Internet Engineering Task Force                        Phil Karn, editor
INTERNET DRAFT                                                  Qualcomm
          Performance Implications of Link Characteristics Working Group

File: draft-ietf-pilc-link-design-09.txt                  February, 2002
                                                 Expires:   August, 2002

                Advice for Internet Subnetwork Designers

Status of this Memo

   This document is an Internet-Draft and is in full conformance with
   all provisions of Section 10 of RFC2026.

   Internet-Drafts are working documents of the Internet Engineering
   Task Force (IETF), its areas, and its working groups.  Note that
   other groups may also distribute working documents as Internet-

   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
   subnetworks (collectively referred to as subnetworks) who wish to
   support the Internet protocols but who may be unfamiliar with
   Internet architecture and the implications of their design choices on
   the performance and efficiency of the Internet.

   This document represents a consensus of the members of the IETF
   Performance Implications of Link Characteristics (PILC) working

Introduction and Overview

   IP, the Internet Protocol [RFC791] is the core protocol of the
   Internet. IP defines a simple "connectionless" packet-switched
   network.  The success of the Internet is largely attributed to IP's
   simplicity, the "end-to-end principle" [SRC81] on which the Internet
   is based, and the resulting ease of carrying IP on a wide variety of
   subnetworks not necessarily designed with IP in mind.

   But while many subnetworks carry IP, they do not necessarily do so
   with maximum efficiency, minimum complexity or minimum cost. Nor do
   they implement certain features to efficiently support newer Internet
   features of increasing importance, such as multicasting and quality
   of service.

   With the explosive growth of the Internet, IP is an increasingly
   large fraction of the traffic carried by the world's
   telecommunications networks. It therefore makes sense to optimize
   both existing and new subnetwork technologies for IP as much as

   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. I.e., subnetwork designers make many more errors of
   commission than errors of omission.  But certain enhanced Internet
   features, such as multicasting and quality-of-service, rely on
   support from the underlying subnetworks beyond that necessary to
   carry "traditional" unicast, best-effort IP.

   A major consideration in the efficient design of any layered
   communication network is the appropriate layer(s) in which to
   implement a given function. This issue was first addressed in the
   seminal paper "End-to-End Arguments in System Design" [SRC81]. 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.

   These functions cannot be provided solely by the concatenation of
   hop-by-hop services, so duplicating these functions at the lower
   protocol layers (i.e., within the subnetwork) can be needlessly
   redundant or even harmful to cost and performance.

   However, partial duplication of functionality in a lower layer can
   *sometimes* be justified by performance, security or availability
   considerations. Examples include link-layer retransmission to improve
   the performance of an unusually lossy channel, e.g., mobile radio;
   link level encryption intended to thwart traffic analysis; and
   redundant transmission links to improve availability, increase
   throughput, or to guarantee performance for certain classes of
   traffic.  Duplication of protocol function should be done only with
   an understanding of system level implications, including possible
   interactions with higher-layer mechanisms.

   The architecture of the Internet was 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

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 packets and are reassembled at
   the destination. The maximum packet size supported by a subnetwork is
   known as its Maximum Transmission Unit (MTU).

   Subnetworks may, but are not required to indicate the length of each
   packet they carry.  One example is Ethernet with the widely used DIX
   [DIX] (not IEEE 802.3 [IEEE8023]) header, which lacks a length field
   to indicate the true data length when the packet is padded to the 60
   byte minimum.  This is not a problem for uncompressed IP because it
   carries its own length field.

   If optional header compression [RFC1144] [RFC2507] [RFC2508]
   [RFC3095] is used, however, it is required that the link framing
   indicate frame length as it is needed for the reconstruction of the
   original header.

   In IP version 4 (the version now in wide use), fragmentation can
   occur at either the sending host or in an intermediate router, and
   fragments can be further fragmented at subsequent routers if

   In IP version 6, fragmentation can occur only at the sending host; it
   cannot occur in a router.

   Both IPv4 and IPv6 provide a "path MTU discovery" procedure [RFC1191]
   [RFC1435] [RFC1981] that allows the sending host to avoid
   fragmentation by discovering the minimum MTU along a given path and
   reducing its packet sizes accordingly. This procedure is optional in
   IPv4 but mandatory in IPv6 where there is no router fragmentation.

   Path MTU discovery is widely deployed, but it sometimes encounters
   problems. Some routers fail to generate the ICMP messages that convey
   path MTU information to the sender, and sometimes the ICMP messages
   are blocked by overly restrictive firewalls.  The result can be a
   "Path MTU Black Hole" [RFC2923] [RFC1435].

   The Path MTU Discovery procedure, the persistence of path MTU black
   holes, and the deletion of router fragmentation in IPv6 reflects a
   consensus of the Internet technical community that IP fragmentation
   is best avoided. This requires that subnetworks support MTUs that are
   "reasonably" large. The smallest MTU permitted in IPv4 by [RFC791] is
   68 bytes, but such a small value would clearly be inefficient.
   Because IPv6 omits router fragmentation, [RFC 2460] specifies a
   larger minimum MTU of 1280 bytes. Any subnetwork with an internal
   packet payload smaller than 1280 bytes MUST implement an internal
   fragmentation/reassembly mechanism if it is to support IPv6.

   If a subnetwork cannot directly support a "reasonable" MTU with
   native framing mechanisms, it should internally fragment. That is, it
   should transparently break IP packets into internal data elements and
   reassemble them at the other end of the subnetwork.

   This leaves the question of what is a "reasonable" MTU.  Ethernet (10
   and 100 Mb/s) has a MTU of 1500 bytes, and because of its ubiquity
   few Internet paths have MTUs larger than this value.  This severely
   limits the utility of larger MTUs provided by other subnetworks. But
   larger MTUs are increasingly desirable on high speed subnetworks to
   reduce the per-packet processing overhead in host computers, and
   implementers are encouraged to provide them even though they may not
   be usable when Ethernet is also in the path.

   Various "tunneling" schemes, such as IP Security [RFC2406] treat IP
   as a subnetwork for IP.  Since tunneling adds header overhead, it can
   trigger fragmentation even when the same physical subnetworks (e.g.,
   Ethernet) are used on both sides of the IP router. Tunneling has made
   it more difficult to avoid IP fragmentation and has increased the
   incidence of path MTU black holes. Larger subnetwork MTUs may help to
   alleviate this problem.

  Choosing the MTU in Slow Networks

   In slow networks, the largest possible packet may take a considerable
   time to send.  Interactive response time should not exceed the well-
   known human factors limit of 100 to 200 ms. This includes all sources
   of delay: electromagnetic propagation delay, queuing delay, and the
   store-and-forward time, i.e,. the time to transmit a packet at link

   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 IP 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) is another example of this
   technique. However, ATM is generally used on high-speed links where
   the store-and-forward delays are already minimal, and it introduces
   significant (~9%) additional overhead due to the addition of 5-byte
   frame headers to each 48-byte ATM cell.

   A third example of data link layer fragmentation is contained in the
   Data Over Cable Service Interface Specification (DOCSIS) version 1.1
   [DOCSIS2].  DOCSIS version 1.1 introduces an internal fragmentation
   and reassembly procedure so that real-time voice packets can avoid
   undue delay when sharing the relatively slow upstream channel of a
   cable modem with large data packets.

   To summarize, there is a fundamental tradeoff between efficiency and
   latency in the design of a subnetwork, and the designer should keep
   this in mind.

Framing on Connection-Oriented Subnetworks

   IP requires that subnetworks mark the beginning and end of each
   variable-length, asynchronous IP packet.  Some examples of links and
   subnetworks that do not provide this as an intrinsic feature include:

   1. leased lines carrying a synchronous bit stream;

   2. ISDN B-channels carrying a synchronous octet stream;

   3. dialup telephone modems carrying an asynchronous octet stream;


   4. Asynchronous Transfer Mode (ATM) networks carrying an asynchronous
   stream of fixed-sized "cells".

   The Internet community has defined packet framing methods for all
   these subnetworks. The Point-To-Point Protocol (PPP) [RFC1661] is
   applicable to bit synchronous, octet synchronous and octet
   asynchronous links (i.e., examples 1-3 above). ATM has its own
   framing methods described in [RFC2684] [RFC2364].

   At high speeds, a subnetwork should provide a framed interface
   capable of carrying asynchronous, variable-length IP datagrams.  The
   maximum packet size supported by this interface is discussed above in
   the MTU/Fragmentation section.  The subnetwork may implement this
   facility in any convenient manner.

