| Internet-Draft | L4S ECN Protocol for V Low Queuing Delay | July 2022 |
| De Schepper & Briscoe | Expires 28 January 2023 | [Page] |
This specification defines the protocol to be used for a new network service called low latency, low loss and scalable throughput (L4S). L4S uses an Explicit Congestion Notification (ECN) scheme at the IP layer that is similar to the original (or 'Classic') ECN approach, except as specified within. L4S uses 'scalable' congestion control, which induces much more frequent control signals from the network and it responds to them with much more fine-grained adjustments, so that very low (typically sub-millisecond on average) and consistently low queuing delay becomes possible for L4S traffic without compromising link utilization. Thus even capacity-seeking (TCP-like) traffic can have high bandwidth and very low delay at the same time, even during periods of high traffic load.¶
The L4S identifier defined in this document distinguishes L4S from 'Classic' (e.g. TCP-Reno-friendly) traffic. It gives an incremental migration path so that suitably modified network bottlenecks can distinguish and isolate existing traffic that still follows the Classic behaviour, to prevent it degrading the low queuing delay and low loss of L4S traffic. This specification defines the rules that L4S transports and network elements need to follow with the intention that L4S flows neither harm each other's performance nor that of Classic traffic. Examples of new active queue management (AQM) marking algorithms and examples of new transports (whether TCP-like or real-time) are specified separately.¶
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This specification defines the protocol to be used for a new network service called low latency, low loss and scalable throughput (L4S). L4S uses an Explicit Congestion Notification (ECN) scheme at the IP layer with the same set of codepoint transitions as the original (or 'Classic') Explicit Congestion Notification (ECN [RFC3168]). RFC 3168 required an ECN mark 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 comparable non-L4S flow under the same conditions. Nonetheless, the much more frequent ECN control signals and the finer responses to these signals result in very low queuing delay without compromising link utilization, and this low delay can be maintained during high load. For instance, queuing delay under heavy and highly varying load with the example DCTCP/DualQ solution cited below on a DSL or Ethernet link is sub-millisecond on average and roughly 1 to 2 milliseconds at the 99th percentile without losing link utilization [DualPI2Linux], [DCttH19]. Note that the inherent queuing delay while waiting to acquire a discontinuous medium such as WiFi has to be minimized in its own right, so it would be additional to the above (see section 6.3 of the L4S architecture [I-D.ietf-tsvwg-l4s-arch]).¶
L4S relies on 'scalable' congestion controls for these delay properties and for preserving low delay as flow rate scales, hence the name. The congestion control used in Data Center TCP (DCTCP) is an example of a scalable congestion control, but DCTCP is applicable solely to controlled environments like data centres [RFC8257], because it is too aggressive to co-exist with existing TCP-Reno-friendly traffic. 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 derived from DCTCP to co-exist with existing traffic, each getting roughly the same flow rate when they compete under similar conditions. Note that a scalable congestion control is still not safe to deploy on the Internet unless it satisfies the requirements listed in Section 4.¶
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 factor that distinguishes L4S from Classic traffic is its behaviour in response to congestion. The transport wire protocol, e.g. TCP, QUIC, SCTP, DCCP, RTP/RTCP, is orthogonal (and therefore not suitable for distinguishing L4S from Classic packets).¶
The L4S identifier defined in this document is the key piece that distinguishes L4S from 'Classic' (e.g. Reno-friendly) traffic. It gives an incremental migration path so that suitably modified network bottlenecks can distinguish and isolate existing Classic traffic from L4S traffic to prevent the former from degrading the very low delay and loss of the new scalable transports, without harming Classic performance at these bottlenecks. Initial implementation of the separate parts of the system has been motivated by the performance benefits.¶
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. If growth in high-throughput latency-sensitive applications continues, periods with solely latency-sensitive traffic will become increasingly common on links where traffic aggregation is low. For instance, on the access links dedicated to individual sites (homes, small enterprises or mobile devices). These links also tend to become the path bottleneck under load. During these periods, if all the traffic were marked for the same treatment, at these bottlenecks Diffserv would make no difference. Instead, it becomes imperative to remove the underlying 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, AQM methods 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, this form of AQM was not widely deployed.¶
More recent state-of-the-art AQM methods, 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 sending window of a Classic congestion control will either cause queuing delay to vary or cause the link to be underutilized. 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, thereby roughly doubling the total round-trip time.¶
If a sender's own behaviour is introducing queuing delay variation, no AQM in the network can 'un-vary' the delay without significantly compromising link utilization. Even flow-queuing (e.g. [RFC8290]), which isolates one flow from another, cannot isolate a flow from the delay variations it inflicts on itself. Therefore those applications that need to seek out high bandwidth but also need low latency will have to migrate to scalable congestion control.¶
Altering host behaviour is not enough on its own though. Even if hosts adopt low latency behaviour (scalable congestion controls), they need to be isolated from the behaviour of existing Classic congestion controls that induce large queue variations. L4S enables that migration by providing latency isolation in the network and distinguishing the two types of packets that need to be isolated: L4S and Classic. L4S isolation can be achieved with a queue per flow (e.g. [RFC8290]) but a DualQ [I-D.ietf-tsvwg-aqm-dualq-coupled] is sufficient, and actually gives better tail latency. Both approaches are addressed in this document.¶
