity of the network. One
approach for balancing performance and fairness is outlined in
11 Delay Characteristics
The TCP sender 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].
Evaluations of TCP's retransmission timer can be found in [AP99] and
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. However 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 affect 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 retransmissions are not needed and 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 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.
12 Bandwidth Asymmetries
Some subnetworks may provide asymmetric bandwidth (or may cause TCP
packet flows to experience asymmetry in the capacity) 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 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 be aware that subnetworks with
significant asymmetry can result in reduced performance, unless
issues are taken to mitigate this [RFC3449].
Several strategies have been identified for reducing the impact of
asymmetry of the network path between two TCP end hosts, e.g.
[RFC3449]. These techniques attempt to reduce the number of ACKs
transmitted over the return path (low bandwidth channel) by changes
at the end host(s), and/or by modification of subnetwork packet
forwarding. While these solutions may mitigate the performance issues
caused by asymmetric subnetworks, they do have associated cost and
may have other implications. A fuller discussion of strategies and
their implications is provided in [RFC3449].
13 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 packet or a subnetwork
frame. 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 Explicit
Congestion Notification [RFC3168] (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 link_bandwidth*link_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., [MBDL99]),
it is widely considered an improvement over tail-drop discard
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 [GM02].
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
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 include flow control mechanisms that effectively require
that the rate of traffic flows be shaped on entry to 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
[RFC3168] through utilizing explicit congestion indicators in subnet
frames. Routers at the edges of the subnet, rather than shaping,
would set the explicit congestion bit in those IP packets that are
received in subnet frames that have an ECN indication. Nodes in the
subnet would need to implement an active queue management protocol
that marks subnet frames instead of dropping them.
ECN is currently a proposed standard, but it is not yet widely
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]). Compression at the subnetwork edges is of no benefit for
any of these cases.
The subnetwork may also process data that has been encrypted by the
application (OpenPGP [RFC2440] or S/MIME [RFC2633]), just above TCP
(SSL, TLS [RFC2246]), or just above IP (IPsec ESP [RFC2406]). Ciphers
generate high entropy 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 strong 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 on dialup modem links. This is
especially beneficial for small, latency-sensitive packets in
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 of 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 [LRKOJ99].
Alternate 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 the end-to-end path 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
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. For example, within a talk
spurt, only extended losses of (depending on the mode chosen) 12 to
64 packets typically cause error propagation.
15 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.
However, 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] [RFC2581] to resend the segment. In
addition, the congestion window is reduced by half, effectively
halving TCP's sending rate. If a subnetwork reorders segments
significantly such that three duplicate ACKs are generated, the TCP
sender needlessly reduces the congestion window and performance
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 [LK00] [BA02].
[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. PPP is frequently used to carry
compressed TCP/IP packets; since it was originally designed for
point-to-point and dialup links it is assumed to provide in-order
delivery. For this reason, subnetwork implementers who provide PPP
interfaces to VPNs and other, more complex subnetworks must also
maintain in-order delivery of PPP frames.
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
Third, traffic destined for the mobile host must be routed to its
current location. This roaming 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 a 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
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 sub-
optimal 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 router 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
Subnetworks connecting more than two systems must provide their own
internal layer-2 forwarding mechanisms, either implicitly (e.g.,
broadcast) or explicitly (e.g., switched). Since routing is the
major function of the Internet layer, the question naturally arises
as to the interaction between routing at the Internet layer and
routing in the subnet, and proper division of function between the
Layer 2 subnetworks can be point-to-point, connecting two systems, or
multipoint. Multipoint subnetworks can be broadcast (e.g., shared
media or emulated) or non-broadcast. Generally, IP considers
multipoint subnetworks as broadcast, with shared-medium Ethernet as
the canonical (and historical) example, and point-to-point
subnetworks as a degenerate case. Non-broadcast subnetworks may
require additional mechanisms, e.g., above IP at the routing layer
IP is ignorant of the topology of the subnetwork layer. In
particular, reconfiguration of subnetwork paths is not tracked by the
IP layer. IP is only affected by whether it can send/receive packets
sent to the remotely connected systems via the subnetwork interface
(i.e. the reachability from one router to another). IP further
considers that subnetworks are largely static - that both their
membership and existence are stable at routing timescales (tens of
seconds); both events are considered re-provisioning, rather than
Routing functionality in a subnetwork is related to addressing in
that subnetwork. Resolution of addresses on subnetwork links is
required for forwarding IP packets across links (e.g., ARP for IPv4,
or ND for IPv6). There is unlikely to be direct interaction between
subnetwork routing and IP routing. Where broadcast is provided or
explicitly emulated, address resolution can be used directly; where
not provided, the link layer routing may interface to a protocol for
resolution, e.g., to the Next-Hop Resolution Protocol [RFC2322] to
provide context-dependent address resolution capabilities.