   IP packet boundaries need not coincide with any framing or
   synchronization mechanisms internal to the subnetwork. When the
   subnetwork implements variable sized data units, the most
   straightforward approach is to place exactly one IP packet into each
   subnetwork data unit (SNDU), and to rely on the subnetwork's existing
   ability to delimit SNDUs to also delimit IP packets.  A good example
   is Ethernet. But some subnetworks have SNDUs of one or more fixed
   sizes, as dictated by switching, forward error correction and/or
   interleaving considerations.  Examples of such subnetworks include
   ATM, with a single cell size of 48 bytes plus a 5-byte header, and
   IS-95 digital cellular, with two "rate sets" of four fixed frame
   sizes each that may be selected on 20 millisecond boundaries.

   Because IP packets are 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. Or only the last SNDU of the
   packet could be marked, as this would implicitly mark the next SNDU
   as the first in a new IP packet. The AAL5 (ATM Adaption Layer 5)
   scheme used with ATM is an example of this approach, though it adds
   other features, including a payload length field and a payload CRC.

   In AAL5, the ATM User-User Indication, which is encoded in the
   Payload Type field of an ATM cel, indicates the end of a packet.  The
   packet control information (trailer) is located at the end of the

   Another framing technique is to insert per-segment overhead to
   indicate the presence of a segment option.  When present, the option
   carries a pointer to the end of the packet.  This differs from AAL5
   in that it permits another packet to follow within the same segment.
   MPEG-2 [EN301] [ISO13181] supports this style of fragmentation, and
   may utilize either padding (limiting each transport stream packet to
   carry only part of one packet), or to allow a second packet to start
   (no padding).

   A third approach is to insert a special flag sequence into the data
   stream between each IP packet, and to pack the resulting data stream
   into SNDUs without regard to SNDU boundaries. The flag sequence can
   also pad unused space at the end of an SNDU. If the special flag
   appears in the user data, it is escaped to an alternate sequence
   (usually larger than a flag) to avoid being misinterpreted as a flag.
   The HDLC-based framing schemes used in PPP are all examples of this

   All three adaptation schemes introduce overhead; how much depends on
   the distribution of IP packet sizes, the size(s) of the SNDUs, and in
   the HDLC-like approaches, the content of the IP packet (since flags
   occurring in the packet must be escaped, which expands them). The
   designer must also weigh implementation complexity in the choice and
   design of an adaptation layer.

Connection-Oriented Subnetworks

   IP has no notion of a "connection"; it is a purely connectionless
   protocol.  When a connection is required by an application, it is
   usually provided by TCP, the Transmission Control Protocol, running
   atop IP on an end-to-end basis.

   Connection-oriented subnetworks can be (and are) widely used to carry
   IP, but often with considerable complexity.  Subnetworks with a few
   nodes can simply open a permanent connection between each pair of
   nodes. This is frequently done with ATM. But the number of
   connections increases as the square of the number of nodes, so this
   is clearly impractical for large subnetworks. A "shim" layer between
   IP and the subnetwork is therefore required to manage connections.
   This is one of the most common functions of a Subnetwork Dependent
   Convergence Function (SNDCF) sublayer between IP and a subnetwork.

   SNDCFs typically open subnetwork connections as needed when an IP
   packet is queued for transmission and close them after an idle
   timeout. There is no relation between subnetwork connections and any
   connections that may exist at higher layers (e.g., TCP).

   Because Internet traffic is typically bursty and transaction-
   oriented, it is often difficult to pick an optimal idle timeout. If
   the timeout is too short, subnetwork connections are opened and
   closed rapidly, possibly over-stressing its call management system
   (especially if was designed for voice traffic holding times). If the
   timeout is too long, subnetwork connections are idle much of the
   time, wasting any resources dedicated to them by the subnetwork.

   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 sharing of capacity inherent in IP.

   In any event, if an SNDCF that opens and closes subnet connections is
   used support IP, care should be taken to make sure that call
   processing in the subnet can keep up with relatively short holding

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
   gateways); a shared link has more than two, using either an
   inherently broadcast media (e.g., Ethernet, radio) or on a switching
   layer hidden from the network layer (switched Ethernet, Myrinet
   [MYR], ATM).  Switched subnetworks handle broadcast by copying
   broadcast packets to ensure each end system receives a copy of each

   Several Internet protocols for IPv4 make use of broadcast
   capabilities, including link-layer address lookup (ARP), auto-
   configuration (RARP, BOOTP, DHCP), and routing (RIP).

   The lack of broadcast 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 use 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, NTP).
   IPv6 additionally relies on multicast for network configuration
   (DHCP-like autoconfiguration) and link-layer address discovery
   (replacing ARP). In the case of IPv6 this allows autoconfiguration
   and address discovery to span across routers, whereas the IPv4
   broadcast-based services cannot (RFC1812).

   Multicast enabled IP routers organize each multicast group into a
   spanning tree, and route multicast packets by making a copy of each
   packet for each output interface that includes at least one
   downstream member of the multicast group.

   Multicasting is considerably more efficient when a subnetwork
   explicitly supports it. For example, a router relaying a multicast
   packet onto an Ethernet subnet need send only one copy, 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 node incurs software overhead for every packet, even if that
   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 media, with
   hardware receiver filters.

   Receivers 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 needs.

   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 that passes all of the multicast (and unicast)
   traffic to which the associated host subscribes, plus some small
   fraction of multicast traffic to which the host does not subscribe.
   Host/router software then has to discard the unwanted packets that
   pass the hardware filter.

   There does not need to be a one-to-one mapping between subnetwork
   multicast address and 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
   software filtering.

   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 uneccessary
   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).

Bandwidth on Demand (BoD) Subnets

   Wireless networks, including both satellite and terrestrial, may use
   Bandwidth on Demand (BoD). Bandwidth on demand, which is implemented
   at the link layer by Demand Assignment Multiple Access (DAMA) in TDMA
   systems, is currently one proposed mechanism to efficiently share
   limited spectrum resources amongst a large number of users.

   The design parameters for BoD are similar to those in connection
   oriented subnetworks, however the implementations may be 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, 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
   hold idle channels open expecting new data to arrive, BoD subnets
   request channel access based on buffer occupancy (or expected buffer
   occupancy) on the sending port. Poor performance will likely result
   if the sender does not anticipate additional traffic arriving at that
   port during the time it takes to grant a transmission request. It is
   recommended that the algorithm have the capability to extend a hold
   on the channel for data that has arrived after the original request
   was generated (this may done by piggybacking new requests on user

   There are a wide variety of BoD protocols available.  However, there
   has been relatively little comprehensive research on the interactions
   between the BoD mechanisms and Internet protocol performance.  One
   item that has been studied is TCP's retransmission timer [KY02].  BoD
   systems can cause spurious timeouts when adjusting from a relativly
   high data rate to a relativly low data rate.  In this case, TCP's
   transmitted data takes longer to get through the network than
   predicted by the retransmission timeout (RTO) and therefore the TCP
   sender is prone to resending a segment prematurely.

Reliability and Error Control

   In the Internet architecture, the ultimate responsibility for error
   recovery is at the end points. The Internet may occasionally drop,
   corrupt, duplicate or reorder packets, and the transport protocol
   (e.g., TCP) or application (e.g., if UDP is used as the transport
   protocol) must recover from these errors on an end-to-end basis.
   Error recovery in the subnetwork is therefore justified only to the
   extent that it can enhance overall performance.  It is important to
   recognize that a subnetwork can go too far in attempting to provide
   error recovery services in the Internet environment.  Subnet
   reliability should be "lightweight", i.e., it only has to be "good
   enough", *not* perfect.

   In this section we discuss how to analyze characteristics of a
   subnetwork to determine what is "good enough".  The discussion below
   focuses on TCP, which is the most widely used transport protocol in
   the Internet.  It is widely believed (and is a stated goal within the
   IETF) that non-TCP transport protocols should attempt to be "TCP-
   friendly" and have many of the same performance characteristics.
   Thus, the discussion below should be applicable even to portions of
   the Internet where TCP may not be the predominant protocol.

 TCP vs Link-Layer Retransmission

   Error recovery involves the generation and transmission of redundant
   information computed from user data. Depending on how much redundant
   information is sent and how it is generated, the receiver can use it
   to reliably detect transmission errors; correct up to some maximum
   number of transmission errors; or both. The general approach is known
   as Error Control Coding, or ECC.

