| Transport Services (tsv) | K. De Schepper |
| Internet-Draft | Nokia Bell Labs |
| Intended status: Experimental | B. Briscoe, Ed. |
| Expires: September 12, 2019 | CableLabs |
| March 11, 2019 |
Identifying Modified Explicit Congestion Notification (ECN) Semantics for Ultra-Low Queuing Delay (L4S)
draft-ietf-tsvwg-ecn-l4s-id-06
This specification defines the identifier to be used on IP packets for a new network service called low latency, low loss and scalable throughput (L4S). It is similar to the original (or 'Classic') Explicit Congestion Notification (ECN). 'Classic' ECN marking was required to be equivalent to a drop, both when applied in the network and when responded to by a transport. Unlike 'Classic' ECN marking, for packets carrying the L4S identifier, the network applies marking more immediately and more aggressively than drop, and the transport response to each mark is reduced and smoothed relative to that for drop. The two changes counterbalance each other so that the throughput of an L4S flow will be roughly the same as a 'Classic' flow under the same conditions. However, the much more frequent control signals and the finer responses to them result in ultra-low queuing delay without compromising link utilization, and low delay is maintained during high load. Examples of new active queue management (AQM) marking algorithms and examples of new transports (whether TCP-like or real-time) are specified separately. The new L4S identifier is the key piece that enables them to interwork and distinguishes them from 'Classic' traffic. It gives an incremental migration path so that existing 'Classic' TCP traffic will be no worse off, but it can be prevented from degrading the ultra-low delay and loss of the new scalable transports.
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This specification defines the identifier to be used on IP packets for a new network service called low latency, low loss and scalable throughput (L4S). It is similar to the original (or 'Classic') Explicit Congestion Notification (ECN [RFC3168]). 'Classic' ECN marking was required to be equivalent to a drop, both when applied in the network and when responded to by a transport. Unlike 'Classic' ECN marking, the network applies L4S marking more immediately and more aggressively than drop, and the transport response to each mark is reduced and smoothed relative to that for drop. The two changes counterbalance each other so that the throughput of an L4S flow will be roughly the same as a 'Classic' flow under the same conditions. Nonetheless, the much more frequent control signals and the finer responses to them result in ultra-low queuing delay without compromising link utilization, and low delay is maintained during high load.
An example of a scalable congestion control that would enable the L4S service is Data Centre TCP (DCTCP), which until now has been applicable solely to controlled environments like data centres [RFC8257], because it is too aggressive to co-exist with existing TCP. The DualQ Coupled AQM, which is defined in a complementary experimental specification [I-D.ietf-tsvwg-aqm-dualq-coupled], is an AQM framework that enables scalable congestion controls like DCTCP to co-exist with existing traffic, each getting roughly the same flow rate when they compete under similar conditions. Note that a transport such as DCTCP is still not safe to deploy on the Internet unless it satisfies the requirements listed in Section 4. Also note that L4S is not only for elastic TCP-like traffic - there are scalable congestion controls for real-time media, such as the L4S variant of the SCReAM [RFC8298] real-time media congestion avoidance technique (RMCAT).
The new L4S identifier is the key piece that enables L4S hosts and L4S network nodes to interwork and distinguishes their traffic from 'Classic' traffic. It gives an incremental migration path so that existing 'Classic' TCP traffic will be no worse off, but it can be prevented from degrading the ultra-low delay and loss of the new scalable congestion controls. The performance improvement is so great that it is motivating initial deployment of the separate parts of this system.
Latency is becoming the critical performance factor for many (most?) applications on the public Internet, e.g. interactive Web, Web services, voice, conversational video, interactive video, interactive remote presence, instant messaging, online gaming, remote desktop, cloud-based applications, and video-assisted remote control of machinery and industrial processes. In the developed world, further increases in access network bit-rate offer diminishing returns, whereas latency is still a multi-faceted problem. In the last decade or so, much has been done to reduce propagation time by placing caches or servers closer to users. However, queuing remains a major intermittent component of latency.
The Diffserv architecture provides Expedited Forwarding [RFC3246], so that low latency traffic can jump the queue of other traffic. However, on access links dedicated to individual sites (homes, small enterprises or mobile devices), often all traffic at any one time will be latency-sensitive. Then Diffserv is of little use. Instead, we need to remove the causes of any unnecessary delay.
The bufferbloat project has shown that excessively-large buffering ('bufferbloat') has been introducing significantly more delay than the underlying propagation time. These delays appear only intermittently—only when a capacity-seeking (e.g. TCP) flow is long enough for the queue to fill the buffer, making every packet in other flows sharing the buffer sit through the queue.
Active queue management (AQM) was originally developed to solve this problem (and others). Unlike Diffserv, which gives low latency to some traffic at the expense of others, AQM controls latency for all traffic in a class. In general, AQMs introduce an increasing level of discard from the buffer the longer the queue persists above a shallow threshold. This gives sufficient signals to capacity-seeking (aka. greedy) flows to keep the buffer empty for its intended purpose: absorbing bursts. However, RED [RFC2309] and other algorithms from the 1990s were sensitive to their configuration and hard to set correctly. So, AQM was not widely deployed.