Subnetwork routing can either complement or compete with IP routing.
It complements IP when a subnetwork encapsulates its internal
routing, and where the effects of that routing are not noticible at
the IP layer. However, if different paths in the subnetwork have
characteristics that affect IP routing, it can affect or even inhibit
the convergence of IP routing.
Routing protocols generally consider layer 2 subnetworks, i.e., with
subnet masks and no intermediate IP hops, to have uniform routing
metrics to all members. Routing can break when a link's
characteristics do not match the routing metric, in this case, e.g.,
when some member pairs have different path characteristics. Consider
a virtual Ethernet subnetwork that includes both nearby (sub-
millisecond latency) and remote (100's of milliseconds away) systems.
Presenting that group as a single subnetwork means that some routing
protocols will assume that all pairs have the same delay, and that it
is small. Because this is not the case, the routing tables
constructed may be suboptimal or may even fail to converge.
When a subnetwork is used to transit between a set of routers, it
conventionally provide the equivalent of a full mesh of point-to-
point links. Simplicity of the internal subnet structure can be used
(e.g., via NHRP [RFC2332]) to reduce the size of address resolution
tables, but routing exchanges will continue to reflect the full mesh
they emulate. In general, subnetworks should not be used as a transit
among a set of routers where routing protocols would break if a full
mesh of equivalent point-to-point links were used.
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. Learning
bridges perform Layer-2 subnetwork forwarding, avoiding the need for
dynamic routing at each subnetwork hop. Another is the "handoff"
mechanism in cellular telephone networks, particularly the "soft
handoff" scheme in IS-95 CDMA.
Subnetworks that cover large geographic areas or include links of
widely-varying capabilities should be avoided. IP routing generally
considers all multipoint subnets equivalent to a local, shared-medium
link with uniform metrics between any pair of systems, and ignores
internal subnetwork topology. Where a subnetwork diverges from that
assumption, it is the obligation of subnetwork designers to provide
compensating mechanisms. Not doing so can affect the scalability and
convergence of IP routing, as noted above.
The subnetwork designer who decides to implement internal routing
should consider whether a custom routing algorithm is warranted, or
if an existing Internet routing algorithm or protocol may suffice.
The designer should consider whether this decision is to reduce the
address resolution table size (possible, but with additional protocol
support required), or is trying to reduce routing table complexity.
The latter may be better achieved by partitioning the subnetwork,
either physically or logically, and using network-layer protocols to
support partitioning (e.g., AS's in BGP). Protocols and routing
algorithms can be notoriously subtle, complex and difficult to
implement correctly. Much work can be avoided if an existing
protocol or existing implementations can be readily used.
18 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.
A corollary of this principle is that end-to-end security (e.g.,
confidentiality, authentication, 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] and S/MIME [RFCs-2630-2634], 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
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.
The clear consensus of the working group held that subnetwork
security mechanisms, especially when weak or incorrectly implemented
[BGW01], 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
The other point of view encourages subnetwork security on the
principle that it is better than the default situation, which all too
often is no security at all. Users of especially vulnerable subnets
(such as consumers who have wireless home networks and/or shared
media Internet access) often have control over at most one endpoint
-- usually a client -- and therefore cannot enforce the use of end-
to-end mechanisms. However, subnet security can be entirely adequate
for protecting low-valued assets against the most likely threats. In
any event, subnet mechanisms do not preclude the use of end-to-end
mechanisms, which are typically used to protect highly-valued assets.