   For largely historical reasons, the use of ECC to detect transmission
   errors so that retransmissions (hopefully without errors) can be
   requested is widely known as "ARQ" (Automatic Repeat Request). ARQ
   has been used for many decades in computer networking protocols.

   When enough ECC information is available to permit the receiver to
   correct transmission errors without a retransmission, the approach is
   known as Forward Error Correction (FEC). Due to the greater
   complexity of the required ECC and the need to tailor its design to
   the characteristics of a specific modem and channel, FEC has
   traditionally been implemented in special purpose hardware integral
   to a modem. This effectively makes it part of the physical layer.

   Unlike ARQ, FEC was seldom used for telecommunications outside of
   deep space links until the 1990s.  It is now nearly universal in
   telephone, cable and DSL modems, digital satellite links and digital
   mobile telephones. FEC is also heavily used in optical and magnetic
   storage where "retransmissions" are not possible.

   Some systems use hybrid combinations of ARQ layered atop FEC; V.90
   dialup modems (in the upstream direction) with V.42bis error control
   are one example. Most errors are corrected by the trellis (FEC) code
   within the V.90 modem, and most that remain are detected and
   corrected by the ARQ mechanisms in V.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. But in this
   discussion we will assume that if FEC is present, it is implemented
   within the physical layer.

   Depending on the layer where it is implemented, error control can
   operate on an end-to-end basis or over a shorter span such as a
   single link.  TCP is the most important example of an end-to-end
   protocol that uses an ARQ strategy.

   Many link-layer protocols use ARQ, usually some flavor of HDLC
   [ISO3309]. Examples include the X.25 link layer, the AX.25 protocol
   used in amateur packet radio, 802.11 wireless LANs, and the reliable
   link layer specified in IEEE 802.2.

   As explained in the introduction, only end-to-end error recovery can
   ensure a reliable service to the application. But some subnetworks
   (e.g., many wireless links) also require link-layer error recovery as
   a performance enhancement.  For example, many cellular links have
   small physical frame sizes (< 100 bytes) and relatively high frame
   loss rates. Relying entirely on end-to-end error recovery clearly
   yields a performance degradation, as retransmissions across the end-
   to-end path take much longer to be received than when link-local
   retransmissions are used. Thus, link-layer error recovery can often
   increase end-to-end performance. As a result, link-layer and end-to-
   end recovery often co-exist; this can lead to the possibility of
   inefficient interactions between the two layers of ARQ protocols.

   This inter-layer "competition" might lead to the following wasteful
   situation. When the link layer retransmits a packet, the link latency
   momentarily increases. Since TCP bases its retransmission timeout on
   prior measurements of end-to-end latency, including that of the link
   in question, this sudden increase in latency may trigger an
   unnecessary retransmission by TCP of a packet that the link layer is
   still retransmitting.  Such spurious end-to-end retransmissions
   generate unnecessary load and reduce end-to-end throughput. One may
   even have multiple copies of the same packet in the same link queue
   at the same time. In general, one could say the competing error
   recovery is caused by an inner control loop (link-layer error
   recovery) reacting to the same signal as an outer control loop (end-
   to-end error recovery) without any coordination between the loops.
   Note that this is solely an efficiency issue; TCP continues to
   provide reliable end-to-end delivery over such links.

   This raises the question of how persistent a link-layer sender should
   be in performing retransmission. We define the link-layer (LL) ARQ
   persistency as the maximum time that a particular link will spend
   trying to transfer a packet before it can be discarded. This
   deliberately simplified definition says nothing about maximum number
   of retransmissions, retransmission strategies, queue sizes, queuing
   disciplines, transmission delays, or the like. The reason we use the
   term LL ARQ persistency instead of a term such as 'maximum link-layer
   packet holding time' is that the definition closely relates to link-
   layer error recovery. For example, on links that implement
   straightforward error recovery strategies, LL ARQ persistency will
   often correspond to a maximum number of retransmissions permitted per
   link-layer frame [ARQ-DRAFT].

   For link layers that do not or cannot differentiate between flows
   (e.g., due to network layer encryption), the LL ARQ persistency
   should be small.  This avoids any harmful effects or performance
   degradation resulting from indiscriminate high persistence.  A
   detailed discussion of these issues is provided in [ARQ-DRAFT].

   However, when a link layer can identify individual flows and apply
   ARQ selectively [ARQ-DRAFT], then the link ARQ persistency should be
   high for a flow using reliable unicast transport protocols (e.g.,
   TCP) and must be low for all other flows.  Setting the link ARQ
   persistency larger than the largest link outage allows TCP to rapidly
   restore transmission without the need to wait for a retransmission
   time out. This generally improves TCP performance in the face of
   transient outages.  However, excessively high persistence may be
   disadvantageous; a practical upper limit of 30-60 seconds may be
   desirable. Implementation of such schemes remains a research issue.
   (See also Section "Recovery from Subnetwork Outages").

   Many subnetwork designers have opportunities to reduce the
   probability of packet loss, e.g., with FEC, ARQ and interleaving, at
   the cost of increased delay. TCP performance improves with decreasing
   loss but worsens with increasing end-to-end delay, so it is important
   to find the proper balance through analysis and simulation.

 Recovery from Subnetwork Outages

   Some types of subnetworks, particularly mobile radio, are subject to
   frequent temporary outages. For example, an active cellular data user
   may drive or walk into an area (such as a tunnel) that is out of
   range of any base station. No packets will be successfully delivered
   until the user returns to an area with coverage.

   The Internet protocols currently provide no standard way for a
   subnetwork to explicitly notify an upper layer protocol (e.g., TCP)
   that it is experiencing an outage rather than severe congestion.

   Under these circumstances TCP will, after each unsuccessful
   retransmission, wait even longer before trying again; this is its
   "exponential back-off" algorithm. And TCP will not discover that the
   subnetwork outage has ended until its next retransmission attempt. If
   TCP has backed off, this may take some time.  This can lead to
   extremely poor TCP performance over such subnetworks.

   It is therefore highly desirable that a subnetwork subject to outages
   not silently discard packets during an outage. Ideally, it should
   define an interface to the next higher layer (i.e., IP) that allows
   it to refuse packets during an outage, and to automatically ask IP
   for new packets when it is again able to deliver them. If it cannot
   do this, then the subnetwork should hold onto at least some of the
   packets it accepts during an outage and attempt to deliver them when
   the subnetwork comes back up. 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 -- a task that as discussed earlier can only be done
   properly on an end-to-end basis.

   Only a single packet per TCP connection, including ACKs, need be held
   in this way to cause the TCP sender to recover from the additional
   losses once the flow resumes [ARQ-DRAFT].

   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.

 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 to
   detect corrupted packets.  The IPv4 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

   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
   link CRC.

   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. This is significantly more robust
   against most patterns of errors than the 16-bit Internet checksum.
   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 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 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 in the subnetwork, not just its
   transmission links.  Examples of this approach include the way in
   which the Ethernet CRC-32 is handled by LAN bridges.  ATM AAL5 also
   uses an edge-to-edge CRC-32.

   Some specific applications may be tolerant of residual errors in the
   data they exchange, however 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. An application also requires 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 such a service.  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 [SRC81].

 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 practiced.

   TCP uses sequence numbering and acknowledgments (ACKs) on an end-to-
   end basis to provide reliable, sequenced, once-only delivery.  TCP
   ACKs are cumulative, i.e., each implicitly ACKs every segment
   received so far.  If an ACK is not received, the acknowledgment value
   carried in the cumulative packet will cease to advance.

   Since the most common cause of packet loss is congestion, TCP treats
   packet loss as an Internet congestion indicator. This happens
   automatically, and the subnetwork need not know anything about IP or
   TCP. It simply drops packets whenever it must, though some packet-
   dropping strategies (e.g., RED) are more fair to competing flows than

   TCP recovers from packet losses in two different ways. The most
   important is the retransmission timeout. If an ACK fails to arrive
   after a certain period of time, TCP retransmits the oldest unacked
   packet. Taking this as a hint that the network is congested, TCP
   waits for the retransmission to be ACKed before it continues, and it
   gradually increases the number of packets in flight as long as a
   timeout does not occur again.

   A retransmission timeout can impose a significant performance
   penalty, as the sender is idle during the timeout interval and
   restarts with a congestion window of 1 following the timeout. To
   allow faster recovery from the occasional lost packet in a bulk
   transfer, an alternate scheme known as "fast recovery" was introduced
   [RFC2581] [RFC2582] [RFC2914] [TCPF98].