More recent state-of-the-art AQMs, e.g. fq_CoDel [RFC8290], PIE [RFC8033], Adaptive RED [ARED01], are easier to configure, because they define the queuing threshold in time not bytes, so it is invariant for different link rates. However, no matter how good the AQM, the sawtoothing rate of TCP will either cause queuing delay to vary or cause the link to be under-utilized. Even with a perfectly tuned AQM, the additional queuing delay will be of the same order as the underlying speed-of-light delay across the network. Flow-queuing can isolate one flow from another, but it cannot isolate a TCP flow from the delay variations it inflicts on itself, and it has other problems - it overrides the flow rate decisions of variable rate video applications, it does not recognise the flows within IPSec VPN tunnels and it is relatively expensive to implement.
Latency is not our only concern: It was known when TCP was first developed that it would not scale to high bandwidth-delay products [TCP-CA]. Given regular broadband bit-rates over WAN distances are already [RFC3649] beyond the scaling range of 'Classic' TCP Reno, 'less unscalable' Cubic [RFC8312] and Compound [I-D.sridharan-tcpm-ctcp] variants of TCP have been successfully deployed. However, these are now approaching their scaling limits. Unfortunately, fully scalable congestion controls such as DCTCP [RFC8257] cause 'Classic' TCP to starve itself, which is why they have been confined to private data centres or research testbeds (until now).
It turns out that a TCP algorithm like DCTCP that solves the latency problem also solves TCP's scalability problem. The finer sawteeth have low amplitude, so they cause very little queuing delay variation and the number of sawteeth per round trip remains invariant, which maintains constant tight control as flow-rate scales. A supporting paper [DCttH15] gives the full explanation of why the design solves both the latency and the scaling problems, both in plain English and in more precise mathematical form. The explanation is summarised without the maths in the L4S architecture document [I-D.ietf-tsvwg-l4s-arch].
The key words "MUST", "MUST NOT", "REQUIRED", "SHALL", "SHALL NOT", "SHOULD", "SHOULD NOT", "RECOMMENDED", "NOT RECOMMENDED", "MAY", and "OPTIONAL" in this document are to be interpreted as described in [RFC2119]. In this document, these words will appear with that interpretation only when in ALL CAPS. Lower case uses of these words are not to be interpreted as carrying RFC-2119 significance.
The new L4S identifier defined in this specification is applicable for IPv4 and IPv6 packets (as for classic ECN [RFC3168]). It is applicable for the unicast, multicast and anycast forwarding modes.
The L4S identifier is an orthogonal packet classification to the Differentiated Services Code Point (DSCP [RFC2474]). Section 5.4 explains what this means in practice.
This document is intended for experimental status, so it does not update any standards track RFCs. Therefore it depends on [RFC8311], which is a standards track specification that:
This subsection briefly records the process that led to a consensus choice of L4S identifier, selected from all the alternatives in Appendix B.
Ideally, the identifier for packets using the Low Latency, Low Loss, Scalable throughput (L4S) service ought to meet the following requirements:
Whether the identifier would be recoverable if the experiment failed is a factor that could be taken into account. However, this has not been made a requirement, because that would favour schemes that would be easier to fail, rather than those more likely to succeed.
It is recognised that the chosen identifier is unlikely to satisfy all these requirements, particularly given the limited space left in the IP header. Therefore a compromise will be necessary, which is why all the requirements are expressed with the word 'SHOULD' not 'MUST'. Appendix B discusses the pros and cons of the compromises made in various competing identification schemes against the above requirements.
On the basis of this analysis, "ECT(1) and CE codepoints" is the best compromise. Therefore this scheme is defined in detail in the following sections, while Appendix B records the rationale for this decision.
The L4S treatment is an experimental track alternative packet marking treatment [RFC4774] to the classic ECN treatment [RFC3168], which has been updated by [RFC8311] to allow this experiment (amongst others). Like classic ECN, L4S ECN identifies both network and host behaviour: it identifies the marking treatment that network nodes are expected to apply to L4S packets, and it identifies packets that have been sent from hosts that are expected to comply with a broad type of sending behaviour.
For a packet to receive L4S treatment as it is forwarded, the sender sets the ECN field in the IP header to the ECT(1) codepoint. See Section 4 for full transport layer behaviour requirements, including feedback and congestion response.
A network node that implements the L4S service normally classifies arriving ECT(1) and CE packets for L4S treatment. See Section 5 for full network element behaviour requirements, including classification, ECN-marking and interaction of the L4S identifier with other identifiers and per-hop behaviours.
A sender that wishes a packet to receive L4S treatment as it is forwarded, MUST set the ECN field in the IP header (v4 or v6) to the ECT(1) codepoint.
In general, a scalable congestion control needs feedback of the extent of CE marking on the forward path. When ECN was added to TCP [RFC3168], the feedback method reported no more than one CE mark per round trip. Some transport protocols derived from TCP mimic this behaviour while others report the accurate extent of TCP marking. This means that some transport protocols will need to be updated as a prerequisite for scalable congestion control. The position for a few well-known transport protocols is given below.
As a condition for a host to send packets with the L4S identifier (ECT(1)), it SHOULD implement a congestion control behaviour that ensures the flow rate is inversely proportional to the proportion of bytes in packets marked with the CE codepoint. This is termed a scalable congestion control, because the number of control signals (ECN marks) per round trip remains roughly constant for any flow rate. As with all transport behaviours, a detailed specification will need to be defined for each type of transport or application, including the timescale over which the proportionality is averaged, and control of burstiness. The inverse proportionality requirement above is worded as a 'SHOULD' rather than a 'MUST' to allow reasonable flexibility when defining these specifications.
