RFC 3819
Advice for Internet Subnetwork Designers
8. Reliability and Error Control
In the Internet architecture, the ultimate responsibility for error recovery is at the end points [SRC81]. The Internet may occasionally drop, corrupt, duplicate, or reorder packets, and the transport protocol (e.g., TCP) or application (e.g., if UDP is used as the transport protocol) must recover from these errors on an end-to-end basis [RFC3155]. Error recovery in the subnetwork is therefore justifiable only to the extent that it can enhance overall performance. It is important to recognize that a subnetwork can go too far in attempting to provide error recovery services in the Internet environment. Subnet reliability should be "lightweight", i.e., it only has to be "good enough", *not* perfect. In this section, we discuss how to analyze characteristics of a subnetwork to determine what is "good enough". The discussion below focuses on TCP, which is the most widely-used transport protocol in the Internet. It is widely believed (and is a stated goal within the IETF) that non-TCP transport protocols should attempt to be "TCP- friendly" and have many of the same performance characteristics. Thus, the discussion below should be applicable, even to portions of the Internet where TCP may not be the predominant protocol.8.1. TCP vs Link-Layer Retransmission
Error recovery involves the generation and transmission of redundant information computed from user data. Depending on how much redundant information is sent and how it is generated, the receiver can use it to reliably detect transmission errors, correct up to some maximum number of transmission errors, or both. The general approach is known as Error Control Coding, or ECC. The use of ECC to detect transmission errors so that retransmissions (hopefully without errors) can be requested is widely known as "ARQ" (Automatic Repeat Request). When enough ECC information is available to permit the receiver to correct some transmission errors without a retransmission, the approach is known as Forward Error Correction (FEC). Due to the greater complexity of the required ECC and the need to tailor its design to the characteristics of a specific modem and channel, FEC
has traditionally been implemented in special-purpose hardware integral to a modem. This effectively makes it part of the physical layer. Unlike ARQ, FEC was rarely used for telecommunications outside of space links prior to the 1990s. It is now nearly universal in telephone, cable and DSL modems, digital satellite links, and digital mobile telephones. FEC is also heavily used in optical and magnetic storage where "retransmissions" are not possible. Some systems use hybrid combinations of ARQ layered atop FEC; V.90 dialup modems (in the upstream direction) with V.42 error control are one example. Most errors are corrected by the trellis (FEC) code within the V.90 modem, and most remaining errors 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] [RFC3450-RFC3453] where conventional ARQ mechanisms are inefficient or difficult to implement. However, in this discussion, we will assume that if FEC is present, it is implemented within the physical layer. Depending on the layer in which it is implemented, error control can operate on an end-to-end basis or over a shorter span, such as a single link. TCP is the most important example of an end-to-end protocol that uses an ARQ strategy. Many link-layer protocols use ARQ, usually some flavor of HDLC [ISO3309]. Examples include the X.25 link layer, the AX.25 protocol used in amateur packet radio, 802.11 wireless LANs, and the reliable link layer specified in IEEE 802.2. Only end-to-end error recovery can ensure reliable service to the application (see Section 8). However, some subnetworks (e.g., many wireless links) also have link-layer error recovery as a performance enhancement [RFC3366]. For example, many cellular links have small physical frame sizes (< 100 bytes) and relatively high frame loss rates. Relying solely on end-to-end error recovery can clearly yield a performance degradation, as retransmissions across the end-to-end path take much longer to be received than when link layer retransmissions are used. Thus, link-layer error recovery can often increase end-to-end performance. As a result, link-layer and end- to-end recovery often co-exist; this can lead to the possibility of inefficient interactions between the two layers of ARQ protocols. This inter-layer "competition" might lead to the following wasteful situation. When the link layer retransmits (parts of) a packet, the link latency momentarily increases. Since TCP bases its
retransmission timeout on prior measurements of total end-to-end latency, including that of the link in question, this sudden increase in latency may trigger an unnecessary retransmission by TCP of a packet that the link layer is still retransmitting. Such spurious end-to-end retransmissions generate unnecessary load and reduce end- to-end throughput. As a result, the link layer may even have multiple copies of the same packet in the same link queue at the same time. In general, one could say the competing error recovery is caused by an inner control loop (link-layer error recovery) reacting to the same signal as an outer control loop (end-to-end error recovery) without any coordination between the loops. Note that this is solely an efficiency issue; TCP continues to provide reliable end-to-end delivery over such links. This raises the question of how persistent a link-layer sender should be in performing retransmission [RFC3366]. We define the link-layer (LL) ARQ persistency as the maximum time that a particular link will spend trying to transfer a packet before it can be discarded. This deliberately simplified definition says nothing about the 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. For link layers that do not or cannot differentiate between flows (e.g., due to network layer encryption), the LL ARQ persistency should be small. This avoids any harmful effects or performance degradation resulting from indiscriminate high persistence. A detailed discussion of these issues is provided in [RFC3366]. However, when a link layer can identify individual flows and apply ARQ selectively [LKJK02], then the link ARQ persistency should be high for a flow using reliable unicast transport protocols (e.g., TCP) and must be low for all other flows. Setting the link ARQ persistency larger than the largest link outage allows TCP to rapidly restore transmission without needing 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 the following 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.8.2. Recovery from Subnetwork Outages