   Fast recovery relies on the fact that when a single packet is lost in
   a bulk transfer, the receiver continues to return ACKs to subsequent
   data packets that do not actually acknowledge any 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-
   of-sequence delivery.  Recent proposals have been made to lower the
   dupack threshold to 2.

   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 this is currently an optional behavior (and, longer
   term, there will always be the possibility of congestion in portions
   of the network which don't support ECN), the lack of an ECN bit MUST
   NEVER be interpreted as a lack of congestion.  Thus, for the
   foreseeable future, TCP MUST interpret a lost packet as a signal of

   The TCP "congestion avoidance" [RFC3168] algorithm maintains a
   congestion window (cwnd) controlling the amount of data TCP may have
   in flight at any moment.  Reducing cwnd reduces the overall bandwidth
   obtained by the connection; similarly, raising cwnd increases the
   performance, up to the limit of the available bandwidth.

   TCP probes for available network bandwidth by setting cwnd at one
   packet and then increasing it by one packet for each ACK returned
   from the receiver. This is TCP's "slow start" mechanism.  When a
   packet loss is detected (or congestion is signaled by other
   mechanisms), cwnd is reset to one and the slow start process is
   repeated until cwnd reaches one half of its previous setting before
   the reset. Cwnd continues to increase past this point, but at a much
   slower rate than before. If no further losses occur, cwnd will
   ultimately reach the window size advertised by the receiver.

   This is an "Additive Increase, Multiplicative Decrease" (AIMD)
   algorithm.  The steep decrease of cwnd in response to congestion
   provides for network stability; the AIMD algorithm also provides for
   fairness between long running TCP connections sharing the same path.

 TCP Performance Characteristics


   Here we present the current "state-of-the-art" understanding of TCP
   performance.  This analysis attempts to characterize the performance
   of TCP connections over links of varying characteristics.

   Link designers may wish to use the techniques in this section to
   predict what performance TCP/IP may achieve over a new link-layer
   design.  Such analysis is encouraged.  Because this is relatively new
   analysis, and the theory is based on single stream TCP connections
   under "ideal" conditions, it should be recognized that the results of
   such analysis may be different than actual performance in the
   Internet.  That being said, we have done the best we can to provide
   information which will help designers get an accurate picture of the
   capabilities and limitations of TCP under various conditions.

  The Formulae

   The performance of TCP's AIMD Congestion Avoidance algorithm has been
   extensively analyzed.  The current best formula for the performance
   of the specific algorithms used by Reno TCP (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 throughput achievable
           MSS  is the segment size being used by the connection
           RTT  is the end-to-end round trip time of the TCP connection
           RTO  is the packet timeout (based on RTT)
           p    is the packet loss rate for the path
                (i.e. .01 if there is 1% packet loss)

   Note that the speed of the links making up the Internet path does not
   explicitly appear in this formula. Attempting to send faster than the
   slowest link in the path causes the queue to grow at the transmitter
   driving the bottleneck. This increases the RTT, which in turn reduces
   the achievable throughput.

   This is currently considered to be the best approximate formula for
   Reno TCP performance.  A further simplification to this formula is
   generally made by assuming that RTO is approximately 5*RTT.

   TCP is constantly being improved.  A simpler formula, which gives an
   upper bound on the performance of any AIMD algorithm which is likely
   to be implemented in TCP in the future, was derived by Ott, et al

                     MSS   1
           BW = C    --- -------
                     RTT sqrt(p)

   where C is 0.93.


   Both formulae assume that the TCP Receiver Window is not limiting the
   performance of the connection.  Because receiver window is entirely
   determined by end-hosts, we assume that hosts will maximize the
   announced receiver window to maximize their network performance.

   Both of these formulae allow BW to become infinite if there is no
   loss.  This is because an Internet path will drop packets at
   bottleneck queues if the load is too high.  Thus, a completely
   lossless TCP/IP network can never occur (unless the network is being

   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
   which are extremely short (<10 packets) and don't lose any packets,
   performance is driven by the TCP slow start algorithm.  For
   connections of medium length, where on average only a few segments
   are lost, single connection performance will actually be slightly
   better than given by the formulae above.

   The difference between the simple and complex formulae above is that
   the complex formula includes the effects of TCP retransmission
   timeouts.  For very low levels of packet loss (significantly less
   than 1%), timeouts are unlikely to occur, and the formulae lead to
   very similar results.  At higher packet losses (1% and above), the
   complex formula gives a more accurate estimate of performance (which
   will always be significantly lower than the result from the simple

   Note that these formulae break down as p approaches 100%.

  Analysis of Link-Layer Effects on TCP Performance

   Consider the following example:

   A designer invents a new wireless link layer which, on average, loses
   1% of IP packets.  The link layer supports packets of up to 1040
   bytes, and has a one-way delay of 20 msec.

   If this link layer were used in the Internet, on a path that
   otherwise had a round trip of of 80 msec, you could compute an upper
   bound on the performance as follows:

   For MSS, use 1000 bytes to exclude the 40 bytes of TCP/IP headers.

   For RTT, use 120 msec (80 msec for the Internet part, plus 20 msec
   each way for the new wireless link).

   For p, use .01.  For C, assume 1.

   The simple formula gives:

   BW = (1000 * 8 bits) / (.120 sec * sqrt(.01)) = 666 kbit/sec

   The more complex formula gives:

   BW = 402.9 kbit/sec

   If this were a 2 Mb/s wireless LAN, the designers might be somewhat

   Some observations on performance:

   1.  We have assumed that the packet losses on the link layer are
   interpreted as congestion by TCP.  This is a "fact of life" that must
   be accepted.

   2.  The equations for TCP performance are all expressed in terms of
   packet loss, but many link-layer designers think in terms of bit-
   error rate.  *If* channel bit errors are independent, then the
   probability of a packet being corrupted is:

   p = 1 - ([1 - BER]^[PACKET_SIZE*8])

   Here we assume PACKET_SIZE is in bytes. It includes the user data and
   all headers (TCP,IP and subnetwork).  If the inequality

   BER * [PACKET_SIZE*8] << 1

   holds, the packet loss probability p can be approximated by:

   p = BER * [PACKET_SIZE*8]

   These equations can be used to apply BER to the performance equations

   Note that PACKET_SIZE can vary from one packet to the next.  Small
   packets (such as TCP acks) generally have a smaller probability of
   packet error than, say, a TCP packet carrying one MSS (maximum
   segment size) of user data.  A flow of small TCP acks can be expected
   to be slightly more reliable than a stream of larger TCP data

   It bears repeating that the above analysis assumes that bit errors
   are statistically independent. Because this is not true for many real
   links, our computation of p is actually an upper bound, not the exact
   probability of packet loss.

   There are many reasons why bit errors are not independent on real
   links.  Many radio links are affected by propagation fading or by
   interference that lasts over many bit times.

   Also, links with Forward Error Correction (FEC) generally have very
   non-uniform bit error distributions that depend on the type of FEC,
   but in general the uncorrected errors tend to occur in bursts even
   when channel symbol errors are independent.  In all such cases our
   computation of p from BER can only place an upper limit on the packet
   loss rate.

   If the distribution of 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 will need to be used to
   determine the actual packet loss rate.

   3.  Note that the packet size plays an important role.  If the
   subnetwork loss characteristics are such that large packets have the
   same probability of loss as smaller packets, then larger packets will
   yield improved performance.

   4.  We have chosen a specific RTT that might occur on a wide-area
   Internet path within the USA.  It is important to recognize that RTTs
   vary considerably in the Internet.

   For example, RTTs are typically less than 10 msec in a wired LAN
   environment.  International connections may have RTTs of 200 msec or
   more.  Modems and other low-capacity links can add considerable delay
   due to their long packet transmission times.

   Links over geostationary repeater satellites have one-way speed-of-
   light delays of around 250ms: 125ms up to the satellite and 125ms
   down. The RTT of an end-to-end TCP connection that includes such a
   link can be expected to be greater than 250ms.

   Queues on heavily-congested links may back up, increasing RTTs.
   Finally, virtual private networks (VPNs) and other forms of
   encryption and tunneling can add significant end-to-end delay to
   network connections.

Quality-of-Service (QoS) considerations

   It is generally recognized that specific service guarantees are
   needed to support real-time multimedia, toll quality telephony and
   other performance critical applications. The provision of such
   Quality of Service guarantees in the Internet is an active area of
   research and standardization. The IETF has not converged on a single
   service model, set of services or 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 the present understanding of the
   challenges in architecting QoS for the Internet.

   There are presently two architectural approaches to providing
   mechanisms for QoS support in the Internet.