The DualQ solution was developed to make very low latency available without requiring per-flow queues at every bottleneck. This was because per-flow-queuing (FQ) has well-known downsides - not least the need to inspect transport layer headers in the network, which makes it incompatible with privacy approaches such as IPSec VPN tunnels, and incompatible with link layer queue management, where transport layer headers can be hidden, e.g. 5G.¶
Latency is not the only concern addressed by L4S: It was known when TCP congestion avoidance was first developed that it would not scale to high bandwidth-delay products (footnote 6 of Jacobson and Karels [TCP-CA]). Given regular broadband bit-rates over WAN distances are already [RFC3649] beyond the scaling range of Reno congestion control, '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] outcompete Classic ECN congestion controls sharing the same queue, which is why they have been confined to private data centres or research testbeds.¶
It turns out that these scalable congestion control algorithms that solve the latency problem can also solve the scalability problem of Classic congestion controls. The finer sawteeth in the congestion window have low amplitude, so they cause very little queuing delay variation and the average time to recover from one congestion signal to the next (the average duration of each sawtooth) remains invariant, which maintains constant tight control as flow-rate scales. A background paper [DCttH19] 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 Section 4 of the L4S architecture [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.¶
Note: The L4S architecture [I-D.ietf-tsvwg-l4s-arch] repeats the following definitions, but if there are accidental differences those below take precedence.¶
The 'L4S' service is intended for traffic from scalable congestion control algorithms, such as TCP Prague [I-D.briscoe-iccrg-prague-congestion-control], which was derived from DCTCP [RFC8257]. The L4S service is for more general traffic than just TCP Prague -- it allows the set of congestion controls with similar scaling properties to Prague to evolve, such as the examples listed above (Relentless, SCReAM). The term 'L4S queue' means a queue providing the L4S service.¶
The terms Classic or L4S can also qualify other nouns, such as 'queue', 'codepoint', 'identifier', 'classification', 'packet', 'flow'. For example: an L4S packet means a packet with an L4S identifier sent from an L4S congestion control.¶
Both Classic and L4S services can cope with a proportion of unresponsive or less-responsive traffic as well, but in the L4S case its rate has to be smooth enough or low enough not to build a queue (e.g. DNS, VoIP, game sync datagrams, etc).¶
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:¶
makes consequent updates to the following additional proposed standard RFCs to reflect the above two bullets:¶
This document is about identifiers that are used for interoperation between hosts and networks. So the audience is broad, covering developers of host transports and network AQMs, as well as covering how operators might wish to combine various identifiers, which would require flexibility from equipment developers.¶
This subsection briefly records the process that led to the chosen L4S identifier.¶
The identifier for packets using the Low Latency, Low Loss, Scalable throughput (L4S) service needs 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 any choice of identifier is unlikely to satisfy all these requirements, particularly given the limited space left in the IP header. Therefore a compromise will always be necessary, which is why all the above requirements are expressed with the word 'SHOULD' not 'MUST'.¶
After extensive assessment of alternative schemes, "ECT(1) and CE codepoints" was chosen as the best compromise. Therefore this scheme is defined in detail in the following sections, while Appendix B records its pros and cons against the above requirements.¶
The L4S treatment is an experimental track alternative packet marking treatment to the Classic ECN treatment in [RFC3168], which has been updated by [RFC8311] to allow experiments such as the one defined in the present specification. [RFC4774] discusses some of the issues and evaluation criteria when defining alternative ECN semantics. 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 always classifies arriving ECT(1) packets for L4S treatment and by default classifies CE packets for L4S treatment unless the heuristics described in Section 5.3 are employed. 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.¶
For a transport protocol to provide scalable congestion control (Section 4.3) it MUST provide 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 ECN 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 that, in steady state, the average duration between induced ECN marks does not increase as flow rate scales up, all other factors being equal. This is termed a scalable congestion control. This invariant duration ensures that, as flow rate scales, the average period with no feedback information about capacity does not become excessive. It also ensures that queue variations remain small, without having to sacrifice utilization.¶
With a congestion control that sawtooths to probe capacity, this duration is called the recovery time, because each time the sawtooth yields, on average it take this time to recover to its previous high point. A scalable congestion control does not have to sawtooth, but it has to coexist with scalable congestion controls that do.¶
For instance, for DCTCP [RFC8257], TCP Prague [I-D.briscoe-iccrg-prague-congestion-control], [PragueLinux] and the L4S variant of SCReAM [RFC8298], the average recovery time is always half a round trip (or half a reference round trip), whatever the flow rate.¶
As with all transport behaviours, a detailed specification (probably an experimental RFC) is expected for each congestion control, following the guidelines for specifying new congestion control algorithms in [RFC5033]. In addition it is expected to document these L4S-specific matters, specifically the timescale over which the proportionality is averaged, and control of burstiness. The recovery time requirement above is worded as a 'SHOULD' rather than a 'MUST' to allow reasonable flexibility for such implementations.¶