This viewpoint recognizes that many security policies implicitly
assume that the entire end-to-end path is composed of a series of
concatenated links that are nominally physically secured. That is,
these policies assume that all endpoints of all links are trusted and
that access to the physical medium by attackers is difficult. To
meet the assumptions of such policies, explicit mechanisms are needed
for links (especially shared medium links) that lack physical
protection. This, for example, is the rationale that underlies Wired
Equivalent Privacy (WEP) in the IEEE 802.11 [IEEE80211] wireless LAN
standard, and the Baseline Privacy Interface in the DOCSIS [DOCSIS1]
[DOCSIS2] data over cable television networks standards.
We therefore recommend that subnetwork designers 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.
In keeping with the "placement principle", a clear consensus exists
for another subnetwork security role: the protection of the
subnetwork itself. 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 can not ensure) end-
One security service can be provided by the subnetwork that will aid
in the solution to an overall Internet problem: subnetwork security
SHOULD provide a mechanism to authenticate the source of a subnetwork
frame. This function is missing in some current protocols, e.g., the
use of ARP [RFC0826] to associate an IPv4 address with a MAC address.
The IPv6 Neighbor Discovery (ND) [RFC2461] performs a similar
There are well known security flaws with this address resolution
mechanism [Wilbur99]. However, the inclusion of subnetwork frame
source authentication will permit a secure subnetwork address
Another potential role for subnetwork security is to protect users
against traffic analysis, i.e., identifying the communicating parties
and determining their communication patterns and volumes even when
their actual contents are protected by strong end-to-end security
mechanisms. Lower-layer security can be more effective against
traffic analysis due to its inherent ability to aggregate the
communications of multiple parties sharing the same physical
facilities while obscuring higher layer protocol information that
indicates specific end points, such as IP addresses and TCP/UDP port
However, traffic analysis is a notoriously subtle and difficult
threat to understand and defeat, far more so than threats to
confidentiality and integrity. We therefore urge extreme care in the
design of subnetwork security mechanisms specifically intended to
thwart traffic analysis.
Subnetwork designers must keep in mind that design and implementation
for security is difficult [Schneier00]. [Schneier95] 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 mistakes that have been made:
1. Use of weak and untested algorithms [Crypto9912] [BGW01]. 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 [BGW01].
6. Incorrect use of a sound mechanism. For example [BGW01], 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.
10. Failure to address "man-in-the-middle" attacks, e.g., with mutual
11. Failure to provide a form of replay detection, e.g., to prevent a
receiver from accepting packets from an attacker that simply resends
previously captured network traffic.
12. Failure to provide integrity mechanisms when providing
confidentiality schemes [Bel98].
This list is by no means comprehensive. Design problems are
difficult to avoid, but expert review is generally invaluable in
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 of the form RFCnnnn are Internet Request for Comments
(RFC) documents available online at www.rfc-editor.org.
[ATMFTM] The ATM Forum, "Traffic Management Specification, Version
4.0", April 1996, document af-tm-0056.000 (www.atmforum.com).
[BGW01] Nikita Borisov, Ian Goldberg and David Wagner, "Intercepting
Mobile Communications: The Insecurity of 802.11," In Proceedings of
ACM MobiCom, July 2001.
[BPK98] Hari Balakrishnan, Venkata Padmanabhan, Randy H. Katz. 'The
Effects of Asymmetry on TCP Performance." ACM Mobile Networks and
Applications (MONET), 1998.
[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
[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.
[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
[RFC791] Jon Postel. "Internet Protocol". September 1981.
[RFC793] Jon Postel. "Transmission Control Protocol", September
[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.
[RFC1661] W. Simpson. "The Point-to-Point Protocol (PPP)". July 1994.
[RFC1812] F. Baker, "Requirements for IP Version 4 Routers". June
[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".
[RFC2309] B. Braden, D. Clark, J. Crowcroft, B. Davie, S. Deering,
D. Estrin, S. Floyd, V. Jacobson, G. Minshall, C. Partridge, L.
Peterson, K. Ramakrishnan, S. Shenker, J. Wroclawski, L. Zhang.