   IP Integrated Services (Intserv) [RFC1633] provides fine-grained
   service guarantees to individual flows.  Flows are identified by a
   flow specification (flowspec), which creates a stateful association
   between individual packets by matching fields in the packet header.
   Bandwidth is reserved for the flow, and appropriate traffic
   conditioning and scheduling is installed in routers along the path.
   The ReSerVation Protocol (RSVP) [RFC2205, RFC2210] usually, but not
   necessarily, is used to install the flow QoS state.  Intserv defines
   two services, in addition to the Default (best effort) service.

     -- Guaranteed Service (GS) [RFC 2212] offers hard upper bounds on
     delay to flows that conform to a traffic specification (TSpec).  It
     uses a fluid flow model to relate the TSpec and reserved bandwidth
     (RSpec) to variable delay.  Non-conforming packets are forwarded on
     a best-effort basis.

     -- Controlled Load Service (CLS) [RFC2211] offers delay and packet
     loss equivalent to that of an unloaded network to flows that
     conform to a TSpec, but no hard bounds. Non-conforming packets are
     forwarded on a best-effort basis.

   Intserv requires installation of state information in every
   participating router. 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 TOS byte of
   the IP header [RFC2474].   The PHB denotes the forwarding behavior to
   be applied to the packet in each node in the Diffserv domain.
   Although there is a "recommended" DSCP associated with each PHB, the
   mappings from DSCPs to PHBs are defined by the DS-domain.  In fact,
   there can be several DSCPs associated with the same PHB.  Diffserv
   presently defines three PHBs.

   The class selector PHB [RFC2474] replaces the IP precedence field of
   the former TOS byte. It offers relative forwarding priorities.

   The Expedited Forwarding (EF) PHB [RFC2598] guarantees that packets
   will have a well-defined minimum departure rate which, if not
   exceeded, ensures that the associated queues are short or empty.  EF
   is intended to support services that offer tightly-bounded loss,
   delay and delay jitter.

   The Assured Forwarding (AF) PHB group [RFC2597] offers different
   levels of forwarding 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.  Or, to put it another way, forwarding behavior in the
     subnet needs to be 'clued' by the forwarding behavior (service or
     PHB) at the IP layer, and

     -- the operation of a set of cooperating mechanisms, such as
     bandwidth reservation and admission control, policy management,
     traffic classification, traffic conditioning (marking, policing
     and/or shaping), selective discard, queuing and scheduling.  Note
     that support for QoS in subnets may require similar mechanisms,
     especially when these subnets are general topology subnets (e.g.,
     ATM, frame relay or MPLS) or shared media subnets.

   Many subnetwork designers face inherent tradeoffs between delay,
   throughput, reliability and cost. Other subnetworks have parameters
   that manage bandwidth, internal connection state, and the like.
   Therefore, the following subnetwork capabilities may be desirable,
   although some might be trivial or moot if the subnet is a simple
   point-to-point link.

     - The subnetwork should have the ability to reserve bandwidth for a
     connection or flow and schedule packets accordingly.

     - Bandwidth reservations should be based on a one- or two- token
     bucket model, depending on whether the service is intended to
     support constant-rate or bursty traffic.

     - If a connection or flow does not use its reserved bandwidth at a
     given time, the unused bandwidth should be available for other

     - 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 subnetwork.

     - The subnet layer should know, and be able to inform the IP layer,
     how much fixed delay and delay jitter it offers for a flow or
     connection.  If the Intserv model is used, the delay jitter
     component may best be expressed in terms of the TSpec/RSpec model
     described in [RFC2212].

     - Support of the Diffserv class selectors [RFC2474] suggests that
     the subnet might consider mechanisms that support priorities.

Fairness vs Performance

   Subnetwork designers should be aware of the tradeoffs between
   fairness and efficiency inherent in many transmission scheduling
   algorithms. For example, many local area networks use contention
   protocols to resolve access to a shared transmission channel.  These
   protocols represent overhead. Limiting the amount of data that a
   station may transmit per contention cycle helps assure each station
   of timely access to the channel, but it also increases contention
   overhead per unit of data sent.

   In some mobile radio networks, capacity is limited by interference,
   which in turn depends on average transmitter power. Some receivers
   may require considerably more transmitter power (generating more
   interference and consuming more channel capacity) than others.

   In each case, the scheduling algorithm designer must balance
   competing objectives: providing a fair share of capacity to each
   station while maximizing the total capacity of the network.

Delay Characteristics

   TCP bases its retransmission timeout (RTO) on measurements of the
   round trip delay experienced by previous packets. This allows TCP to
   adapt automatically to the very wide range of delays found on the
   Internet. The recommended algorithms are described in [RFC2988].

   These algorithms model the delay along an Internet path as a
   normally-distributed random variable with slowly-varying mean and
   standard deviation. TCP estimates these two parameters by
   exponentially smoothing individual delay measurements, and it sets
   the RTO to the estimated mean delay plus some fixed number of
   standard deviations. (The algorithm actually uses mean deviation as
   an approximation to standard deviation, as it is easier to compute.)

   The goal is to compute a RTO that is small enough to detect and
   recover from packet losses while minimizing unnecessary ("spurious")
   retransmissions when packets are unexpectedly delayed but not lost.
   Although these goals conflict, the algorithm works well when the
   delay variance along the Internet path is low, or the packet loss
   rate is low.

   If the path delay variance is high, TCP sets a RTO that is much
   larger than the mean of the measured delays. But if the packet loss
   rate is low, the large RTO is of little consequence, as timeouts
   occur only rarely.  Conversely, if the path delay variance is low,
   then TCP recovers quickly from lost packets; again, the algorithm
   works well.

   But when delay variance and the packet loss rate are both high, these
   algorithms perform poorly, especially when the mean delay is also

   Because TCP uses returning acknowledgments as a "clock" to time the
   transmission of additional data, excessively high delays (even if the
   delay variance is low) also affects TCP's ability to fully utilize a
   high speed transmission pipe. It also slows down the recovery of lost
   packets even when delay variance is small.

   Subnetwork designers should therefore minimize all three parameters
   (delay, delay variance and packet loss) as much as possible.

   In many subnetworks, these parameters are inherently in conflict.
   For example, on a mobile radio channel the subnetwork designer can
   use retransmission (ARQ) and/or forward error correction (FEC) to
   trade off delay, delay variance and packet loss in an effort to
   improve TCP performance. For example, while ARQ increases delay
   variance, FEC does not. However, FEC (especially when combined with
   interleaving) often increases mean delay even on good channels where
   ARQ would not increase either the delay or the delay variance.

   The tradeoffs among these error control mechanisms and their
   interactions with TCP can be quite complex, and they are the subject
   of much ongoing research. We therefore recommend that subnetwork
   designers provide as much flexibility as possible in the
   implementation of these mechanisms, and to provide access to them as
   discussed above in the section on Quality of Service.

Bandwidth Asymmetries

   Some subnetworks may provide asymmetric bandwidth (or may cause TCP
   packet flows to be 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 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 [asym-id].

   Several strategies have been identified for reducing the impact of
   asymmetry of the network path between two TCP end hosts, e.g. [asym-
   id].  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 [asym-id].

Buffering, flow & congestion control

   Many subnets include multiple links with varying traffic demands and
   possibly different transmission speeds. At each link there must be a
   queuing system, including buffering, scheduling and a capability to
   discard excess subnet packets.  These queues may also be part of a
   subnet flow control or congestion control scheme.

   For the purpose of this discussion, we talk about packets without
   regard to whether they refer to a complete IP datagram or a
   subnetwork packet.  At each queue, a packet experiences a delay that
   depends on competing traffic and the scheduling discipline, and is
   subjected to a local discarding policy.

   Some subnets may have flow or congestion control mechanisms in
   addition to packet dropping.  Such mechanisms can operate on
   components in the subnet layer, such as schedulers, shapers or
   discarders, and can  affect the operation of IP forwarders at the
   edges of the subnet.  However, with the exception of  RFC3168
   explicit congestion notification (discussed below), IP has no way to
   pass explicit congestion or flow control signals to TCP.

   TCP traffic, especially aggregated TCP traffic, is bursty.  As a
   result, instantaneous queue depths can vary dramatically, even in
   nominally stable networks.  For optimal performance, packets should
   be dropped in a controlled fashion, not just when buffer space is
   unavailable.  How much buffer space should be supplied is still a
   matter of debate, but as a rule of thumb, each node should have
   enough buffering to hold one 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., see
   [MBDL99]), it is widely considered an improvement over tail-drop
   discard policies.