Some types of subnetworks, particularly mobile radio, are subject to frequent temporary outages. For example, an active cellular data user may drive or walk into an area (such as a tunnel) that is out of range of any base station. No packets will be delivered successfully until the user returns to an area with coverage. The Internet protocols currently provide no standard way for a subnetwork to explicitly notify an upper layer protocol (e.g., TCP) that it is experiencing an outage rather than severe congestion. Under these circumstances TCP will, after each unsuccessful retransmission, wait even longer before trying again; this is its "exponential back-off" algorithm. Furthermore, TCP will not discover that the subnetwork outage has ended until its next retransmission attempt. If TCP has backed off, this may take some time. This can lead to extremely poor TCP performance over such subnetworks. It is therefore highly desirable that a subnetwork subject to outages does not silently discard packets during an outage. Ideally, the subnetwork should define an interface to the next higher layer (i.e., IP) that allows it to refuse packets during an outage, and to automatically ask IP for new packets when it is again able to deliver them. If it cannot do this, then the subnetwork should hold onto at least some of the packets it accepts during an outage and attempt to deliver them when the outage ends. When packets are discarded, IP should be notified so that the appropriate ICMP messages can be sent. Note that it is *not* necessary to completely avoid dropping packets during an outage. The purpose of holding onto a packet during an outage, either in the subnetwork or at the IP layer, is so that its eventual delivery will implicitly notify TCP that the subnetwork is again operational. This is to enhance performance, not to ensure reliability -- reliability, as discussed earlier, can only be ensured on an end-to-end basis. Only a few packets per TCP connection, including ACKs, need be held in this way to cause the TCP sender to recover from the additional losses once the flow resumes [RFC3366].
Because it would be a layering violation (and possibly a performance hit) for IP or a subnetwork layer to look at TCP headers (which would in any event be impossible if IPsec encryption [RFC2401] is in use), it would be reasonable for the IP or subnetwork layers to choose, as a design parameter, some small number of packets that will be retained during an outage.8.3. CRCs, Checksums and Error Detection
The TCP [RFC793], UDP [RFC768], ICMP, and IPv4 [RFC791] protocols all use the same simple 16-bit 1's complement checksum algorithm [RFC1071] to detect corrupted packets. The IPv4 header checksum protects only the IPv4 header, while the TCP, ICMP, and UDP checksums provide end-to-end error detection for both the transport pseudo header (including network and transport layer information) and the transport payload data. Protection of the data is optional for applications using UDP [RFC768] for IPv4, but is required for IPv6. The Internet checksum is not very strong from a coding theory standpoint, but it is easy to compute in software, and various proposals to replace the Internet checksums with stronger checksums have failed. However, it is known that undetected errors can and do occur in packets received by end hosts [SP2000]. To reduce processing costs, IPv6 has no IP header checksum. The destination host detects "important" errors in the IP header, such as the delivery of the packet to the wrong destination. This is done by including the IP source and destination addresses (pseudo header) in the computation of the checksum in the TCP or UDP header, a practice already performed in IPv4. Errors in other IPv6 header fields may go undetected within the network; this was considered a reasonable price to pay for a considerable reduction in the processing required by each router, and it was assumed that subnetworks would use a strong 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 for error detection. This is significantly more robust against most patterns of errors than the 16-bit Internet checksum. Note that the error detection properties of a specific CRC code diminish with increasing frame size. The Point-to-Point Protocol [RFC1662] requires support of a 16-bit CRC