"Recommendations on Queue Management and Congestion Avoidance in the
Internet". April 1998.
[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".
[RFC2395] R. Friend, R. Monsour. "IP Payload Compression Using LZS".
[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".
[RFC2406] S. Kent, R. Atkinson. "IP Encapsulating Security Payload
(ESP)". November 1998.
[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",
[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.
[RFC2914] S. Floyd. "Congestion Control Principles". September 2000
[RFC2923] K. Lahey. "TCP Problems with Path MTU Discovery".
[RFC2988] V.Paxson, M. Allman. "Computing TCP's Retransmission
Timer". November 2000.
[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.
[RFC3096] M. Degermark, ed., "Requirements for robust IP/UDP/RTP
header compression", July 2001.
[RFC3168] K. Ramakrishnan, S. Floyd, D. Black, "The Addition of
Explicit Congestion Notification (ECN) to IP", September 2001.
[Schneier95] Schneier, Bruce, Applied Cryptography: Protocols,
Algorithms and Source Code in C (John Wiley and Sons, October 1995).
[Schneier00] Schneier, Bruce, Secrets and Lies: Digital Security in a
Networked World (John Wiley & Sons, August 2000).
[SRC81] Jerome H. Saltzer, David P. Reed and David D. Clark, "End-to-
End Arguments in System Design". Second International Conference on
Distributed Computing Systems (April, 1981) pages 509-512. Published
with minor changes in ACM Transactions in Computer Systems 2, 4,
November, 1984, pages 277-288. Reprinted in: Craig Partridge, editor
Innovations in internetworking. Artech House, Norwood, MA, 1988,
pages 195-206. ISBN 0-89006-337-0.
[SSL2] Hickman, Kipp, "The SSL Protocol", Netscape Communications
Corp., Feb 9, 1995.
[SSL3] A. Frier, P. Karlton, and P. Kocher, "The SSL 3.0 Protocol",
Netscape Communications Corp., Nov 18, 1996.
References of the form RFCnnnn are Internet Request for Comments
(RFC) documents available online at www.rfc-editor.org.
[802.1D] Information Technology Telecommunications and information
exchange between systems Local and metropolitan area networks, Common
specifications Media access control (MAC) bridges, IEEE 802.1D, 1998.
[802.1p] IEEE, 802.1p, Standard for Local and Metropolitan Area
Networks - Supplement to Media Access Control (MAC) Bridges: Traffic
Class Expediting and Multicast
[AP99] M. Allman, V. Paxson, On Estimating End-to-End Network Path
Properties, In Proceedings of ACM SIGCOMM 99.
[AR02] G. Acar and C. Rosenberg, Weighted Fair Bandwidth-on-Demand
(WFBoD) for Geo-Stationary Satellite Networks with On-Board
Processing, Computer Networks, 39(1), 2002.
[BA02] Ethan Blanton, Mark Allman. On Making TCP More Robust to
Packet Reordering. ACM Computer Communication Review, 32(1), January
[Bel98] Steven M. Bellovin, "Cryptography and the Internet", in
Proceedings of CRYPTO '98, August 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.
[CGMP] Farinacci D., Tweedly A., Speakman T., "Cisco Group Management
Protocol (CGMP)", 1996/1997
[ITU-I363] ITU-T I.363.5 B-ISDN ATM Adaptation Layer Specification
Type AAL5, International Standards Organisation (ISO), 1996.
[RFC3366] Fairhurst, G., and L. Wood, Advice to link designers on
link Automatic Repeat reQuest (ARQ), August 2002.
[RFC3449] H. Balakrishnan, V. N. Padmanabhan, G. Fairhurst, M,
Sooriyabandara. "TCP Performance Implications of Network Path
Asymmetry", December 2002.