   One form of active queue management is the Random Early Detection
   (RED) algorithm [RED93], actually a family of related algorithms. In
   one version of RED, an exponentially-weighted moving average of the
   queue depth is maintained:

     When this average queue depth is between a maximum threshold
     max_th, and a minimum threshold min_th, packets are dropped with a
     probability which is proportional to the amount by which the
     average queue depth exceeds min_th.

     When this average queue depth is equal to max_th, the drop
     probability is equal to a configurable parameter max_p.

     When this average queue depth is greater than max_th, packets are
     always dropped.  Numerous variants on RED appear in the literature,
     and there are other active queue management algorithms which claim
     various advantages over RED [MG01].

   With an active queue management algorithm, dropped packets become a
   feedback signal to trigger more appropriate congestion behavior by
   the TCPs in the end hosts.  Randomization of dropping tends to break
   up the observed tendency of TCP windows belonging to different TCP
   connections to become synchronized by correlated drops, and it also
   imposes a degree of fairness on those connections that properly
   implement TCP congestion avoidance.  Another important property of
   active queue management algorithms is that they attempt to keep
   average queue depths short while accommodating large short term

   Since TCP neither knows nor cares whether congestive packet loss
   occurs at the IP layer or in a subnet, it may be advisable for
   subnets that perform queuing and discarding to consider implementing
   some form of active queue management.  This is especially true if
   large aggregates of TCP connections are likely to share the same
   queue.  However, active queue management may be less effective in the
   case of many queues carrying smaller aggregates of TCP connections,
   e.g., in an ATM switch that implements per-VC queuing.

   Note that the performance of active queue management algorithms is
   highly sensitive to settings of configurable parameters, and also to
   factors such as RTT [MBB00] [FB00].

   Some subnets, most notably ATM, perform segmentation and reassembly
   at the subnetwork edges.  Care should be taken here in designing
   discard policies.  If the subnet discards a fragment of an IP packet,
   then the remaining fragments become an unproductive load on the
   subnet that can markedly degrade end-to-end performance [RF95].
   Subnetworks should therefore attempt to discard these extra fragments
   whenever one of them must be discarded.  If the IP packet has already
   been partially forwarded when discarding becomes necessary, then
   every remaining fragment except the one marking the end of the IP
   packet should also be discarded.  For ATM subnets, this specifically
   means using Early Packet Discard and Partial Packet Discard [ATMFTM].

   Some subnets might include flow control mechanisms that effectively
   require that the rate of traffic flows be shaped as they enter the
   subnet.  One example of such a subnet mechanism is in the ATM
   Available Bit rate (ABR) service category [ATMFTM].  Such flow
   control mechanisms have the effect of making the subnet nearly
   lossless by pushing congestion into the IP routers at the edges of
   the subnet.  In such a case, adequate buffering and discard policies
   are needed in these routers to deal with a subnet that appears to
   have varying bandwidth.  Whether there is benefit in this kind of
   flow control is controversial; there are numerous simulation and
   analytical studies that go both ways.  It appears that some of the
   issues that lead to such different results include sensitivity to ABR
   parameters, use of binary rather than explicit rate feedback, use (or
   not) of per-VC queuing, and the specific ATM switch algorithms
   selected for the study.  Anecdotally, some large networks have used
   IP over ABR to carry TCP traffic, have claimed it to be successful,
   but have published no results.

   Another possible approach to flow control in the subnet would be to
   work with TCP Explicit Congestion Notification (ECN) semantics
   [RFC3168].  Routers at the edges of the subnet, rather than shaping,
   would set the ECN bit in those IP packets that are received in subnet
   packets that have an ECN indication.  Nodes in the subnet would need
   to implement an active queue management protocol that marks subnet
   packets instead of dropping them.

   ECN is currently a proposed standard, 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]).  In any of these cases, compression in the subnetwork is
   of no benefit.

   The subnetwork may also process data that has been encrypted by the
   application (OpenPGP [RFC2440] or S/MIME [RFCs-2630-2634]), just
   above TCP (SSL, TLS [RFC2246]), or just above IP (IPSEC ESP
   [RFC2406]). Ciphers generate random-looking bit streams lacking any
   patterns that can be exploited by a compression algorithm.

   However, much data is still transmitted uncompressed over the
   Internet, so subnetwork compression may be beneficial.  Any
   subnetwork compression algorithm must not expand uncompressible data,
   e.g., data that has already been compressed or encrypted.

   We make a stronger recommendation that subnetworks operating at low
   speed or with small MTUs compress IP and transport-level headers (TCP
   and UDP) using several header compression schemes developed within
   the IETF. An uncompressed 40-byte TCP/IP header takes about 33
   milliseconds to send at 9600 bps.  "VJ" TCP/IP header compression
   [RFC1144] compresses most headers to 3-5 bytes, reducing transmission
   time to several milliseconds. This is especially beneficial for
   small, latency-sensitive packets in interactive sessions.

   Similarly, RTP compression schemes such as CRTP [RFC2508] and ROHC
   [RFC3095] compress most IP/UDP/RTP headers to one to four bytes.  The
   resulting savings are especially significant when audio packets are
   kept small to minimize store-and-forward latency.

   Designers should consider the effect of the subnetwork error rate on
   the performance of header compression. TCP ordinarily recovers from
   lost packets by retransmitting only those packets that were actually
   lost; packets arriving correctly after a packet loss are kept on a
   resequencing queue and do not need to be retransmitted.  In VJ TCP/IP
   [RFC1144] header compression, however, the receiver cannot explicitly
   notify a sender about data corruption and subsequent loss of
   synchronization between compressor and decompressor. It relies
   instead on TCP retransmission to re-synchronize the decompressor.
   After a packet is lost, the decompressor must discard every
   subsequent packet, even if the subnetwork makes no further errors,
   until the sending TCP retransmits to re-synchronize the decompressor.
   This effect can substantially magnify the effect of subnetwork packet
   losses if the sending TCP window is large, as it will often be on a
   path with a large bandwidth*delay product.

   Alternative header compression schemes such as those described in
   [RFC2507] include an explicit request for retransmission of an
   uncompressed packet to allow decompressor resynchronization without
   waiting for a TCP retransmission.  However, these schemes are not yet
   in widespread use.

   Both TCP header compression schemes do not compress widely-used TCP
   options such as selective acknowledgements (SACK).  Both fail to
   compress TCP traffic that makes use of explicit congestion
   notification (ECN).  Work is under way in the IETF ROHC WG to address
   these shortcomings in a ROHC header compression scheme for TCP.
   [RFC3095] [RFC3096]

   The subnetwork error rate also is important for RTP header
   compression.  CRTP uses delta encoding, so a packet loss on the link
   causes uncertainty about the subsequent packets, which often must be
   discarded until the decompressor has notified the compressor and the
   compressor has sent re-synchronizing information.  This typically
   takes slightly more than a round-trip time.  For links that combine
   significant error rates with latencies that require multiple packets
   to be in flight at a time, this leads to significant error
   propagation, i.e. subsequent losses caused by an initial loss.

   For links that are both high-latency (multiple packets in flight from
   a typical RTP stream) and error-prone, RTP ROHC provides a more
   robust way of RTP header compression, at a cost of higher complexity
   at the compressor and decompressor.  Within a talk spurt, only
   extended losses of (depending on the mode chosen) 12 to 64 packets
   typically cause error propagation.

Packet Reordering

   The Internet architecture does not guarantee that packets will arrive
   in the same order in which they were originally transmitted, and
   transport protocols like TCP must take this into account.

   But reordering does come at a cost with TCP as it is currently
   defined. Because TCP returns a cumulative acknowledgment (ACK)
   indicating the last in-order segment that has arrived, out-of-order
   segments cause a TCP receiver to transmit a duplicate acknowledgment.
   When the TCP sender notices three duplicate acknowledgments it
   assumes that a segment was dropped by the network and uses the fast
   retransmit algorithm [Jac90] [APS99] to resend the segment.  In
   addition, the congestion window is reduced by half, effectively
   halving TCP's sending rate.  If a subnetwork badly re-orders segments
   such that three duplicate ACKs are generated, the TCP sender
   needlessly reduces the congestion window and performance suffers.

   Packet reordering does frequently occur in parts of the Internet, and
   it seems to be difficult or impossible to eliminate [BPS99].  For
   this reason, research has begun into improving TCP's behavior in the
   face of packet reordering.