The condition 'all other factors being equal', allows the recovery time to be different for different round trip times, as long as it does not increase with flow rate for any particular RTT.¶
Saying that the recovery time remains roughly invariant is equivalent to saying that the number of ECN CE marks per round trip remains invariant as flow rate scales, all other factors being equal. For instance, an average recovery time of half of 1 RTT is equivalent to 2 ECN marks per round trip. For those familiar with steady-state congestion response functions, it is also equivalent to say that the congestion window is inversely proportional to the proportion of bytes in packets marked with the CE codepoint (see section 2 of [PI2]).¶
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:¶
In uncontrolled environments, monitoring MUST be implemented to support detection of problems with an ECN-capable AQM at the path bottleneck that appears not to support L4S and might be in a shared queue. Such monitoring SHOULD be applied to live traffic that is using Scalable congestion control. Alternatively, monitoring need not be applied to live traffic, if monitoring has been arranged to cover the paths that live traffic takes through uncontrolled environments.¶
A function to detect the above problems with an ECN-capable AQM MUST also be implemented and used. The detection function SHOULD be capable of making the congestion control adapt its ECN-marking response in real-time to coexist safely with Classic congestion controls such as standard Reno [RFC5681], as required by [RFC5033]. This could be complemented by more detailed offline detection of potential problems. If only offline detection is used and potential problems with such an AQM are detected on certain paths, the scalable congestion control MUST be replaced by a Classic congestion control, at least for the problem paths.¶
See Section 4.3.1, Appendix A.1.5 and the L4S operational guidance [I-D.ietf-tsvwg-l4sops] for rationale.¶
Note that a scalable congestion control is not expected to change to setting ECT(0) while it transiently adapts to coexist with Classic congestion controls, whereas a replacement congestion control that solely behaves in the Classic way will set ECT(0).¶
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.¶
Later (Section 5.4.1.1) this document discusses the conditions for mixing other "'Safe' Unresponsive Traffic" (e.g. DNS, LDAP, NTP, voice, game sync packets) with L4S traffic. To be clear, although such traffic can share the same queue as L4S traffic, it is not appropriate for the sender to tag it as ECT(1), except in the (unlikely) case that it satisfies the above conditions.¶
RFC 3168 requires the congestion responses to a CE-marked packet and a dropped packet to be the same. RFC 8311 is a standards-track update to RFC 3168 intended to enable experimentation with ECN, including the L4S experiment. RFC 8311 allows an experimental congestion control's response to a CE-marked packet to differ from the response to a dropped packet, provided that the differences are documented in an experimental RFC, such as the present document.¶
BCP 124 [RFC4774] gives guidance to protocol designers, when specifying alternative semantics for the ECN field. RFC 8311 explained that it did not need to update the best current practice in BCP 124 in order to relax the 'equivalence with drop' requirement because, although BCP 124 quotes the same requirement from RFC 3168, the BCP does not impose requirements based on it. BCP 124 describes three options for incremental deployment, with Option 3 (in Section 4.3 of BCP 124) best matching the L4S case. Option 3's requirement for end-nodes is that they respond to CE marks "in a way that is friendly to flows using IETF-conformant congestion control." This echoes other general congestion control requirements in the RFC series, for example [RFC5033], which says "...congestion controllers that have a significantly negative impact on traffic using standard congestion control may be suspect", or [RFC8085] concerning UDP congestion control says "Bulk-transfer applications that choose not to implement TFRC or TCP-like windowing SHOULD implement a congestion control scheme that results in bandwidth (capacity) use that competes fairly with TCP within an order of magnitude."¶
The third normative bullet in Section 4.3 above (which concerns L4S response to congestion from a Classic ECN AQM) aims to ensure that these 'coexistence' requirements are satisfied, but it makes some compromises. This subsection highlights and justifies those compromises and Appendix A.1.5 and the L4S operational guidance [I-D.ietf-tsvwg-l4sops] give detailed analysis, examples and references (the normative text in that bullet takes precedence if any informative elaboration leads to ambiguity). The approach is based on an assessment of the risk of harm, which is a combination of the prevalence of the conditions necessary for harm to occur, and the potential severity of the harm if they do.¶
There are three cases:¶
Classic ECN [RFC3168]: The compromises centre around cases where the bottleneck supports Classic ECN but not L4S. But it depends on which sub-case:¶
Per-Flow-queues with Classic ECN: Most AQMs with per-flow-queuing (FQ) deployed from 2012 onwards had Classic ECN enabled by default, specifically FQ-CoDel [RFC8290] and COBALT [COBALT]. But the compromises only apply to the second of two further sub-cases:¶
To summarize, the coexistence problem is confined to cases of imperfect flow isolation in an FQ, or in potential cases where a Classic ECN AQM has been deployed in a shared queue (see the L4S operational guidance [I-D.ietf-tsvwg-l4sops] for further details including recent surveys attempting to quantify prevalence). Further, if one of these cases does occur, the coexistence problem does not arise unless sources of Classic and L4S flows are simultaneously sharing the same bottleneck queue (e.g. different applications in the same household) and flows of each type have to be large enough to coincide for long enough for any throughput imbalance to have developed.¶
These throughput ratios can clearly fall well outside current RFC guidance on coexistence. However, the tendency towards leaving a greater share for Classic flows at lower link rate and the very limited prevalence of the conditions necessary for harm to occur led to the possibility of allowing the RFC requirements to be compromised, albeit briefly:¶
[I-D.ietf-tsvwg-l4sops] also gives example ways in which L4S congestion controls can be rolled out initially in lower risk scenarios.¶