for each link frame, with a 32-bit CRC as an option. (PPP is often used in conjunction with a dialup modem, which provides its own error control). Other subnetworks, including 802.3/Ethernet, AAL5/ATM, FDDI, Token Ring, and PPP over SONET/SDH all use a 32-bit CRC. Many subnetworks can also use other mechanisms to enhance the error detection capability of the link CRC (e.g., FEC in dialup modems, mobile radio and satellite channels). Any new subnetwork designed to carry IP should therefore provide error detection for each IP packet that is at least as strong as the 32-bit CRC specified in [ISO3309]. While this will achieve a very low undetected packet error rate due to transmission errors, it will not (and need not) achieve a very low packet loss rate as the Internet protocols are better suited to dealing with lost packets than to dealing with corrupted packets [SRC81]. Packet corruption may be, and is, also caused by bugs in host and router hardware and software. Even if every subnetwork implemented strong error detection, it is still essential that end-to-end checksums are used at the receiving end host [SP2000]. Designers of complex subnetworks consisting of internal links and packet switches should consider implementing error detection on an edge-to-edge basis to cover an entire SNDU (or IP packet). A CRC would be generated at the entry point to the subnetwork and checked at the exit endpoint. This may be used instead of, or in combination with, error detection at the interface to each physical link. An edge-to-edge check has the significant advantage of protecting against errors introduced anywhere within the subnetwork, not just within its transmission links. Examples of this approach include the way in which the Ethernet CRC-32 is handled by LAN bridges [802.1D]. ATM AAL5 [ITU-I363] also uses an edge-to-edge CRC-32. Some specific applications may be tolerant of residual errors in the data they exchange, but removal of the link CRC may expose the network to an undesirable increase in undetected errors in the IP and transport headers. Applications may also require a high level of error protection for control information exchanged by protocols acting above the transport layer. One example is a voice codec, which is robust against bit errors in the speech samples. For such mechanisms to work, the receiving application must be able to tolerate receiving corrupted data. This also requires that an application uses a mechanism to signal that payload corruption is permitted and to indicate the coverage (headers and data) required to be protected by the subnetwork CRC. The UDP-Lite protocol [RFC3828] is the first Internet standards track transport protocol supporting partial payload protection. Receipt of corrupt data by arbitrary
application protocols carries a serious danger that a subnet delivers data with errors that remain undetected by the application and hence corrupt the communicated data [SRC81].8.4. How TCP Works
One of TCP's functions is end-host based congestion control for the Internet. This is a critical part of the overall stability of the Internet, so it is important that link-layer designers understand TCP's congestion control algorithms. TCP assumes that, at the most abstract level, the network consists of links and queues. Queues provide output-buffering on links that are momentarily oversubscribed. They smooth instantaneous traffic bursts to fit the link bandwidth. When demand exceeds link capacity long enough to fill the queue, packets must be dropped. The traditional action of dropping the most recent packet ("tail dropping") is no longer recommended [RFC2309] [RFC2914], but it is still widely practiced. TCP uses sequence numbering and acknowledgments (ACKs) on an end-to-end basis to provide reliable, sequenced delivery. TCP ACKs are cumulative, i.e., each implicitly ACKs every segment received so far. If a packet with an unexpected sequence number is received, the ACK field in the packets returned by the receiver will cease to advance. Using an optional enhancement, TCP can send selective acknowledgments (SACKs) [RFC2018] to indicate which segments have arrived at the receiver. Since the most common cause of packet loss is congestion, TCP treats packet loss as an indication of potential Internet congestion along the path between TCP end hosts. This happens automatically, and the subnetwork need not know anything about IP or TCP. A subnetwork node simply drops packets whenever it must, though some packet-dropping strategies (e.g., RED) are more fair to competing flows than others. TCP recovers from packet losses in two different ways. The most important mechanism is the retransmission timeout. If an ACK fails to arrive after a certain period of time, TCP retransmits the oldest unacked packet. Taking this as a hint that the network is congested, TCP waits for the retransmission to be ACKed before it continues, and it gradually increases the number of packets in flight as long as a timeout does not occur again. A retransmission timeout can impose a significant performance penalty, as the sender is idle during the timeout interval and restarts with a congestion window of one TCP segment 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. A technique called "Explicit Congestion Notification" (ECN) [RFC3168] allows routers to directly signal congestion to hosts without dropping packets. This is done by setting a bit in the IP header. Since ECN support is likely to remain optional, the lack of an ECN bit must *never* be interpreted as a lack of congestion. Thus, for the foreseeable future, TCP must interpret a lost packet as a signal of congestion. The TCP "congestion avoidance" [RFC2581] algorithm maintains a congestion window (cwnd) controlling the amount of data TCP may have in flight at any moment. Reducing cwnd reduces the overall bandwidth obtained by the connection; similarly, raising cwnd increases performance, up to the limit of the available capacity. TCP probes for available network capacity by initially setting cwnd to one or two packets and then increasing cwnd by one packet for each ACK returned from the receiver. This is TCP's "slow start" mechanism. When a packet loss is detected (or congestion is signaled by other mechanisms), cwnd is reset to one and the slow start process is repeated until cwnd reaches one half of its previous setting before the reset. Cwnd continues to increase past this point, but at a much slower rate than before. If no further losses occur, cwnd will ultimately reach the window size advertised by the receiver. This is an "Additive Increase, Multiplicative Decrease" (AIMD) algorithm. The steep decrease of cwnd in response to congestion provides for network stability; the AIMD algorithm also provides for fairness between long running TCP connections sharing the same path.