Data Center TCP (DCTCP [RFC8257]) and the L4S variant of SCReAM [RFC8298] are examples of a scalable congestion controls.
Each sender in a session can use a scalable congestion control independently of the congestion control used by the receiver(s) when they send data. Therefore there might be ECT(1) packets in one direction and ECT(0) or Not-ECT in the other.
In order to coexist safely with other Internet traffic, a scalable congestion control MUST NOT tag its packets with the ECT(1) codepoint unless it complies with the following bulleted requirements. The specification of a particular scalable congestion control MUST describe in detail how it satisfies each requirement:
A network node that implements the L4S service MUST classify arriving ECT(1) packets for L4S treatment and, other than in the exceptional case referred to next, it MUST classify arriving CE packets for L4S treatment as well. CE packets might have originated as ECT(1) or ECT(0), but the above rule to classify them as if they originated as ECT(1) is the safe choice (see Appendix B.1 for rationale). The exception is where some flow-aware in-network mechanism happens to be available for distinguishing CE packets that originated as ECT(0), as described in Section 5.3, but there is no implication that such a mechanism is necessary.
An L4S AQM treatment follows similar codepoint transition rules to those in RFC 3168. Specifically, the ECT(1) codepoint MUST NOT be changed to any other codepoint than CE, and CE MUST NOT be changed to any other codepoint. An ECT(1) packet is classified as ECN-capable and, if congestion increases, an L4S AQM algorithm will mark the ECN field as CE for an increasing proportion of packets, otherwise forwarding packets unchanged as ECT(1). Necessary conditions for an L4S marking treatment are defined in Section 5.2. Under persistent overload an L4S marking treatment SHOULD turn off ECN marking, using drop as a congestion signal until the overload episode has subsided, as recommended for all AQMs in [RFC7567] (Section 4.2.1), which follows the similar advice in RFC 3168 (Section 7).
For backward compatibility in uncontrolled environments, a network node that implements the L4S treatment MUST also implement a classic AQM treatment. It MUST classify arriving ECT(0) and Not-ECT packets for treatment by the Classic AQM (see the discussion of the classifier for the dual-queue coupled AQM in [I-D.ietf-tsvwg-aqm-dualq-coupled]). Classic treatment means that the AQM will mark ECT(0) packets under the same conditions as it would drop Not-ECT packets [RFC3168].
The likelihood that an AQM drops a Not-ECT Classic packet (p_C) MUST be roughly proportional to the square of the likelihood that it would have marked it if it had been an L4S packet (p_L) {ToDo cross-ref to new section in l4s-arch that explains the rationale for the square}. That is
The constant of proportionality (k) does not have to be standardised for interoperability, but a value of 2 is RECOMMENDED.
[I-D.ietf-tsvwg-aqm-dualq-coupled] specifies the essential aspects of an L4S AQM, as well as recommending other aspects. It gives example implementations in appendices.
The term 'likelihood' is used above to allow for marking and dropping to be either probabilistic or deterministic. The example AQMs in [I-D.ietf-tsvwg-aqm-dualq-coupled] drop and mark probabilistically, so the drop probability is arranged to be the square of the marking probability. Nonetheless, an alternative AQM that dropped and marked deterministically would be valid, as long as the dropping frequency was proportional to the square of the marking frequency.
Note that, contrary to RFC 3168, a Dual AQM implementing the L4S and Classic treatments does not mark an ECT(1) packet under the same conditions that it would have dropped a Not-ECT packet, as allowed by [RFC8311], which updates RFC 3168. However, it does mark an ECT(0) packet under the same conditions that it would have dropped a Not-ECT packet.
To implement the L4S treatment, a network node does not need to identify transport-layer flows. Nonetheless, if an implementer is willing to identify transport-layer flows at a network node, and if the most recent ECT packet in the same flow was ECT(0), the node MAY classify CE packets for classic ECN [RFC3168] treatment. In all other cases, a network node MUST classify all CE packets for L4S treatment. Examples of such other cases are: i) if no ECT packets have yet been identified in a flow; ii) if it is not desirable for a network node to identify transport-layer flows; or iii) if the most recent ECT packet in a flow was ECT(1).
If an implementer uses flow-awareness to classify CE packets, to determine whether the flow is using ECT(0) or ECT(1) it only uses the most recent ECT packet of a flow (this advice will need to be verified as part of L4S experiments). This is because a sender might switch from sending ECT(1) (L4S) packets to sending ECT(0) (Classic) packets, or back again, in the middle of a transport-layer flow (e.g. it might manually switch its congestion control module mid-connection, or it might be deliberately attempting to confuse the network).
In a typical case for the public Internet a network element that implements L4S might want to classify some low-rate but unresponsive traffic (e.g. DNS, voice, game sync packets) into the low latency queue to mix with L4S traffic. Such non-ECN-based packet types MUST be safe to mix with L4S traffic without harming the low latency service, where 'safe' is explained in Section 5.4.1.1.1 below.