   [BPS99] cites reasons why it may even be undesirable to eliminate
   reordering. There are situations where average packet latency can be
   reduced, link efficiency can be increased, and/or reliability can be
   improved if reordering is permitted.  Examples include certain high
   speed switches within the Internet backbone and the parallel links
   used over many Internet paths for load splitting and redundancy.

   This suggests that subnetwork implementers should try to avoid packet
   reordering whenever possible, but not if doing so compromises
   efficiency, impairs reliability or increases average packet delay.

   Note that every header compression scheme currently standardized for
   the Internet requires in-order packet delivery on the link between
   compressor and decompressor. 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
   (IPCP) [RFC1332].

   Third, traffic destined for the mobile host must be routed to its
   current location. This function is the most common meaning of the
   term "Internet mobility".

   Internet mobility can be provided at any of several layers in the
   Internet protocol stack, and there is ongoing debate as to which are
   the most appropriate and efficient. Mobility is already an feature of
   certain application layer protocols; the Post Office Protocol (POP)
   [RFC1939] and the Internet Message Access Protocol (IMAP) [RFC2060]
   were created specifically to provide mobility in the receipt of
   electronic mail.

   Mobility can also be provided at the IP layer [RFC2002]. This
   mechanism provides greater transparency, viz., IP addresses that
   remain fixed as the nodes move, but at the cost of potentially
   significant network overhead and increased delay because of the non-
   optimum network routing and tunneling involved.

   Some subnetworks may provide internal mobility, transparent to IP, as
   a feature of their own internal routing mechanisms. To the extent
   that these simplify routing at the IP layer, reduce the need for
   mechanisms like Mobile IP, or exploit mechanisms unique to the
   subnetwork, this is generally desirable. This is especially true when
   the subnetwork covers a relatively small geographic area and the
   users move rapidly between the attachment points within that area.
   Examples of internal mobility schemes include Ethernet switching and
   intra-system handoff in cellular telephony.

   However, if the subnetwork is physically large and connects to other
   parts of the Internet at multiple geographic points, care should be
   taken to optimize the wide-area routing of packets between nodes on
   the external Internet and nodes on the subnet. This is generally done
   with "nearest exit" routing strategies. Because a given subnetwork
   may be unaware of the actual physical location of a destination on
   another subnetwork, it simply routes packets bound for the other
   subnetwork to the nearest gateway between the two. This implies some
   awareness of IP addressing and routing within the subnetwork. The
   subnetwork may wish to use IP routing internally for wide area
   routing and restrict subnetwork-specific routing to constrained
   geographic areas where the effects of suboptimal routing are


   Many subnetworks provide their own internal routing mechanisms.
   Since routing is the major function of the Internet layer, the
   question naturally arises as to the proper division of function
   between routing at the Internet layer and routing in the subnet.

   In general, routing in a subnetwork and at IP is more complementary
   than competitive. Routing algorithms often have difficulty scaling to
   very large networks, and a division of labor between IP and a large
   subnetwork can often make the routing problem more tractable for

   Some subnetworks have special features that allow the use of more
   effective or responsive routing mechanisms that cannot be implemented
   in IP because of its need for generality. One example is the self-
   learning bridge algorithm widely used in Ethernet networks. Another
   is the "handoff" mechanism in cellular telephone networks,
   particularly the "soft handoff" scheme in IS-95 CDMA.

   On the other hand, routing optimality can suffer when a subnetwork's
   routing architecture hides internal structure that an IP router could
   have used to make more efficient decisions. Such situations occur
   most often when the subnetwork covers a large geographic area and
   includes links of widely varying capacities, but presents itself to
   IP as a single, fully-connected network with uniform metrics between
   border nodes.

   The subnetwork designer who decides to implement internal routing
   should also consider whether a custom routing algorithm is warranted,
   or if an existing Internet routing algorithm or protocol may suffice.
   Protocols and routing algorithms can be notoriously subtle, complex
   and difficult to implement correctly. Much work can be avoided if an
   existing protocol or off-the-shelf product can be readily used.

Security Considerations

   Security has become a high priority in the design and operation of
   the Internet. The Internet is vast, and countless organizations and
   individuals own and operate its various components.  A consensus has
   emerged for what might be called a "security placement principle": a
   security mechanism is most effective when it is placed as close as
   possible to, and under the direct control of the owner of, the asset
   that it protects.

   The most important conclusion that follows from this principle is
   that end-to-end security (e.g., confidentiality, integrity and access
   control) cannot be ensured with subnetwork security mechanisms.  Not
   only are end-to-end security mechanisms much more closely associated
   with the end-user assets they protect, they are also much more
   comprehensive. For example, end-to-end security mechanisms cover gaps
   that can appear when otherwise good subnetwork mechanisms are
   concatenated.  This is an important application of the end-to-end
   principle [SRC81].

   Several security mechanisms that can be used end-to-end have already
   been deployed in the Internet and are enjoying increasing use. The
   most important are the Secure Sockets Layer (SSL) [SSL2] [SSL3] and
   TLS [RFC2246] primarily used to protect web commerce; Pretty Good
   Privacy (PGP) [RFC1991], primarily used to protect and authenticate
   email and software distributions; the Secure Shell (SSH) [SSH], used
   for secure remote access and file transfer; and IPSEC [RFC2401], a
   general purpose encryption and authentication mechanism that sits
   just above IP and can be used by any IP application. (IPSEC can
   actually be used either on an end-to-end basis or between security
   gateways that do not include either or both end systems.)

   Nonetheless, end-to-end security mechanisms are not used as widely as
   might be desired. However, the group could not reach consensus on
   whether subnetwork designers should be actively encouraged to
   implement mechanisms to protect user data.

   The clear consensus of the working group held that subnetwork
   security mechanisms, especially when weak or incorrectly implemented
   [BGW], may actually be counterproductive.  The argument is that
   subnetwork security mechanisms can lull end users into a false sense
   of security, diminish the incentive to deploy effective end-to-end
   mechanisms, and encourage "risky" uses of the Internet that would not
   be made if users understood the inherent limits of subnetwork
   security mechanisms.

   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 high 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 media by attackers is difficult.  To meet
   the assumptions of such policies, explicit mechanisms are needed for
   links (especially shared media links) that lack physical protection.
   This, for example, is the rationale that underlies Wired Equivalent
   Privacy (WEP) in the IEEE 802.11 wireless LAN standard, and the
   Baseline Privacy Interface in the DOCSIS 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 NOT ensure) end-to-
   end security.

   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 [Schneier2] [Schneier3].  [Schneier1]
   describes protocols and algorithms which are considered well
   understood and believed to be sound.

   Poor design process, subtle design errors and flawed implementation
   can result in gaping vulnerabilities.  In recent years, a number of
   subnet standards have had problems exposed.  The following are
   examples of the mistakes that have been made:

   1. Use of weak and untested algorithms [Crypto9912] [BGW].  For a
   variety of reasons, algorithms were chosen which had subtle flaws
   that made them vulnerable to a variety of attacks.

   2. Use of 'security by obscurity' [Schneier4] [Crypto9912].  One
   common mistake is to assume that keeping cryptographic algorithms
   secret makes them more secure.  This is intuitive, but wrong.  Full
   public disclosure early in the design process attracts peer review by
   knowledgeable cryptographers.  Exposure of flaws by this review far
   outweighs any imagined benefit from forcing attackers to reverse
   engineer security algorithms.

   3. Inclusion of trapdoors [Schneier4] [Crypto9912].  Trapdoors are
   flaws surreptitiously left in an algorithm to allow it to be broken.
   This might be done to recover lost keys or to permit surreptitious
   access by governmental agencies.  Trapdoors can be discovered and
   exploited by malicious attackers.

   4. Sending passwords or other identifying information as clear text.
   For many years, analog cellular telephones could be cloned and used
   to steal service.  The cloners merely eavesdropped on the
   registration protocols that exchanged everything in clear text.

   5. Keys which are common to all systems on a subnet [BGW].

   6. Incorrect use of a sound mechanism.  For example [BGW], one subnet
   standard includes an initialization vector which is poorly designed
   and poorly specified.  A determined attacker can easily recover
   multiple ciphertexts encrypted with the same key stream and perform
   statistical attacks to decipher them.

   7. Identifying information sent in clear text that can be resolved to
   an individual, identifiable device. This creates a vulnerability to
   attacks targeted to that device (or its owner).

   8. Inability to renew and revoke shared secret information.

   9. Insufficient key length [Blaze96].

   10. Failure to address "man-in-the-middle" attacks, e.g., with mutual

   This list is by no means comprehensive.  Design problems are
   difficult to avoid, but expert review is generally invaluable in
   avoiding problems.