8.5. TCP Performance Characteristics
Caveat Here we present a current "state-of-the-art" understanding of TCP performance. This analysis attempts to characterize the performance of TCP connections over links of varying characteristics. Link designers may wish to use the techniques in this section to predict what performance TCP/IP may achieve over a new link-layer design. Such analysis is encouraged. Because this is a relatively new analysis, and the theory is based on single-stream TCP connections under "ideal" conditions, it should be recognized that the results of such analysis may differ from actual performance in the Internet. That being said, we have done our best to provide the designers with helpful information to get an accurate picture of the capabilities and limitations of TCP under various conditions.8.5.1. The Formulae
The performance of TCP's AIMD Congestion Avoidance algorithm has been extensively analyzed. The current best formula for the performance of the specific algorithms used by Reno TCP (i.e., the TCP specified in [RFC2581]) is given by Padhye, et al. [PFTK98]. This formula is: MSS BW = -------------------------------------------------------- RTT*sqrt(1.33*p) + RTO*p*[1+32*p^2]*min[1,3*sqrt(.75*p)] where BW is the maximum TCP throughout achievable by an individual TCP flow MSS is the TCP segment size being used by the connection RTT is the end-to-end round trip time of the TCP connection RTO is the packet timeout (based on RTT) p is the packet loss rate for the path (i.e., .01 if there is 1% packet loss) Note that the speed of the links making up the Internet path does not explicitly appear in this formula. Attempting to send faster than the slowest link in the path causes the queue to grow at the transmitter driving the bottleneck. This increases the RTT, which in turn reduces the achievable throughput. This is currently considered to be the best approximate formula for Reno TCP performance. A further simplification of this formula is generally made by assuming that RTO is approximately 5*RTT.
TCP is constantly being improved. A simpler formula, which gives an upper bound on the performance of any AIMD algorithm which is likely to be implemented in TCP in the future, was derived by Ott, et al. [MSMO97]. MSS 1 BW = C --- ------- RTT sqrt(p) where C is 0.93.8.5.2. Assumptions
Both formulae assume that the TCP Receiver Window is not limiting the performance of the connection. Because the receiver window is entirely determined by end-hosts, we assume that hosts will maximize the announced receiver window to maximize their network performance. Both of these formulae allow BW to become infinite if there is no loss. However, an Internet path will drop packets at bottlenecked queues if the load is too high. Thus, a completely lossless TCP/IP network can never occur (unless the network is being underutilized). The RTT used is the arithmetic average, including queuing delays. The formulae are for a single TCP connection. If a path carries many TCP connections, each will follow the formulae above independently. The formulae assume long-running TCP connections. For connections that are extremely short (<10 packets) and don't lose any packets, performance is driven by the TCP slow-start algorithm. For connections of medium length, where on average only a few segments are lost, single connection performance will actually be slightly better than given by the formulae above. The difference between the simple and complex formulae above is that the complex formula includes the effects of TCP retransmission timeouts. For very low levels of packet loss (significantly less than 1%), timeouts are unlikely to occur, and the formulae lead to very similar results. At higher packet losses (1% and above), the complex formula gives a more accurate estimate of performance (which will always be significantly lower than the result from the simple formula). Note that these formulae break down as p approaches 100%.
8.5.3. Analysis of Link-Layer Effects on TCP Performance
Consider the following example: A designer invents a new wireless link layer which, on average, loses 1% of IP packets. The link layer supports packets of up to 1040 bytes, and has a one-way delay of 20 msec. If this link were to be used on an Internet path with a round trip time greater than 80ms, the upper bound may be computed by: For MSS, use 1000 bytes to exclude the 40 bytes of minimum IPv4 and TCP headers. For RTT, use 120 msec (80 msec for the Internet part, plus 20 msec each way for the new wireless link). For p, use .01. For C, assume 1. The simple formula gives: BW = (1000 * 8 bits) / (.120 sec * sqrt(.01)) = 666 kbit/sec The more complex formula gives: BW = 402.9 kbit/sec If this were a 2 Mb/s wireless LAN, the designers might be somewhat disappointed. Some observations on performance: 1. We have assumed that the packet losses on the link layer are interpreted as congestion by TCP. This is a "fact of life" that must be accepted. 2. The equations for TCP performance are all expressed in terms of packet loss, but many subnetwork designers think in terms of bit-error ratio. *If* channel bit errors are independent, then the probability of a packet being corrupted is: p = 1 - ([1 - BER]^[FRAME_SIZE*8]) Here we assume FRAME_SIZE is in bytes and "^" represents exponentiation. It includes the user data and all headers (TCP,IP and subnetwork). (Note: this analysis assumes the
subnetwork does not perform ARQ or transparent fragmentation
[RFC3366].) If the inequality
BER * [FRAME_SIZE*8] << 1
holds, the packet loss probability p can be approximated by:
p = BER * [FRAME_SIZE*8]
These equations can be used to apply BER to the performance
equations above.