In this case it would not be appropriate to call the queue an L4S queue, because it is shared by L4S and non-L4S traffic. Instead it will be called the low latency or L queue. The L queue then offers two different treatments:
To identify packets for just the scheduling treatment, it would be inappropriate to use the L4S ECT(1) identifier, because such traffic is unresponsive to ECN marking. Therefore, a network element that implements L4S MAY classify additional packets into the L queue if they carry certain non-ECN identifiers. For instance:
Of course, a packet that carried both the ECT(1) codepoint and a relevant non-ECN identifier would also be classified into the L queue.
For clarity, non-ECN identifiers, such as the examples itemized above, might be used by some network operators who believe they identify non-L4S traffic that would be safe to mix with L4S traffic. They are not alternative ways for a host to indicate that it is sending L4S packets. Only the ECT(1) ECN codepoint indicates to a network element that a host is sending L4S packets (and CE indicates that it could be). Specifically ECT(1) indicates that the host claims its behaviour satisfies the per-requisite transport requirements in Section 4.
The above section requires unresponsive traffic to be 'safe' to mix with L4S traffic. Ideally this means that the sender never sends any sequence of packets at a data rate that exceeds the available capacity of the bottleneck link. However, typically an unresponsive transport does not even know the bottleneck capacity of the path, let alone its available capacity. Nonetheless, an application can be considered safe enough if it paces packets out (not necessarily completely regularly) such that its maximum instantaneous data rate from packet to packet stays well below a typical broadband access rate.
This is a vague but useful definition, because it encompasses many low latency applications of interest, such as DNS, voice, game sync packets, RPC, ACKs, keep-alives, etc.
To extend the above example, an operator might want to exclude some traffic from the L4S treatment for policy reason, e.g. security (traffic from malicious sources) or commercial (initially the operator may wish to confine the benefits of L4S to business customers).
In this exclusion case, the operator MUST classify on the relevant locally-used identifiers (e.g. source addresses) before classifying the non-matching traffic on the end-to-end L4S ECN identifier.
The operator MUST NOT re-mark the end-to-end L4S identifier, because its decision to exclude certain traffic from L4S treatment is local-only. The end-to-end L4S identifier then survives for other operators to use, or indeed, they can apply their own policy, independently based on their own choice of locally-used identifiers. This approach also allows any operator to remove its locally-applied exclusions in future, e.g. if it wishes to widen the benefit of the L4S treatment to all its customers.
L4S concerns low latency, which it can provide for all traffic without differentiation and without affecting bandwidth allocation. Diffserv provides for differentiation of both bandwidth and low latency, but its control of latency depends on its control of bandwidth. The two can be combined if a network operator wants to control bandwidth allocation but it also wants to provide low latency - for any amount of traffic within one of these allocations of bandwidth (rather than only providing low latency by limiting bandwidth) [I-D.briscoe-tsvwg-l4s-diffserv].
The examples above were framed in the context of providing the default Best Efforts Per-Hop Behaviour (PHB) using two queues - a Low Latency (L) queue and a Classic (C) Queue. This single DualQ structure is expected to be by far the most common and useful arrangement. But, more generally, an operator might choose to control bandwidth allocation through a hierarchy of Diffserv PHBs at a node, and to offer one (or more) of these PHBs with a low latency and a classic variant.
In the first case, if we assume that there are no other PHBs except the DualQ, if a packet carries ECT(1) or CE, a network element would classify it for the L4S treatment irrespective of its DSCP. And, if a packet carried (say) the EF DSCP, the network element could classify it into the L queue irrespective of its ECN codepoint. However, where the DualQ is in a hierarchy of other PHBs, the classifier would classify some traffic into other PHBs based on DSCP before classifying between the latency and classic queues (based on ECT(1), CE and perhaps also the EF DSCP or other identifiers as in the above example). [I-D.briscoe-tsvwg-l4s-diffserv] gives a number of examples of such arrangements to address various requirements.
[I-D.briscoe-tsvwg-l4s-diffserv] describes how an operator might use L4S to offer low latency for all L4S traffic as well as using Diffserv for bandwidth differentiation. It identifies two main types of approach, which can be combined: the operator might split certain Diffserv PHBs between L4S and a corresponding Classic service. Or it might split the L4S and/or the Classic service into multiple Diffserv PHBs. In any of these cases, a packet would have to be classified on its Diffserv and ECN codepoints.
In summary, there are numerous ways in which the L4S ECN identifier (ECT(1) and CE) could be combined with other identifiers to achieve particular objectives. The following categorization articulates those that are valid, but it is not necessarily exhaustive. Those tagged 'Global-use' could be set by the sending host or a network. Those tagged 'Local-use' would only be set by a network:
[I-D.ietf-tsvwg-aqm-dualq-coupled] sets operational and management requirements for experiments with DualQ Coupled AQMs. General operational and management requirements for experiments with L4S congestion controls are given in Section 4 and Section 5 above, e.g. co-existence and scaling requirements, incremental deployment arrangements. The specification of each scalable congestion control will need to include protocol-specific requirements for configuration and monitoring performance during experiments. Appendix A of [RFC5706] provides a helpful checklist.
This specification contains no IANA considerations.
Approaches to assure the integrity of signals using the new identifer are introduced in Appendix C.1.
The requirement to detect loss in time units prevents the ACK-splitting attacks described in [Savage-TCP].