Section 5.2 below specifies that an L4S AQM is expected to signal L4S ECN immediately, to avoid introducing delay due to filtering or smoothing. This contrasts with a Classic AQM, which filters out variations in the queue before signalling ECN marking or drop. In the L4S architecture [I-D.ietf-tsvwg-l4s-arch], responsibility for smoothing out these variations shifts to the sender's congestion control.¶
This shift of responsibility has the advantage that each sender can smooth variations over a timescale proportionate to its own RTT. Whereas, in the Classic approach, the network doesn't know the RTTs of any of the flows, so it has to smooth out variations for a worst-case RTT to ensure stability. For all the typical flows with shorter RTT than the worst-case, this makes congestion control unnecessarily sluggish.¶
This also gives an L4S sender the choice not to smooth, depending on its context (start-up, congestion avoidance, etc). Therefore, this document places no requirement on an L4S congestion control to smooth out variations in any particular way. Implementers are encouraged to openly publish the approach they take to smoothing, and the results and experience they gain during the L4S experiment.¶
A network node that implements the L4S service:¶
MUST classify arriving CE packets for L4S treatment as well, unless overridden by a another classifier or unless the exception referred to next applies;¶
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 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 increasingly mark the ECN field as CE, 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 MUST begin applying drop to L4S traffic until the overload episode has subsided, as recommended for all AQM methods in [RFC7567] (Section 4.2.1), which follows the similar advice in RFC 3168 (Section 7). During overload, it MUST apply the same drop probability to L4S traffic as it would to Classic traffic.¶
Where an L4S AQM is transport-aware, this requirement could be satisfied by using drop in only the most overloaded individual per-flow AQMs. In a DualQ with flow-aware queue protection (e.g. [I-D.briscoe-docsis-q-protection]), this could be achieved by redirecting packets in those flows contributing most to the overload out of the L4S queue so that they are subjected to drop in the Classic queue.¶
For backward compatibility in uncontrolled environments, a network node that implements the L4S treatment MUST also implement an AQM treatment for the Classic service as defined in Section 1.2. This Classic AQM treatment need not mark ECT(0) packets, but if it does, see Section 5.2 for the strengths of the markings relative to drop. It MUST classify arriving ECT(0) and Not-ECT packets for treatment by this Classic AQM (for the DualQ Coupled AQM, see the extensive discussion on classification in Sections 2.3 and 2.5.1.1 of [I-D.ietf-tsvwg-aqm-dualq-coupled]).¶
In case unforeseen problems arise with the L4S experiment, it MUST be possible to configure an L4S implementation to disable the L4S treatment. Once disabled, all packets of all ECN codepoints will receive Classic treatment and ECT(1) packets MUST be treated as if they were Not-ECT.¶
The relative strengths of L4S CE and drop are irrelevant where AQMs are implemented in separate queues per-application-flow, which are then explicitly scheduled (e.g. with an FQ scheduler as in FQ-CoDel [RFC8290]). Nonetheless, the relationship between them needs to be defined for the coupling between L4S and Classic congestion signals in a DualQ Coupled AQM [I-D.ietf-tsvwg-aqm-dualq-coupled], as below.¶
Unless an AQM node schedules application flows explicitly, the likelihood that the 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). That is¶
The constant of proportionality (k) does not have to be standardised for interoperability, but a value of 2 is RECOMMENDED. The term 'likelihood' is used above to allow for marking and dropping to be either probabilistic or deterministic.¶
This formula ensures that Scalable and Classic flows will converge to roughly equal congestion windows, for the worst case of Reno congestion control. This is because the congestion windows of Scalable and Classic congestion controls are inversely proportional to p_L and sqrt(p_C) respectively. So squaring p_C in the above formula counterbalances the square root that characterizes Reno-friendly flows.¶
Note that, contrary to RFC 3168, an 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, if it marks ECT(0) packets, it does so under the same conditions that it would have dropped a Not-ECT packet [RFC3168].¶
Also, In the L4S architecture [I-D.ietf-tsvwg-l4s-arch], the sender, not the network, is responsible for smoothing out variations in the queue. So, an L4S AQM MUST signal congestion as soon as possible. Then, an L4S sender generally interprets CE marking as an unsmoothed signal.¶
This requirement does not prevent an L4S AQM from mixing in additional congestion signals that are smoothed, such as the signals from a Classic smoothed AQM that are coupled with unsmoothed L4S signals in the coupled DualQ [I-D.ietf-tsvwg-aqm-dualq-coupled]. But only as long as the onset of congestion can be signalled immediately, and can be interpreted by the sender as if it has been signalled immediately, which is important for interoperability¶
To implement L4S packet classification, a network node does not need to identify transport-layer flows. Nonetheless, if an L4S network node classifies packets by their transport-layer flow ID and their ECN field, and if all the ECT packets in a flow have been ECT(0), the node MAY classify any CE packets in the same flow as if they were Classic ECT(0) packets. In all other cases, a network node MUST classify all CE packets as if they were ECT(1) packets. 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 some ECT packets in a flow have been ECT(1) (this advice will need to be verified as part of L4S experiments).¶
The examples in this section concern how additional identifiers might complement the L4S identifier to classify packets between class-based queues. Firstly Section 5.4.1 considers two queues, L4S and Classic, as in the Coupled DualQ AQM [I-D.ietf-tsvwg-aqm-dualq-coupled], either alone (Section 5.4.1.1) or within a larger queuing hierarchy (Section 5.4.1.2). Then Section 5.4.2 considers schemes that might combine per-flow 5-tuples with other identifiers.¶