Note that FRAME_SIZE can vary from one packet to the next. Small
packets (such as TCP acks) generally have a smaller probability
of packet error than, say, a TCP packet carrying one MSS (maximum
segment size) of user data. A flow of small TCP acks can be
expected to be slightly more reliable than a stream of larger TCP
data segments.
It bears repeating that the above analysis assumes that bit
errors are statistically independent. Because this is not true
for many real links, our computation of p is actually an upper
bound, not the exact probability of packet loss.
There are many reasons why bit errors are not independent on real
links. Many radio links are affected by propagation fading or by
interference that lasts over many bit times. Also, links with
Forward Error Correction (FEC) generally have very non-uniform
bit error distributions that depend on the type of FEC, but in
general the uncorrected errors tend to occur in bursts even when
channel symbol errors are independent. In all such cases, our
computation of p from BER can only place an upper limit on the
packet loss rate.
If the distribution of errors under the FEC scheme is known, one
could apply the same type of analysis as above, using the correct
distribution function for the BER. It is more likely in these
FEC cases, however, that empirical methods are needed to
determine the actual packet loss rate.
3. Note that the packet size plays an important role. If the
subnetwork loss characteristics are such that large packets have
the same probability of loss as smaller packets, then larger
packets will yield improved performance.
4. We have chosen a specific RTT that might occur on a wide-area
Internet path within the USA. It is important to recognize that
a variety of RTT values are experienced in the Internet.
For example, RTTs are typically less than 10 msec in a wired LAN
environment when communicating with a local host. International
connections may have RTTs of 200 msec or more. Modems and other
low-capacity links can add considerable delay due to their long
packet transmission (serialisation) times.
Links over geostationary repeater satellites have one-way speed-
of-light delays of around 250ms, a minimum of 125ms propagation
delay up to the satellite and 125ms down. The RTT of an end-to-
end TCP connection that includes such a link can be expected to
be greater 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.
9. Quality-of-Service (QoS) considerations
It is generally recognized that specific service guarantees are
needed to support real-time multimedia, toll-quality telephony, and
other performance-critical applications. The provision of such
Quality of Service guarantees in the Internet is an active area of
research and standardization. The IETF has not converged on a single
service model, set of services, or single mechanism that will offer
useful guarantees to applications and be scalable to the Internet.
Indeed, the IETF does not have a single definition of Quality of
Service. [RFC2990] represents a current understanding of the
challenges in architecting QoS for the Internet.
There are presently two architectural approaches to providing
mechanisms for QoS support in the Internet.
IP Integrated Services (Intserv) [RFC1633] provides fine-grained
service guarantees to individual flows. Flows are identified by a
flow specification (flowspec), which creates a stateful association
between individual packets by matching fields in the packet header.
Capacity is reserved for the flow, and appropriate traffic
conditioning and scheduling is installed in routers along the path.
The ReSerVation Protocol (RSVP) [RFC2205] [RFC2210] is usually, but
need not necessarily be, used to install the flow QoS state. Intserv
defines two services, in addition to the Default (best effort)
service.
1. Guaranteed Service (GS) [RFC2212] 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. 2. Controlled Load Service (CLS) [RFC2211] offers delay and packet loss equivalent to that of an unloaded network to flows that conform to a TSpec, but no hard bounds. Non-conforming packets are forwarded on a best-effort basis. Intserv requires installation of state information in every participating router. Performance guarantees cannot be made unless this state is present in every router along the path. This, along with RSVP processing and the need for usage-based accounting, is believed to have scalability problems, particularly in the core of the Internet [RFC2208]. IP Differentiated Services (Diffserv) [RFC2475] provides a "toolkit" offering coarse-grained controls to aggregates of flows. Diffserv in itself does *not* provide QoS guarantees, but can be used to construct services with QoS guarantees across a Diffserv domain. Diffserv attempts to address the scaling issues associated with Intserv by requiring state awareness only at the edge of a Diffserv domain. At the edge, packets are classified into flows, and the flows are conditioned (marked, policed, or shaped) to a traffic conditioning specification (TCS). A Diffserv Codepoint (DSCP), identifying a per-hop behavior (PHB), is set in each packet header. The DSCP is carried in the DS-field, subsuming six bits of the former Type-of-Service (ToS) byte [RFC791] of the IP header [RFC2474]. The PHB denotes the forwarding behavior to be applied to the packet in each node in the Diffserv domain. Although there is a "recommended" DSCP associated with each PHB, the mappings from DSCPs to PHBs are defined by the DS-domain. In fact, there can be several DSCPs associated with the same PHB. Diffserv presently defines three PHBs. 1. The class selector PHB [RFC2474] replaces the IP precedence field of the former ToS byte. It offers relative forwarding priorities. 2. The Expedited Forwarding (EF) PHB [RFC3246] [RFC3248] 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.