Thanks to Richard Scheffenegger, John Leslie, David Täht, Jonathan Morton, Gorry Fairhurst, Michael Welzl, Mikael Abrahamsson and Andrew McGregor for the discussions that led to this specification. Ing-jyh (Inton) Tsang was a contributor to the early drafts of this document. And thanks to Mikael Abrahamsson, Lloyd Wood, Nicolas Kuhn, Greg White, David Black and Gorry Fairhurst for providing help and reviewing this draft and to Ingemar Johansson for reviewing and providing substantial text. Appendix A listing the Prague L4S Requirements is based on text authored by Marcelo Bagnulo Braun that was originally an appendix to [I-D.ietf-tsvwg-l4s-arch]. That text was in turn based on the collective output of the attendees listed in the minutes of a 'bar BoF' on DCTCP Evolution during IETF-94 [TCPPrague].
The authors' contributions were part-funded by the European Community under its Seventh Framework Programme through the Reducing Internet Transport Latency (RITE) project (ICT-317700). Bob Briscoe was also part-funded by the Research Council of Norway through the TimeIn project. The views expressed here are solely those of the authors.
| [RFC2119] | Bradner, S., "Key words for use in RFCs to Indicate Requirement Levels", BCP 14, RFC 2119, DOI 10.17487/RFC2119, March 1997. |
| [RFC3168] | Ramakrishnan, K., Floyd, S. and D. Black, "The Addition of Explicit Congestion Notification (ECN) to IP", RFC 3168, DOI 10.17487/RFC3168, September 2001. |
| [RFC4774] | Floyd, S., "Specifying Alternate Semantics for the Explicit Congestion Notification (ECN) Field", BCP 124, RFC 4774, DOI 10.17487/RFC4774, November 2006. |
| [RFC6679] | Westerlund, M., Johansson, I., Perkins, C., O'Hanlon, P. and K. Carlberg, "Explicit Congestion Notification (ECN) for RTP over UDP", RFC 6679, DOI 10.17487/RFC6679, August 2012. |
This appendix is informative, not normative. It gives a list of modifications to current scalable congestion controls so that they can be deployed over the public Internet and coexist safely with existing traffic. The list complements the normative requirements in Section 4 that a sender has to comply with before it can set the L4S identifier in packets it sends into the Internet. As well as necessary safety improvements (requirements) this appendix also includes preferable performance improvements (optimizations).
These recommendations have become know as the Prague L4S Requirements, because they were originally identified at an ad hoc meeting during IETF-94 in Prague [TCPPrague]. The wording has been generalized to apply to all scalable congestion controls, not just TCP congestion control specifically. They were originally called the 'TCP Prague Requirements', but they are not solely applicable to TCP, so the name has been generalized, and TCP Prague is now used for a specific implementation of the requirements.
DCTCP [RFC8257] is currently the most widely used scalable transport protocol. In its current form, DCTCP is specified to be deployable only in controlled environments. Deploying it in the public Internet would lead to a number of issues, both from the safety and the performance perspective. The modifications and additional mechanisms listed in this section will be necessary for its deployment over the global Internet. Where an example is needed, DCTCP is used as a base, but it is likely that most of these requirements equally apply to other scalable congestion controls.
Description: A scalable congestion control needs to distinguish the packets it sends from those sent by classic congestion controls.
Motivation: It needs to be possible for a network node to classify L4S packets without flow state into a queue that applies an L4S ECN marking behaviour and isolates L4S packets from the queuing delay of classic packets.
Description: The transport protocol for a scalable congestion control needs to provide timely, accurate feedback about the extent of ECN marking experienced by all packets.
Motivation: Classic congestion controls only need feedback about the existence of a congestion episode within a round trip, not precisely how many packets were marked with ECN or dropped. Therefore, in 2001, when ECN feedback was added to TCP [RFC3168], it could not inform the sender of more than one ECN mark per RTT. Since then, requirements for more accurate ECN feedback in TCP have been defined in [RFC7560] and [I-D.ietf-tcpm-accurate-ecn] specifies an experimental change to the TCP wire protocol to satisfy these requirements. Most other transport protocols already satisfy this requirement.
Description: A scalable congestion control needs to react to packet loss in a way that will coexist safely with a TCP Reno congestion control [RFC5681].
Motivation: Part of the safety conditions for deploying a scalable congestion control on the public Internet is to make sure that it behaves properly when it builds a queue at a network bottleneck that has not been upgraded to support L4S. Packet loss can have many causes, but it usually has to be conservatively assumed that it is a sign of congestion. Therefore, on detecting packet loss, a scalable congestion control will need to fall back to classic congestion control behaviour. If it does not comply with this requirement it could starve classic traffic.
A scalable congestion control can be used for different types of transport, e.g. for real-time media or for reliable bulk transport like TCP. Therefore, the particular classic congestion control behaviour to fall back on will need to be part of the congestion control specification of the relevant transport. In the particular case of DCTCP, the current DCTCP specification states that "It is RECOMMENDED that an implementation deal with loss episodes in the same way as conventional TCP." For safe deployment of a scalable congestion control in the public Internet, the above requirement would need to be defined as a "MUST".
Packet loss might (rarely) occur in the case that the bottleneck is L4S capable. In this case, the sender may receive a high number of packets marked with the CE bit set and also experience a loss. Current DCTCP implementations react differently to this situation. At least one implementation reacts only to the drop signal (e.g. by halving the CWND) and at least another DCTCP implementation reacts to both signals (e.g. by halving the CWND due to the drop and also further reducing the CWND based on the proportion of marked packet). We believe that further experimentation is needed to understand what is the best behaviour for the public Internet, which may or not be one of these existing approaches.