[Crypto9912] Schneier, Bruce "European Cellular Encryption
Algorithms" Crypto-Gram (December 15, 1999)
[DIX82] 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, European Standard
(Telecommunications Series) EN 301 192 v1.2.1(1999-06)
[ES00] David A. Eckhardt and Peter Steenkiste, "Effort-limited Fair
(ELF) Scheduling for Wireless Networks, Proceedings of IEEE Infocom
[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
[IEEE80211] IEEE 802.11 Wireless LAN standard. Available from
[ISO13818] ISO/IEC, ISO/IEC 13818-1:2000(E) Information Technology
- Generic coding of moving pictures and associated audio information:
Systems, Second edition, 2000-12-01 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:
[KY02] F. Khafizov, M. Yavuz. Running TCP Over IS-2000, Proceedings
of IEEE ICC, 2002.
[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.
[LKJK02] R. Ludwig, A. Konrad, A. D. Joseph, R. H. Katz, "Optimizing
the End-to-End Performance of Reliable Flows over Wireless Links",
Kluwer/ACM Wireless Networks Journal, Vol. 8, Nos. 2/3, pp. 289-299,
[LRKOJ99] R. Ludwig, B. Rathonyi, A. Konrad, K. Oden, A. Joseph,
Multi-Layer Tracing of TCP over a Reliable Wireless Link, pp.
144-154, In Proceedings of ACM SIGMETRICS 99.
[LS00] R. Ludwig, K. Sklower, The Eifel Retransmission Timer, ACM
Computer Communication Review, Vol. 30, No. 3, July 2000.
[MAGMA-PROXY] Work In Progress, MAGMA WG, draft-ietf-magma-igmp-
[MAGMA-SNOOP] Work In Progress, MAGMA WG, draft-ietf-magma-
[MBB00] May, M., Bonald, T., and Bolot, J-C., "Analytic Evaluation of
RED Performance", INFOCOM 2000.
[MBDL99] May, M., Bolot, J., Diot, C., and Lyles, B., "Reasons not to
deploy RED", Proc. of 7th. International Workshop on Quality of
Service (IWQoS'99), June 1999.
[GM02] Luigi Alfredo Grieco1, Saverio Mascolo, "TCP Westwood and Easy
RED to Improve Fairness in High-Speed Networks", Proceedings of the
7th International Workshop on Protocols for High-Speed Networks,
[MBONED-GAP] Meyer, D. and B. Nickless, Work In Progress, MBoned WG,
[MYR95] Nanette J. Boden, Danny Cohen, Robert E. Felderman, Alan E.
Kulawik, Charles L. Seitz, et al. MYRINET: A Gigabit per Second
Local Area Network, IEEE-Micro,Vol.15, No.1, February 1995, pp.29-36.
[RF95] Romanow, A., and Floyd, S., "Dynamics of TCP Traffic over ATM
Networks". IEEE Journal of Selected Areas in Communication, V. 13 N.
4, May 1995, p. 633-641.
[RFC0826] Plummer, D.C., "Ethernet Address Resolution Protocol: Or
converting network protocol addresses to 48-bit Ethernet address for
transmission on Ethernet hardware," STD 37, RFC 826, November 1982.
[RFC1071] R. Braden, D. Borman, C. Partridge, "Computing the Internet
Checksum", September 1988.
[RFC1112] S. Deering, "Host Extensions for IP Multicasting", August
[RFC1750] D. Eastlake, S. Crocker, J. Schiller, "Randomness
Recommendations for Security", December 1994.
[RFC2018] M. Mathis, J. Mahdavi, S. Floyd, A. Romanow. "TCP
Selective Acknowledgement Options". October 1996.
[RFC2236] W. Fenner, Internet Group Management Protocol, Version 2.,
[RFC2328] J. Moy, "OSPF Version 2", April 1998.
[RFC2401] S. Kent, R. Atkinson, "Security Architecture for the
Internet Protocol". November 1998.
[RFC2440] J. Callas et al. "OpenPGP Message Format". November 1998.
[RFC2460] S. Deering, R. Hinden. "Internet Protocol, Version 6
(IPv6) Specification". December 1998.
[RFC2461] T. Narten, E. Nordmark, W. Simpson. "Neighbor Discovery
for IP Version 6 (IPv6)". December 1998.
[RFC2616] R. Fielding et al. "Hypertext Transfer Protocol --
HTTP/1.1". June 1999.
[RFC2630] R. Housley. "Cryptographic Message Syntax". June 1999.