   In addition, well-designed security protocols can be compromised by
   implementation defects.  Examples of such defects include use of
   predictable pseudo-random numbers [RFC1750], vulnerability to buffer
   overflow attacks due to unsafe use of certain I/O system calls
   [WFBA2000], and inadvertent exposure of secret data.


   References of the form RFCnnnn are Internet Request for Comments
   (RFC) documents available online at

   [asym-id] H. Balakrishnan, V. N. Padmanabhan, G. Fairhurst, M,
   Sooriyabandara. "TCP Performance Implications of Network Path
   Asymmetry", work in progress as internet-draft, draft-ietf-pilc-
   asym-07.txt, November 2001.

   [APS99] Mark Allman, Vern Paxson, W. Richard Stevens.  "TCP
   Congestion Control". April 1999.  RFC 2581.

   [ARQ-DRAFT] Fairhurst, G., and L. Wood, Advice to link designers on
   link Automatic Repeat reQuest (ARQ), work in progress as internet-
   draft, draft-ietf-pilc-link-arq-issues-03.txt, August 2001.

   [ATMFTM] The ATM Forum, "Traffic Management Specification, Version
   4.0", April 1996, document af-tm-0056.000 (

   [BGW] Nikita Borisov, Ian Goldberg, David Wagner, "Security of the
   WEP Algorithm"

   [Blaze96] M. Blaze, W. Diffie, R. Rivest, B. Schneier, T. Shimomura,
   E. Thompson, M. Weiner, "Minimal Key Lengths for Symmetric Ciphers to
   Provide Adequate Commercial Security", available at

   [BPK98] Hari Balakrishnan, Venkata Padmanabhan, Randy H. Katz.  'The
   Effects of Asymmetry on TCP Performance."  ACM Mobile Networks and
   Applications (MONET), 1998.

   [BPS99] "Packet Reordering is Not Pathological Network Behavior", Jon
   C. R. Bennet, Craig Partridge, Nicholas Shectman, IEEE/ACM
   Transactions on Networking, Vol 7, No. 6, December 1999.

   [Crypto9912] Schneier, Bruce "European Cellular Encryption
   Algorithms" Crypto-Gram (December 15, 1999)

   [DIX] Digital Equipment Corp, Intel Corp, Xerox Corp, Ethernet Local
   Area Network Specification Version 2.0, November 1982.

   [DOCSIS2] Data-Over-Cable Service Interface Specifications, Radio
   Frequency Interface Specification 1.1, SP-RFIv1.1-I05-000714, July
   2000, Cable Television Laboratories, Inc.

   [EN301] ETSI, (European Broadcasting Union), Digital Video
   Broadcasting (DVB); DVB Specification for Data Broadcasting, 1997.
   Draft ETSI Standard EN 301 192 v1.1.1 (August 1997).

   [FB00] Firoiu V., and Borden M., "A Study of Active Queue Management
   for Congestion Control" to appear in Infocom 2000

   [IEEE8023] IEEE 802.3 CSMA/CD Access Method. Available from

   [IEEE80211] IEEE 802.11 Wireless LAN standard. Available from

   [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

   [ISO13181] ISO/IEC, ISO/IEC 13181-1: Information Technology - Generic
   coding of moving pictures and associated audio information,  1995,
   International Organization for Standardization and International
   Electrotechnical Commission.

   [Jac90] Van Jacobson.  Modified TCP Congestion Avoidance Algorithm.
   Email to the end2end-interest mailing list, April 1990.  URL:

   [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.

   [LKJK01] R. Ludwig, A. Konrad, A. D. Joseph, R. H. Katz, "Optimizing
   the End-to-End Performance of Reliable Flows over Wireless Links", To
   appear in ACM/Baltzer Wireless Networks Journal (Special issue:
   Selected papers from ACM/IEEE MOBICOM 99), available at

   [MBB00] May, M., Bonald, T., and Bolot, J-C., "Analytic Evaluation of
   RED Performance" to appear INFOCOM 2000

   [MBDL99] May, M., Bolot, J., Diot, C., and Lyles, B., "Reasons not to
   deploy RED", technical report, June 1999.

   [MG01] S. Mascolo, A. Grieco, "Easy RED: Improving Fairness via
   Simple Early Discard", to appear in IEEE Communication Letters.
   Available as

   [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.

   [MYR] Myricom Corporation home page,

   [OKM96] T. Ott, J.H.B. Kemperman, M. Mathis, "The Stationary Behavior
   of Ideal TCP Congestion Avoidance".

   [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,

   [RF95] Romanow, A., and Floyd, S., "Dynamics of TCP Traffic over ATM
   Networks".  IEEE JSAC, V. 13 N. 4, May 1995, p. 633-641.

   [RFC791] Jon Postel.  "Internet Protocol". September 1981.

   [RFC1144] Jacobson, V., "Compressing TCP/IP Headers for Low-Speed
   Serial Links," RFC 1144, February 1990.

   [RFC1191] J. Mogul, S. Deering. "Path MTU Discovery". November 1990.

   [RFC1435] S. Knowles. "IESG Advice from Experience with Path MTU
   Discovery".  March 1993.

   [RFC1577] M. Laubach.  "Classical IP and ARP over ATM". January 1994.

   [RFC1661] W. Simpson. "The Point-to-Point Protocol (PPP)". July 1994.

   [RFC1750] D. Eastlake, S. Crocker, J. Schiller, "Randomness
   Recommendations for Security", December 1994 (this will be obsoleted
   soon... check before submitting to RFC editor)

   [RFC1981] J. McCann, S. Deering, J. Mogul. "Path MTU Discovery for IP
   version 6".  August 1996.

   [RFC2246] T. Dierks, C. Allen. "The TLS Protocol Version 1.0".
   January 1999.

   [RFC2364] G. Gross et al. "PPP Over AAL5". July 1998.

   [RFC2393] A. Shacham et al. "IP Payload Compression Protocol
   (IPComp)". December 1998.

   [RFC2394] R. Pereira. "IP Payload Compression Using DEFLATE".
   December 1998.

   [RFC2395] R. Friend, R. Monsour. "IP Payload Compression Using LZS".
   December 1998.

   [RFC2440] J. Callas et al. "OpenPGP Message Format". November 1998.

   [RFC2507] M. Degermark, B. Nordgren, S. Pink. "IP Header
   Compression".  February 1999.

   [RFC2508] S. Casner, V. Jacobson. "Compressing IP/UDP/RTP Headers for
   Low-Speed Serial Links". February 1999.

   [RFC2581] M. Allman, V. Paxson, W. Stevens. "TCP Congestion Control".
   April 1999.

   [RFC2406] S. Kent, R. Atkinson. "P Encapsulating Security Payload
   (ESP)". November 1998.

   [RFC2616] R. Fielding et al. "Hypertext Transfer Protocol --
   HTTP/1.1". June 1999.

   [RFC2684] D. Grossman, J. Heinanen. "Multiprotocol Encapsulation over
   ATM Adaptation Layer 5". September 1999.

   [RFC2686] C. Bormann, "The Multi-Class Extension to Multi-Link PPP",
   September 1999.

   [RFC2687] C. Bormann, "PPP in a Real-time Oriented HDLC-like
   Framing", September 1999.

   [RFC2689] C. Bormann, "Providing Integrated Services over Low-bitrate
   Links", September 1999.

   [RFC2914] S. Floyd. "Congestion Control Principles". September 2000

   [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.

   [Schneier1] Schneier, Bruce, Applied Cryptography: Protocols,
   Algorithms and Source Code in C (John Wiley and Sons, October 1995).

   [Schneier2] Schneier, Bruce, Secrets and Lies: Digital Security in a
   Networked World (John Wiley & Sons, August 2000).

   [Schneier3] Schneier, Bruce "Why Cryptography is Harder Than it

   [SP2000] "When the CRC and TCP Checksum Disagree", Jonathan Stone &
   Craig Partridge, ACM CCR p309-321, September 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.

   [SSH] OpenSSH project web page,

   [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.

   [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
   from: "

   [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


   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 to this document: Carsten Bormann, Gorry
   Fairhurst, Aaron Falk, Dan Grossman, Reiner Ludwig, Jamshid Mahdavi,
   Saverio Mascolo, Gabriel Montenegro, Marie-Jose Montpetit, Joe Touch
   and Lloyd Wood.

Authors'  Addresses:

   Phil Karn, editor (