3. 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 [RFC2990]) 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: 1. the concatenation of the QoS guarantees offered by the subnets along the path of a flow. This implies that a subnet may wish to offer multiple services (with different QoS guarantees) to the IP layer, which can then determine which flows use which subnet service. To put it another way, forwarding behavior in the subnet needs to be "clued" by the forwarding behavior (service or PHB) at the IP layer, and 2. the operation of a set of cooperating mechanisms, such as bandwidth reservation and admission control, policy management, traffic classification, traffic conditioning (marking, policing and/or shaping), selective discard, queuing, and scheduling. Note that support for QoS in subnets may require similar mechanisms, especially when these subnets are general topology subnets (e.g., ATM, frame relay, or MPLS) or shared media subnets. Many subnetwork designers face inherent tradeoffs between delay, throughput, reliability, and cost. Other subnetworks have parameters that manage bandwidth, internal connection state, and the like. Therefore, the following subnetwork capabilities may be desirable, although some might be trivial or moot if the subnet is a dedicated point-to-point link. 1. The subnetwork should have the ability to reserve bandwidth for a connection or flow and schedule packets accordingly. 2. 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.
3. If a connection or flow does not use its reserved bandwidth at a
given time, the unused bandwidth should be available for other
flows.
4. 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.
5. 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 selection of more or less error control
and/or to selection of particular error control mechanisms within
the subnetwork.
6. 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 be best expressed in terms of the TSpec/RSpec model
described in [RFC2212].
7. Support of the Diffserv class selectors [RFC2474] suggests that
the subnet might consider mechanisms that support priorities.
10. Fairness vs Performance
Subnetwork designers should be aware of the tradeoffs between
fairness and efficiency inherent in many transmission scheduling
algorithms. For example, many local area networks use contention
protocols to resolve access to a shared transmission channel. These
protocols represent overhead. While limiting the amount of data that
a subnet node may transmit per contention cycle helps assure timely
access to the channel for each subnet node, it also increases
contention overhead per unit of data sent.
In some mobile radio networks, capacity is limited by interference,
which in turn depends on average transmitter power. Some receivers
may require considerably more transmitter power (generating more
interference and consuming more channel capacity) than others.
In each case, the scheduling algorithm designer must balance
competing objectives: providing a fair share of capacity to each
subnet node while maximizing the total capacity of the network. One
approach for balancing performance and fairness is outlined in
[ES00].
11. Delay Characteristics
The TCP sender bases its retransmission timeout (RTO) on measurements of the round trip delay experienced by previous packets. This allows TCP to adapt automatically to the very wide range of delays found on the Internet. The recommended algorithms are described in [RFC2988]. Evaluations of TCP's retransmission timer can be found in [AP99] and [LS00]. These algorithms model the delay along an Internet path as a normally-distributed random variable with a 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, because it is easier to compute.) The goal is to compute an 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 an RTO that is much larger than the mean of the measured delays. If the packet loss rate is low, the large RTO is of little consequence, as timeouts occur only rarely. Conversely, if the path delay variance is low, then TCP recovers quickly from lost packets; again, the algorithm works well. However, when delay variance and the packet loss rate are both high, these algorithms perform poorly, especially when the mean delay is also high. Because TCP uses returning acknowledgments as a "clock" to time the transmission of additional data, excessively high delays (even if the delay variance is low) also affect TCP's ability to fully utilize a high-speed transmission pipe. It also slows 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. While ARQ increases delay variance, FEC
does not. However, FEC (especially when combined with interleaving) often increases mean delay, even on good channels where ARQ retransmissions are not needed and ARQ would not increase either the delay or the delay variance. The tradeoffs among these error control mechanisms and their interactions with TCP can be quite complex, and are the subject of much ongoing research. We therefore recommend that subnetwork designers provide as much flexibility as possible in the implementation of these mechanisms, and provide access to them as discussed above in the section on Quality of Service.12. Bandwidth Asymmetries
Some subnetworks may provide asymmetric bandwidth (or may cause TCP packet flows to experience asymmetry in the capacity) and the Internet protocol suite will generally still work fine. However, there is a case when such a scenario reduces TCP performance. Since TCP data segments are "clocked" out by returning acknowledgments, TCP senders are limited by the rate at which ACKs can be returned [BPK98]. Therefore, when the ratio of the available capacity of the Internet path carrying the data to the bandwidth of the return path of the acknowledgments is too large, the slow return of the ACKs directly impacts performance. Since ACKs are generally smaller than data segments, TCP can tolerate some asymmetry, but as a general rule, designers of subnetworks should be aware that subnetworks with significant asymmetry can result in reduced performance, unless issues are taken to mitigate this [RFC3449]. Several strategies have been identified for reducing the impact of asymmetry of the network path between two TCP end hosts, e.g., [RFC3449]. These techniques attempt to reduce the number of ACKs transmitted over the return path (low bandwidth channel) by changes at the end host(s), and/or by modification of subnetwork packet forwarding. While these solutions may mitigate the performance issues caused by asymmetric subnetworks, they do have associated cost and may have other implications. A fuller discussion of strategies and their implications is provided in [RFC3449].13. Buffering, flow and congestion control
Many subnets include multiple links with varying traffic demands and possibly different transmission speeds. At each link there must be a queuing system, including buffering, scheduling, and a capability to discard excess subnet packets. These queues may also be part of a subnet flow control or congestion control scheme.