Description: A scalable congestion control needs to react to ECN marking from a non-L4S but ECN-capable bottleneck in a way that will coexist with a TCP Reno congestion control [RFC5681].
Motivation: Similarly to the requirement in Appendix A.1.3, this requirement is a safety condition to ensure a scalable congestion control behaves properly when it builds a queue at a network bottleneck that has not been upgraded to support L4S. On detecting classic ECN marking (see below), a scalable congestion control will need to fall back to classic congestion control behaviour. If it does not comply with this requirement it could starve classic traffic.
It would take time for endpoints to distinguish classic and L4S ECN marking. An increase in queuing delay or in delay variation would be a tell-tale sign, but it is not yet clear where a line would be drawn between the two behaviours. It might be possible to cache what was learned about the path to help subsequent attempts to detect the type of marking.
Description: A scalable congestion control needs to reduce or eliminate RTT bias over as wide a range of RTTs as possible, or at least over the typical range of RTTs that will interact in the intended deployment scenario.
Motivation: Classic TCP's throughput is known to be inversely proportional to RTT, so one would expect flows over very low RTT paths to nearly starve flows over larger RTTs. However, Classic TCP has never allowed a very low RTT path to exist because it induces a large queue. For instance, consider two paths with base RTT 1ms and 100ms. If Classic TCP induces a 100ms queue, it turns these RTTs into 101ms and 200ms leading to a throughput ratio of about 2:1. Whereas if a Scalable TCP induces only a 1ms queue, the ratio is 2:101, leading to a throughput ratio of about 50:1.
Therefore, with very small queues, long RTT flows will essentially starve, unless scalable congestion controls comply with this requirement.
Description: A scalable congestion control needs to remain responsive to congestion when RTTs are significantly smaller than in the current public Internet.
Motivation: As currently specified, the minimum required congestion window of TCP (and its derivatives) is set to 2 maximum segment sizes (MSS) (see equation (4) in [RFC5681]). Once the congestion window reaches this minimum, all current TCP algorithms become unresponsive to congestion signals. No matter how much drop or ECN marking, the congestion window no longer reduces. Instead, TCP forces the queue to grow, overriding any AQM and increasing queuing delay.
L4S mechanisms significantly reduce queueing delay so, over the same path, the RTT becomes lower. Then this problem becomes surprisingly common [TCP-sub-mss-w]. This is because, for the same link capacity, smaller RTT implies a smaller window. For instance, consider a residential setting with an upstream broadband Internet access of 8 Mb/s, assuming a max segment size of 1500 B. Two upstream flows will each have the minimum window of 2 MSS if the RTT is 6ms or less, which is quite common when accessing a nearby data centre. So, any more than two such parallel TCP flows will become unresponsive and increase queuing delay.
Unless scalable congestion controls are required to comply with this requirement from the start, they will frequently become unresponsive, negating the low latency benefit of L4S, for themselves and for others. One possible sub-MSS window mechanism is described in [TCP-sub-mss-w], and other approaches are likely to be feasible.
Description: A scalable congestion control needs to detect loss by counting in time-based units, which is scalable, rather than counting in units of packets, which is not.
Motivation: If it is known that all L4S senders using a link obey this rule, then link technologies that support L4S can remove the head-of-line blocking delay they have to introduce while trying to keep packets in tight order to avoid triggering loss detection based on counting packets.
End-systems cannot know whether a missing packet is due to loss or reordering, except in hindsight - if it appears later. If senders deem that loss has occurred by counting reordered packets (e.g. the 3 Duplicate ACK rule of Classic TCP), the time over which the network has to keep packets in order scales down as packet rates scale up over the years. In contrast, if senders allow a reordering window in time-based units before they deem there has been a loss, the time over which the network has to keep packets in order stays constant.
The potential benefit for links comes in two parts: i) switching the unit from packet count to time-based; ii) potentially relaxing the amount of time available for re-ordering. The initial switch to time-based units offers no immediate benefit, but as the years progress it stops the reordering requirement becoming tighter. The secondary relaxation might be possible where some transport protocols find they can tolerate more re-ordering (e.g. more than the 3 DupACK rule, perhaps because it was reasonable when packet rates were low, but it is now far too tight).
Tolerance of reordering over a small duration could allow parallel (e.g. bonded-channel) link technologies to relax their need to deliver packets strictly in order. Such links typically give arriving packets a link-level sequence number and introduce delay while buffering packets at the receiving end until they can be delivered in the same order. For radio links, this delay usually includes the time allowed for link-layer retransmissions.
Note, however, that a link technology will only be able to relax its ordering requirement if it is certain that it will not degrade the transport /most/ sensitive to reordering that might use the link. Also note that in some controlled environments no reordering is tolerated by some transports (e.g. RoCEv2 ussed for RDMA), therefore a switch to time-based units could not be exploited to relax reordering.