In a typical case for the public Internet a network element that implements L4S in a shared queue might want to classify some low-rate but unresponsive traffic (e.g. DNS, LDAP, NTP, voice, game sync packets) into the low latency queue to mix with L4S traffic. 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. Examples of relevant non-ECN identifiers are:¶
In summary, a network element that implements L4S in a shared queue MAY classify additional types of packets into the L queue based on identifiers other than the ECN field, but the types SHOULD be 'safe' to mix with L4S traffic, where 'safe' is explained in Section 5.4.1.1.1.¶
A packet that carries one of these non-ECN identifiers to classify it into the L queue would not be subject to the L4S ECN marking treatment, unless it also carried an ECT(1) or CE codepoint. The specification of an L4S AQM MUST define the behaviour for packets with unexpected combinations of codepoints, e.g. a non-ECN-based classifier for the L queue, but ECT(0) in the ECN field (for examples see section 2.5.1.1 of the DualQ spec [I-D.ietf-tsvwg-aqm-dualq-coupled]).¶
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 have originated as ECT(1)). Specifically ECT(1) indicates that the host claims its behaviour satisfies the prerequisite transport requirements in Section 4.¶
In order to include non-L4S packets in the L queue, a network node MUST NOT alter Not-ECT or ECT(0) in the IP-ECN field to an L4S identifier. This ensures that these codepoints survive for any potential use later on the network path. If a non-compliant or malicious network node did swap ECT(0) to ECT(1), the packet could subsequently be ECN-marked by a downstream L4S AQM, but the sender would respond to congestion indications thinking it had sent a Classic packet. This could result in the flow yielding excessively to other L4S flows sharing the downstream bottleneck.¶
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 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 rate from packet to packet stays well below a typical broadband access rate.¶
This is a vague but useful definition, because many low latency applications of interest, such as DNS, voice, game sync packets, RPC, ACKs, keep-alives, could match this description.¶
Low rate streams such as voice and game sync packets, might not use continuously adapting ECN-based congestion control, but they ought to at least use a 'circuit-breaker' style of congestion response [RFC8083]. If the volume of traffic from unresponsive applications is high enough to overload the link, this will at least protect the capacity available to responsive applications. However, queuing delay in the L queue will probably rise to that controlled by the Classic (drop-based) AQM. If a network operator considers that such self-restraint is not enough, it might want to police the L queue (see Section 8.2 of the L4S architecture [I-D.ietf-tsvwg-l4s-arch]).¶
To extend the above example, an operator might want to exclude some traffic from the L4S treatment for a policy reason, e.g. security (traffic from malicious sources) or commercial (e.g. initially the operator may wish to confine the benefits of L4S to business customers).¶
In this exclusion case, the classifier 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.¶
A network node MUST NOT alter the end-to-end L4S ECN identifier from L4S to Classic, because an operator 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. If a non-compliant or malicious network node did swap ECT(1) to ECT(0), the packet could subsequently be ECN-marked by a downstream Classic ECN AQM. L4S senders are required to detect and handle such treatment (Section 4.3 item 3), but that does not make this swap OK, because such detection is not known to be perfect or immediate.¶
A network node that supports L4S but excludes certain packets carrying the L4S identifier from L4S treatment MUST still apply marking or dropping that is compatible with an L4S congestion response. For instance, it could either drop such packets with the same likelihood as Classic packets or it could ECN-mark them with a likelihood appropriate to L4S traffic (e.g. the coupled probability in a DualQ coupled AQM) but aiming for the Classic delay target. It MUST NOT ECN-mark such packets with a Classic marking probability, which could confuse the sender.¶
L4S concerns low latency, which it can provide for all traffic without differentiation and without necessarily 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 DualQ examples so far have been 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 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 using a pair of queues for a low latency and a Classic variant of the PHB.¶
In the first case, if we assume that a network element provides no PHBs except the DualQ, if a packet carries ECT(1) or CE, the 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 low 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 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 either 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 'Recommended-standard-use' could be set by the sending host or a network. Those tagged 'Local-use' would only be set by a network:¶
At a node with per-flow queueing (e.g. FQ-CoDel [RFC8290]), the L4S identifier could complement the Layer-4 flow ID as a further level of flow granularity (i.e. Not-ECT and ECT(0) queued separately from ECT(1) and CE packets). "Risk of reordering Classic CE packets" in Appendix B discusses the resulting ambiguity if packets originally marked ECT(0) are marked CE by an upstream AQM before they arrive at a node that classifies CE as L4S. It argues that the risk of reordering is vanishingly small and the consequence of such a low level of reordering is minimal.¶
Alternatively, it could be assumed that it is not in a flow's own interest to mix Classic and L4S identifiers. Then the AQM could use the ECN field to switch itself between a Classic and an L4S AQM behaviour within one per-flow queue. For instance, for ECN-capable packets, the AQM might consist of a simple marking threshold and an L4S ECN identifier might simply select a shallower threshold than a Classic ECN identifier would.¶
As well as senders needing to limit packet bursts (Section 4.3), links need to limit the degree of burstiness they introduce. In both cases (senders and links) this is a tradeoff, because batch-handling of packets is done for good reason, e.g. processing efficiency or to make efficient use of medium acquisition delay. Some take the attitude that there is no point reducing burst delay at the sender below that introduced by links (or vice versa). However, delay reduction proceeds by cutting down 'the longest pole in the tent', which turns the spotlight on the next longest, and so on.¶