For the purpose of this discussion, we talk about packets without regard to whether they refer to a complete IP packet or a subnetwork frame. At each queue, a packet experiences a delay that depends on competing traffic and the scheduling discipline, and is subjected to a local discarding policy. Some subnets may have flow or congestion control mechanisms in addition to packet dropping. Such mechanisms can operate on components in the subnet layer, such as schedulers, shapers, or discarders, and can affect the operation of IP forwarders at the edges of the subnet. However, with the exception of Explicit Congestion Notification [RFC3168] (discussed below), IP has no way to pass explicit congestion or flow control signals to TCP. TCP traffic, especially aggregated TCP traffic, is bursty. As a result, instantaneous queue depths can vary dramatically, even in nominally stable networks. For optimal performance, packets should be dropped in a controlled fashion, not just when buffer space is unavailable. How much buffer space should be supplied is still a matter of debate, but as a rule of thumb, each node should have enough buffering to hold one link_bandwidth*link_delay product's worth of data for each TCP connection sharing the link. This is often difficult to estimate, since it depends on parameters beyond the subnetwork's control or knowledge. Internet nodes generally do not implement admission control policies, and cannot limit the number of TCP connections that use them. In general, it is wise to err in favor of too much buffering rather than too little. It may also be useful for subnets to incorporate mechanisms that measure propagation delays to assist in buffer sizing calculations. There is a rough consensus in the research community that active queue management is important to improving fairness, link utilization, and throughput [RFC2309]. Although there are questions and concerns about the effectiveness of active queue management (e.g., [MBDL99]), it is widely considered an improvement over tail- drop discard policies. One form of active queue management is the Random Early Detection (RED) algorithm [RED93], 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, the probability of packets that are dropped is proportional to the amount by which the average queue depth exceeds min_th.
When this average queue depth is equal to max_th, the drop
probability is equal to a configurable parameter max_p.
When this average queue depth is greater than max_th, packets are
always dropped.
Numerous variants on RED appear in the literature, and there are
other active queue management algorithms which claim various
advantages over RED [GM02].
With an active queue management algorithm, dropped packets become a
feedback signal to trigger more appropriate congestion behavior by
the TCPs in the end hosts. Randomization of dropping tends to break
up the observed tendency of TCP windows belonging to different TCP
connections to become synchronized by correlated drops, and it also
imposes a degree of fairness on those connections that implement TCP
congestion avoidance properly. Another important property of active
queue management algorithms is that they attempt to keep average
queue depths short while accommodating large short-term bursts.
Since TCP neither knows nor cares whether congestive packet loss
occurs at the IP layer or in a subnet, it may be advisable for
subnets that perform queuing and discarding to consider implementing
some form of active queue management. This is especially true if
large aggregates of TCP connections are likely to share the same
queue. However, active queue management may be less effective in the
case of many queues carrying smaller aggregates of TCP connections,
e.g., in an ATM switch that implements per-VC queuing.
Note that the performance of active queue management algorithms is
highly sensitive to settings of configurable parameters, and also to
factors such as RTT [MBB00] [FB00].
Some subnets, most notably ATM, perform segmentation and reassembly
at the subnetwork edges. Care should be taken here in designing
discard policies. If the subnet discards a fragment of an IP packet,
then the remaining fragments become an unproductive load on the
subnet that can markedly degrade end-to-end performance [RF95].