For receivers that need their packets in order, it would seem that relaxing network ordering would simply shift this reordering delay from the network to the receiver. However, that is not true in the general case because links generally do not recognize transport layer flows and often cannot even see application layer streams within the flows (as in SCTP, HTTP/2 or QUIC). So a link will often be holding back packets from one flow or stream while waiting for those from another. Relaxing strict ordering in the network will remove this head-of-line blocking delay. {ToDo: this is being quantified experimentally - will need to add the figures here.}
Classic TCP implementations are switching over to the time-based approach of RACK (Recent ACKnowledgements [I-D.ietf-tcpm-rack]). However, it will be many years (decades?) before networks no longer have to allow for the presence of traditional TCP senders still using the 3 DupACK rule. This specification (Section 4.3) says that senders are not entitled to identify packets as L4S in the IP/ECN field unless they use the time-based approach. Then networks that identify L4S traffic separately (e.g. using [I-D.ietf-tsvwg-aqm-dualq-coupled]) can know for certain that all L4S traffic is using the scalable time-based approach.
This will allow networks to remove the head-of-line blocking delay of resequencing straight away, but only for L4S traffic. Classic traffic will have to wait for many years until incremental deployment of RACK has become near-universal. Nonetheless, experience with RACK will determine how much reordering tolerance networks will be reasonable for L4S traffic.
Performance Optimization as well as Safety Improvement: The delay benefit would be lost if any L4S sender did not follow the time-based approach. Therefore, the time-based approach is made a normative requirement (a necessary safety improvement). Nonetheless, the time-based approach also enables a throughput benefit that a flow can enjoy independently of others (a performance optimization), explained next.
Given the requirement for a scalable congestion control to fall-back to Reno or Cubic on a loss (see Appendix A.1.3), it is important that a scalable congestion control does not deem that a loss has occurred too soon. If, later within the same round trip, an out-of-order acknowledgement fills the gap, the sender would have halved its rate spuriously (as well as retransmitting spuriously). With a RACK-like approach, allowing longer before a loss is deemed to have occurred maintains higher throughput in the presence of reordering {ToDo: Quantify this statement}.
On the other hand, it is also important not to wait too long before deeming that a gap is due to a loss (termed a long reordering window), otherwise loss recovery would be slow.
The speed of loss recovery is much more significant for short flows than long, therefore a good compromise would adapt the reordering window; from a small fraction of the RTT at the start of a flow, to a larger fraction of the RTT for flows that continue for many round trips. This is the approach adopted by TCP RACK (Recent ACKnowledgements) [I-D.ietf-tcpm-rack] and recommended for all L4S senders, whether using TCP or another transport protocol.
The requirement to detect loss in time units also prevents the ACK-splitting attacks described in [Savage-TCP].
Description: This item only concerns TCP and its derivatives (e.g. SCTP), because the original specification of ECN for TCP precluded the use of ECN on control packets and retransmissions. To improve performance, scalable transport protocols ought to enable ECN at the IP layer in TCP control packets (SYN, SYN-ACK, pure ACKs, etc.) and in retransmitted packets. The same is true for derivatives of TCP, e.g. SCTP.
Motivation: RFC 3168 prohibits the use of ECN on these types of TCP packet, based on a number of arguments. This means these packets are not protected from congestion loss by ECN, which considerably harms performance, particularly for short flows. [I-D.ietf-tcpm-generalized-ecn] counters each argument in RFC 3168 in turn, showing it was over-cautious. Instead it proposes experimental use of ECN on all types of TCP packet as long as AccECN feedback [I-D.ietf-tcpm-accurate-ecn] is available (which is itself a prerequisite for using a scalable congestion control).
Description: It would improve performance if scalable congestion controls did not limit their congestion window increase to the traditional additive increase of 1 MSS per round trip [RFC5681] during congestion avoidance. The same is true for derivatives of TCP congestion control, including similar approaches used for real-time media.
Motivation: As currently defined, DCTCP uses the traditional TCP Reno additive increase in congestion avoidance phase. When the available capacity suddenly increases (e.g. when another flow finishes, or if radio capacity increases) it can take very many round trips to take advantage of the new capacity. In the steady state, DCTCP induces about 2 ECN marks per round trip, so it should be possible to quickly detect when these signals have disappeared and seek available capacity more rapidly. It will of course be necessary to minimize the impact on other flows (classic and scalable).
TCP Cubic was designed to solve this problem, but as flow rates have continued to increase, the delay accelerating into available capacity has become prohibitive. For instance, with RTT=20 ms, to increase flow rate from 100Mb/s to 200Mb/s Cubic takes between 50 and 100 round trips. Every 8x increase in flow rate leads to 2x more acceleration delay.
Description: Particularly when a flow starts, scalable congestion controls need to converge (reach their steady-state share of the capacity) at least as fast as classic TCP and preferably faster. This does not just affect TCP Prague, but also the flow start behaviour of any L4S congestion control derived from a Classic transport that uses TCP slow start, including those for real-time media.
Motivation: As an example, a new DCTCP flow takes longer than classic TCP to obtain its share of the capacity of the bottleneck when there are already ongoing flows using the bottleneck capacity. In a data centre environment DCTCP takes about a factor of 1.5 to 2 longer to converge due to the much higher typical level of ECN marking that DCTCP background traffic induces, which causes new flows to exit slow start early [Alizadeh-stability]. In testing for use over the public Internet the convergence time of DCTCP relative to regular TCP is even less favourable [Paced-Chirping]). It is exacerbated by the typically greater mismatch between the link rate of the sending host and typical Internet access bottlenecks, in combination with the shallow ECN marking threshold needed for L4S. This problem is detrimental in general, but would particularly harm the performance of short flows relative to classic TCP.