This document does not set any quantified requirements for links to limit burst delay, primarily because link technologies are outside the remit of L4S specifications. Nonetheless, the following two subsections outline opportunities for addressing bursty links in the process of L4S implementation and deployment.¶
It would not make sense to implement an L4S AQM that feeds into a particular link technology without also reviewing opportunities to reduce any form of burst delay introduced by that link technology. This would at least limit the bursts that the link would otherwise introduce into the onward traffic, which would cause jumpy feedback to the sender as well as potential extra queuing delay downstream. This document does not presume to even give guidance on an appropriate target for such burst delay until there is more industry experience of L4S. However, as suggested in Section 4.3 it would not seem necessary to limit bursts lower than roughly 10% of the minimum base RTT expected in the typical deployment scenario (e.g. 250 us burst duration for links within the public Internet).¶
The initial scope of the L4S experiment is to deploy L4S AQMs at bottlenecks and L4S congestion controls at senders. This is expected to highlight interactions with the most bursty upstream links and lead operators to tune down the burstiness of those links in their network that are configurable, or failing that, to have to compromise on the delay target of some L4S AQMs. It might also require specific redesign work relevant to the most problematic link types. Such knock-on effects of initial L4S deployment would all be part of the learning from the L4S experiment.¶
The details of such link changes are beyond the scope of the present document. Nonetheless, where L4S technology is being implemented on an outgoing interface of a device, it would make sense to consider opportunities for reducing bursts arriving at other incoming interface(s). For instance, where an L4S AQM is implemented to feed into the upstream WAN interface of a home gateway, there would be opportunities to alter the WiFi profiles sent out of any WiFi interfaces from the same device, in order to mitigate incoming bursts of aggregated WiFi frames from other WiFi stations.¶
The L4S identifier is expected to work through and within any tunnel without modification, as long as the tunnel propagates the ECN field in any of the ways that have been defined since the first variant in the year 2001 [RFC3168]. L4S will also work with (but does not rely on) any of the more recent updates to ECN propagation in [RFC4301], [RFC6040] or [I-D.ietf-tsvwg-rfc6040update-shim]. However, it is likely that some tunnels still do not implement ECN propagation at all. In these cases, L4S will work through such tunnels, but within them the outer header of L4S traffic will appear as Classic.¶
AQMs are typically implemented where an IP-layer buffer feeds into a lower layer, so they are agnostic to link layer encapsulations. Where a bottleneck link is not IP-aware, the L4S identifier is still expected to work within any lower layer encapsulation without modification, as long it propagates the ECN field as defined for the link technology, for example for MPLS [RFC5129] or TRILL [I-D.ietf-trill-ecn-support]. In some of these cases, e.g. layer-3 Ethernet switches, the AQM accesses the IP layer header within the outer encapsulation, so again the L4S identifier is expected to work without modification. Nonetheless, the programme to define ECN for other lower layers is still in progress [I-D.ietf-tsvwg-ecn-encap-guidelines].¶
If a mix of L4S and Classic packets is sent into the same security association (SA) of a virtual private network (VPN), and if the VPN egress is employing the optional anti-replay feature, it could inappropriately discard Classic packets (or discard the records in Classic packets) by mistaking their greater queuing delay for a replay attack (see "Dropped Packets for Tunnels with Replay Protection Enabled" in [Heist21] for the potential performance impact). This known problem is common to both IPsec [RFC4301] and DTLS [RFC6347] VPNs, given they use similar anti-replay window mechanisms. The mechanism used can only check for replay within its window, so if the window is smaller than the degree of reordering, it can only assume there might be a replay attack and discard all the packets behind the trailing edge of the window. The specifications of IPsec AH [RFC4302] and ESP [RFC4303] suggest that an implementer scales the size of the anti-replay window with interface speed, and DTLS 1.3 [I-D.ietf-tls-dtls13] says "The receiver SHOULD pick a window large enough to handle any plausible reordering, which depends on the data rate." However, in practice, the size of a VPN's anti-replay window is not always scaled appropriately.¶
If a VPN carrying traffic participating in the L4S experiment experiences inappropriate replay detection, the foremost remedy would be to ensure that the egress is configured to comply with the above window-sizing requirements.¶
If an implementation of a VPN egress does not support a sufficiently large anti-replay window, e.g. due to hardware limitations, one of the temporary alternatives listed in order of preference below might be feasible instead:¶
If the above is not possible and it is necessary to use L4S, either of the following might be appropriate as a last resort:¶
Modification to VPN implementations is outside the present scope, which is why this section has so far focused on reconfiguration. Although this document does not define any requirements for VPN implementations, determining whether there is a need for such requirements could be one aspect of L4S experimentation.¶
This section describes open questions that L4S Experiments ought to focus on. This section also documents outstanding open issues that will need to be investigated as part of L4S experimentation, given they could not be fully resolved during the WG phase. It also lists metrics that will need to be monitored during experiments (summarizing text elsewhere in L4S documents) and finally lists some potential future directions that researchers might wish to investigate.¶