Subnetworks should therefore attempt to discard these extra fragments
whenever one of them must be discarded. If the IP packet has already
been partially forwarded when discarding becomes necessary, then
every remaining fragment except the one marking the end of the IP
packet should also be discarded. For ATM subnets, this specifically
means using Early Packet Discard and Partial Packet Discard [ATMFTM].
Some subnets include flow control mechanisms that effectively require
that the rate of traffic flows be shaped upon entry to the subnet.
One example of such a subnet mechanism is in the ATM Available Bit
rate (ABR) service category [ATMFTM]. Such flow control mechanisms have the effect of making the subnet nearly lossless by pushing congestion into the IP routers at the edges of the subnet. In such a case, adequate buffering and discard policies are needed in these routers to deal with a subnet that appears to have varying bandwidth. Whether there is a 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 leading 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 that used IP over ABR to carry TCP traffic have claimed it to be successful, but have published no results. Another possible approach to flow control in the subnet would be to work with TCP Explicit Congestion Notification (ECN) semantics [RFC3168] through utilizing explicit congestion indicators in subnet frames. Routers at the edges of the subnet, rather than shaping, would set the explicit congestion bit in those IP packets that are received in subnet frames that have an ECN indication. Nodes in the subnet would need to implement an active queue management protocol that marks subnet frames instead of dropping them. ECN is currently a proposed standard, but it is not yet widely deployed.14. Compression
Application data compression is a function that can usually be omitted in the subnetwork. The endpoints typically have more CPU and memory resources to run a compression algorithm and a better understanding of what is being compressed. End-to-end compression benefits every network element in the path, while subnetwork-layer compression, by definition, benefits only a single subnetwork. Data presented to the subnetwork layer may already be in a 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 [RFC3173] supports DEFLATE [RFC2394] and LZS [RFC2395]). Compression at the subnetwork edges is of no benefit for any of these cases. The subnetwork may also process data that has been encrypted by the application (OpenPGP [RFC2440] or S/MIME [RFC2633]), just above TCP (SSL, TLS [RFC2246]), or just above IP (IPsec ESP [RFC2406]).
Ciphers generate high-entropy bit streams lacking any patterns that can be exploited by a compression algorithm. However, much data is still transmitted uncompressed over the Internet, so subnetwork compression may be beneficial. Any subnetwork compression algorithm must not expand uncompressible data, e.g., data that has already been compressed or encrypted. We make a strong recommendation that subnetworks operating at low speed or with small MTUs compress IP and transport-level headers (TCP and UDP) using several header compression schemes developed within the IETF [RFC3150]. An uncompressed 40-byte TCP/IP header takes about 33 milliseconds to send at 9600 bps. "VJ" TCP/IP header compression [RFC1144] compresses most headers to 3-5 bytes, reducing transmission time to several milliseconds on dialup modem links. This is especially beneficial for small, latency-sensitive packets in interactive sessions. Similarly, RTP compression schemes, such as CRTP [RFC2508] and ROHC [RFC3095], compress most IP/UDP/RTP headers to 1-4 bytes. The resulting savings are especially significant when audio packets are kept small to minimize store-and-forward latency. Designers should consider the effect of the subnetwork error rate on the performance of header compression. TCP ordinarily recovers from lost packets by retransmitting only those packets that were actually lost; packets arriving correctly after a packet loss are kept on a resequencing queue and do not need to be retransmitted. In VJ TCP/IP [RFC1144] header compression, however, the receiver cannot explicitly notify a sender of data corruption and subsequent loss of synchronization between compressor and decompressor. It relies instead on TCP retransmission to re-synchronize the decompressor. After a packet is lost, the decompressor must discard every subsequent packet, even if the subnetwork makes no further errors, until the sending TCP retransmits to re-synchronize the decompressor. This effect can substantially magnify the effect of subnetwork packet losses if the sending TCP window is large, as it will often be on a path with a large bandwidth*delay product [LRKOJ99]. Alternate header compression schemes, such as those described in [RFC2507], include an explicit request for retransmission of an uncompressed packet to allow decompressor resynchronization without waiting for a TCP retransmission. However, these schemes are not yet in widespread use. Both TCP header compression schemes do not compress widely-used TCP options such as selective acknowledgements (SACK). Both fail to compress TCP traffic that makes use of explicit congestion
notification (ECN). Work is under way in the IETF ROHC WG to address these shortcomings in a ROHC header compression scheme for TCP [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 the end-to-end path round-trip time. For links that combine significant error rates with latencies that require multiple packets to be in flight at a time, this leads to significant error propagation, i.e., subsequent losses caused by an 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. For example, within a talk spurt, only extended losses of (depending on the mode chosen) 12-64 packets typically cause error propagation.
(page 36 continued on part 3)