This appendix is informative, not normative. It records the pros and cons of various alternative ways to identify L4S packets to record the rationale for the choice of ECT(1) (Appendix B.1) as the L4S identifier. At the end, Appendix B.6 summarises the distinguishing features of the leading alternatives. It is intended to supplement, not replace the detailed text.
The leading solutions all use the ECN field, sometimes in combination with the Diffserv field. Both the ECN and Diffserv fields have the additional advantage that they are no different in either IPv4 or IPv6. A couple of alternatives that use other fields are mentioned at the end, but it is quickly explained why they are not serious contenders.
Definition:
Cons:
It is extremely unlikely that a set of 5 eventualities that are each unusual in themselves would all happen simultaneously. But, even if they did, it would only cause spurious retransmission of a packet.
Pros:
Definition:
Use of a DSCP is the only approach for alternative ECN semantics given as an example in [RFC4774]. However, it was perhaps considered more for controlled environments than new end-to-end services;
Cons:
Pros:
Definition:
This approach would only be feasible if
Cons:
Pros:
It has been suggested that a new ID in the IPv4 Protocol field or the IPv6 Next Header field could identify L4S packets. However this approach is ruled out by numerous problems:
Locally, a network operator could arrange for L4S service to be applied based on source or destination addressing, e.g. packets from its own data centre and/or CDN hosts, packets to its business customers, etc. It could use addressing at any layer, e.g. IP addresses, MAC addresses, VLAN IDs, etc. Although addressing might be a useful tactical approach for a single ISP, it would not be a feasible approach to identify an end-to-end service like L4S. Even for a single ISP, it would require packet classifiers in buffers to be dependent on changing topology and address allocation decisions elsewhere in the network. Therefore this approach is not a feasible solution.
Table 1 provides a very high level summary of the pros and cons detailed against the schemes described respectively in Appendix B.2, Appendix B.3 and Appendix B.1, for six issues that set them apart.
| Issue | DSCP + ECN | ECN | ECT(1) + CE |
|---|---|---|---|
| initial eventual | initial | initial eventual | |
| end-to-end | N . . . ? . | . . Y | . . Y . . Y |
| tunnels | . O . . O . | . . ? | . . ? . . Y |
| lower layers | N . . . ? . | . O . | . O . . . ? |
| codepoints | N . . . . ? | . . Y | N . . . . ? |
| reordering | . . Y . . Y | . . Y | . O . . . ? |
| ctrl pkts | . . Y . . Y | . O . | . O . . . ? |
| Note 1 |
Note 1: Only feasible if classic ECN is obsolete.
The schemes are scored based on both their capabilities now ('initial') and in the long term ('eventual'). The 'ECN' scheme shares the 'eventual' scores of the 'ECT(1) + CE' scheme. The scores are one of 'N, O, Y', meaning 'Poor', 'Ordinary', 'Good' respectively. The same scores are aligned vertically to aid the eye. A score of "?" in one of the positions means that this approach might optimisitically become this good, given sufficient effort. The table summarises the text and is not meant to be understandable without having read the text.
The ECT(1) codepoint of the ECN field has already been assigned once for the ECN nonce [RFC3540], which has now been categorized as historic [RFC8311]. ECN is probably the only remaining field in the Internet Protocol that is common to IPv4 and IPv6 and still has potential to work end-to-end, with tunnels and with lower layers. Therefore, ECT(1) should not be reassigned to a different experimental use (L4S) without carefully assessing competing potential uses. These fall into the following categories:
Receiving hosts can fool a sender into downloading faster by suppressing feedback of ECN marks (or of losses if retransmissions are not necessary or available otherwise).
The historic ECN nonce protocol [RFC3540] proposed that a TCP sender could set either of ECT(0) or ECT(1) in each packet of a flow and remember the sequence it had set. If any packet was lost or congestion marked, the receiver would miss that bit of the sequence. An ECN Nonce receiver had to feed back the least significant bit of the sum, so it could not suppress feedback of a loss or mark without a 50-50 chance of guessing the sum incorrectly.
It is highly unlikely that ECT(1) will be needed for integrity protection in future. The ECN Nonce RFC [RFC3540] as been reclassified as historic, partly because other ways have been developed to protect TCP feedback integrity [RFC8311] that do not consume a codepoint in the IP header. For instance:
Various researchers have proposed to use ECT(1) as a less severe congestion notification than CE, particularly to enable flows to fill available capacity more quickly after an idle period, when another flow departs or when a flow starts, e.g. VCP [VCP], Queue View (QV) [QV].
Before assigning ECT(1) as an identifer for L4S, we must carefully consider whether it might be better to hold ECT(1) in reserve for future standardisation of rapid flow acceleration, which is an important and enduring problem [RFC6077].
Pre-Congestion Notification (PCN) is another scheme that assigns alternative semantics to the ECN field. It uses ECT(1) to signify a less severe level of pre-congestion notification than CE [RFC6660]. However, the ECN field only takes on the PCN semantics if packets carry a Diffserv codepoint defined to indicate PCN marking within a controlled environment. PCN is required to be applied solely to the outer header of a tunnel across the controlled region in order not to interfere with any end-to-end use of the ECN field. Therefore a PCN region on the path would not interfere with any of the L4S service identifiers proposed in Appendix B.