In addition to this section, the DualQ spec [I-D.ietf-tsvwg-aqm-dualq-coupled] sets operational and management requirements for experiments with DualQ Coupled AQMs; and 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 the guidelines in [RFC5706] provides a helpful checklist.¶
L4S experiments would be expected to answer the following questions:¶
Have all the parts of L4S been deployed, and if so, what proportion of paths support it?¶
Did use of L4S over the Internet result in improvements to the following metrics:¶
Was per-flow queue protection typically (un)necessary?¶
Monitoring for harm to other traffic, specifically bandwidth starvation or excess queuing delay, will need to be conducted alongside all early L4S experiments. It is hard, if not impossible, for an individual flow to measure its impact on other traffic. So such monitoring will need to be conducted using bespoke monitoring across flows and/or across classes of traffic.¶
Researchers might find that L4S opens up the following interesting areas for investigation:¶
The 01 codepoint of the ECN Field of the IP header is specified by the present Experimental RFC. The process for an experimental RFC to assign this codepoint in the IP header (v4 and v6) is documented in Proposed Standard [RFC8311], which updates the Proposed Standard [RFC3168].¶
When the present document is published as an RFC, IANA is asked to update the 01 entry in the registry, "ECN Field (Bits 6-7)" to the following (see https://www.iana.org/assignments/dscp-registry/dscp-registry.xhtml#ecn-field ):¶
| Binary | Keyword | References |
|---|---|---|
| 01 | ECT(1) (ECN-Capable Transport(1))[1] | [RFC8311] [RFC Errata 5399] [RFCXXXX] |
[XXXX is the number that the RFC Editor assigns to the present document (this sentence to be removed by the RFC Editor)].¶
Approaches to assure the integrity of signals using the new identifier are introduced in Appendix C.1. See the security considerations in the L4S architecture [I-D.ietf-tsvwg-l4s-arch] for further discussion of mis-use of the identifier, as well as extensive discussion of policing rate and latency in regard to L4S.¶
Defining ECT(1) as the L4S identifier introduces a difference between the effects of ECT0 and ECT1 that RFC 3168 previously defined as distinct but with equivalent effect. For L4S ECN, a network node is still required not to swap one to the other, even if the network operator chooses to locally apply the treatment associated with the opposite codepoint (see Sections 5.4.1.1 and 5.4.1.2). These sections also describe the potential effects if a non-compliant or malicious network node does swap one to the other. The present specification does not change the effects of other unexpected transitions of the ECN field, which are still as described in Section 18 of RFC 3168.¶
If the anti-replay window of a VPN egress is too small, it will mistake deliberate delay differences as a replay attack, and discard higher delay packets (e.g. Classic) carried within the same security association (SA) as low delay packets (e.g. L4S). Section 6.2 recommends that VPNs used in L4S experiments are configured with a sufficiently large anti-replay window, as required by the relevant specifications. It also discusses other alternatives.¶
If a user taking part in the L4S experiment sets up a VPN without being aware of the above advice, and if the user allows anyone to send traffic into their VPN, they would open up a DoS vulnerability in which an attacker could induce the VPN's anti-replay mechanism to discard enough of the user's Classic (C) traffic (if they are receiving any) to cause a significant rate reduction. While the user is actively downloading C traffic, the attacker sends C traffic into the VPN to fill the remainder of the bottleneck link, then sends intermittent L4S packets to maximize the chance of exceeding the VPN's replay window. The user can prevent this attack by following the recommendations in Section 6.2.¶
The recommendation to detect loss in time units prevents the ACK-splitting attacks described in [Savage-TCP].¶
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 rationale for safety improvements (the requirements in Section 4) this appendix also includes preferable performance improvements (optimizations).¶
The requirements and recommendations in Section 4) have become know as the Prague L4S Requirements, because they were originally identified at an ad hoc meeting during IETF-94 in Prague [TCPPrague]. They were originally called the 'TCP Prague Requirements', but they are not solely applicable to TCP, so the name and wording has been generalized for all transport protocols, and the name 'TCP Prague' is now used for a specific implementation of the requirements.¶
At the time of writing, DCTCP [RFC8257] is 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 the requirements in Section 4 apply equally to other scalable congestion controls, covering adaptive real-time media, etc., not just capacity-seeking behaviours.¶
Description: A scalable congestion control needs to distinguish the packets it sends from those sent by Classic congestion controls (see the precise normative requirement wording in Section 4.1).¶
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 (see the precise normative requirement wording in Section 4.2).¶
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 a change to the TCP protocol to satisfy these requirements. Most other transport protocols already satisfy this requirement (see Section 4.2).¶
Description: It needs to be possible to replace the implementation of a scalable congestion control with a Classic control (see the precise normative requirement wording in Section 4.3).¶
Motivation: L4S is an experimental protocol, therefore it seems prudent to be able to disable it at source in case of insurmountable problems, perhaps due to some unexpected interaction on a particular sender; over a particular path or network; with a particular receiver or even ultimately an insurmountable problem with the experiment as a whole.¶
Description: As well as responding to ECN markings in a scalable way, a scalable congestion control needs to react to packet loss in a way that will coexist safely with a Reno congestion control [RFC5681] (see the precise normative requirement wording in Section 4.3).¶
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, 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 transport like TCP. Therefore, the particular Classic congestion control behaviour to fall back on will need to be dependent on the specific congestion control implementation. In the particular case of DCTCP, the DCTCP specification [RFC8257] states that "It is RECOMMENDED that an implementation deal with loss episodes in the same way as conventional TCP." For safe deployment, Section 4.3 requires any specification of a scalable congestion control for the public Internet to define the above requirement as a "MUST".¶