[RFC2631] E. Rescorla. "Diffie-Hellman Key Agreement Method". June
[RFC2632] B. Ramsdell. "S/MIME Version 3 Certificate Handling".
[RFC2633] B. Ramsdell. "S/MIME Version 3 Message Specification".
[RFC2710] S. Deering, W. Fenner, B. Haberman, Multicast Listener
Discovery (MLD) for IPv6, October 1999.
[RFC2784] D. Farinacci, T. Li, S. Hanks, D. Meyer, P. Traina.
"Generic Routing Encapsulation (GRE)". March 2000.
[RFC2923] K. Lahey. "TCP Problems with Path MTU Discovery".
[RFC3048] B. Whetten, L. Vicisano, R. Kermode, M. Handley, S. Floyd,
M. Luby. "Reliable Multicast Transport Building Blocks for One-to-
Many Bulk-Data Transfer". January 2001.
[RFC3376] B. Cain, S. Deering, I. Kouvelas, B. Fenner, A.
Thyagarajan, Internet Group Management Protocol, Version 3, October
[RFC3488] Cisco Systems Router-port Group Management Protocol (RGMP).
I. Wu, T. Eckert. February 2003.
[RFC3590] B. Haberman, Source Address Selection for the Multicast
Listener Discovery (MLD) Protocol, September 2003.
[SP2000] "When the CRC and TCP Checksum Disagree", Jonathan Stone &
Craig Partridge, ACM SIGCOMM, September 2000.
[Stevens94] R. Stevens, "TCP/IP Illustrated, Volume 1," Addison-
Wesley, 1994 (section 2.10).
[TCPF98] Dong Lin and H.T. Kung, "TCP Fast Recovery Strategies:
Analysis and Improvements", IEEE Infocom, March 1998. Available
[WFBA2000] David Wagner, Jeffrey S. Foster, Eric Brewer and Alexander
Aiken, "A First Step Toward Automated Detection of Buffer Overrun
Vulnerabilities", Proceedings of NDSS2000, or
[Wilbur89] Wilbur, Steve R., Jon Crowcroft, and Yuko Murayama. "MAC
layer Security Measures in Local Area Networks, " Local Area Network
Security, Workshop LANSEC '89 Proceedings, Springer-Verlag, April
Phil Karn, Editor Qualcomm 5775 Morehouse Drive San Diego CA 92121
858 587 1121 firstname.lastname@example.org
Carsten Bormann Universitaet Bremen FB3 TZI Postfach 330440 D-28334
Bremen, GERMANY +49 421 218 7024 email@example.com
Godred (Gorry) Fairhurst Department of Engineering University of
Aberdeen Aberdeen, AB24 3UE UK firstname.lastname@example.org
Dan Grossman Motorola, Inc. 20 Cabot Blvd. Mansfield, MA 02048
Reiner Ludwig Ericsson Research Ericsson Allee 1 52134 Herzogenrath,
Germany +49 2407 575 719 Reiner.Ludwig@ericsson.com
Jamshid Mahdavi Volera, Inc. 2211 N. 1st St. San Jose, CA 95131
Gabriel Montenegro Sun Microsystems Laboratories, Europe 29, chemin
du Vieux Chene 38240 Meylan, FRANCE email@example.com
Joe Touch USC/ISI 4676 Admiralty Way Marina del Rey CA 90292 310 448
9151 firstname.lastname@example.org http://www.isi.edu/touch
Lloyd Wood Global Defense and Space Group, Cisco Systems 9 New Square
Park, Bedfont Lakes Feltham TW14 8HA, United Kingdom +44 (0)20 8824
4236 email@example.com http://www.ee.surrey.ac.uk/Personal/L.Wood/
Aaron Falk USC Information Sciences Institute 4676 Admiralty Way
Marina Del Rey, CA 90292 310-448-9327 firstname.lastname@example.org
Saverio Mascolo Dipartimento di Elettrotecnica ed Elettronica,
Politecnico di Bari Via Orabona 4, 70125 Bari, Italy +39 080 596
3621 email@example.com http://www-dee.poliba.it/dee-
Marie-Jose Montpetit firstname.lastname@example.org
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