Even though a bottleneck is L4S capable, it might still become overloaded and have to drop packets. In this case, the sender may receive a high proportion of packets marked with the CE bit set and also experience loss. Current DCTCP implementations each 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). A third approach for the public Internet has been proposed that adjusts the loss response to result in a halving when combined with the ECN response. 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: Monitoring has to be in place so that a non-L4S but ECN-capable AQM can be detected at path bottlenecks. This is in case such an AQM has been implemented in a shared queue, in which case any long-running scalable flow would predominate over any simultaneous long-running Classic flow sharing the queue. The precise requirement wording in Section 4.3 is written so that such a problem could either be resolved in real-time, or via administrative intervention.¶
Motivation: Similarly to the discussion in Appendix A.1.4, this requirement in Section 4.3 is a safety condition to ensure an L4S congestion control coexists well with Classic flows when it builds a queue at a shared network bottleneck that has not been upgraded to support L4S. Nonetheless, if necessary, it is considered reasonable to resolve such problems over management timescales (possibly involving human intervention) because:¶
Therefore, the relevant normative requirement (Section 4.3) is divided into three stages: monitoring, detection and action:¶
This involves the act of switching the sender to a Classic congestion control. This might occur in real-time within the congestion control for the subsequent duration of a flow, or it might involve administrative action to switch to Classic congestion control for a specific interface or for a certain set of destination addresses.¶
Instead of the sender taking action itself, the operator of the sender (e.g. a CDN) might prefer to ask the network operator to modify the Classic AQM's treatment of L4S packets; or to ensure L4S packets bypass the AQM; or to upgrade the AQM to support L4S (see the L4S operational guidance [I-D.ietf-tsvwg-l4sops]). Once L4S flows no longer shared the Classic ECN AQM they would obviously no longer detect it, and the requirement to act on it would no longer apply.¶
The whole set of normative requirements concerning Classic ECN AQMs in Section 4.3 is worded so that it does not apply in controlled environments, such as private networks or data centre networks. CDN servers placed within an access ISP's network can be considered as a single controlled environment, but any onward networks served by the access network, including all the attached customer networks, would be unlikely to fall under the same degree of coordinated control. Monitoring is expressed as a 'MUST' for these uncontrolled segments of paths (e.g. beyond the access ISP in a home network), because there is a possibility that there might be a shared queue Classic ECN AQM in that segment. Nonetheless, the intent of the wording is to only require occasional monitoring of these uncontrolled regions, and not to burden CDN operators if monitoring never uncovers any potential problems.¶
More detailed discussion of all the above options and alternatives can be found in the L4S operational guidance [I-D.ietf-tsvwg-l4sops].¶
Having said all the above, the approach recommended in Section 4.3 is to monitor, detect and act in real-time on live traffic. A passive monitoring algorithm to detect a Classic ECN AQM at the bottleneck and fall back to Classic congestion control is described in an extensive technical report [ecn-fallback], which also provides a link to Linux source code, and a large online visualization of its evaluation results. Very briefly, the algorithm primarily monitors RTT variation using the same algorithm that maintains the mean deviation of TCP's smoothed RTT, but it smooths over a duration of the order of a Classic sawtooth. The outcome is also conditioned on other metrics such as the presence of CE marking and congestion avoidance phase having stabilized. The report also identifies further work to improve the approach, for instance improvements with low capacity links and combining the measurements with a cache of what had been learned about a path in previous connections. The report also suggests alternative approaches.¶
Although using passive measurements within live traffic (as above) can detect a Classic ECN AQM, it is much harder (perhaps impossible) to determine whether or not the AQM is in a shared queue. Nonetheless, this is much easier using active test traffic out-of-band, because two flows can be used. Section 4 of the same report [ecn-fallback] describes a simple technique to detect a Classic ECN AQM and determine whether it is in a shared queue, summarized here.¶
An L4S-enabled test server could be set up so that, when a test client accesses it, it serves a script that gets the client to open two parallel long-running flows. It could serve one with a Classic congestion control (C, that sets ECT(0)) and one with a scalable CC (L, that sets ECT(1)). If neither flow induces any ECN marks, it can be presumed the path does not contain a Classic ECN AQM. If either flow induces some ECN marks, the server could measure the relative flow rates and round trip times of the two flows. Table 2 shows the AQM that can be inferred for various cases (presuming the AQM behaviours known at the time of writing).¶
| Rate | RTT | Inferred AQM |
|---|---|---|
| L > C | L = C | Classic ECN AQM (FIFO) |
| L = C | L = C | Classic ECN AQM (FQ) |
| L = C | L < C | FQ-L4S AQM |
| L ~= C | L < C | Coupled DualQ AQM |
Finally, we motivate the recommendation in Section 4.3 that a scalable congestion control is not expected to change to setting ECT(0) while it adapts its behaviour to coexist with Classic flows. This is because the sender needs to continue to check whether it made the right decision - and switch back if it was wrong, or if a different link becomes the bottleneck:¶
Note that if a flow is configured to only use a Classic congestion control, it is then entirely appropriate not to use ECT(1).¶
Description: A scalable congestion control needs to reduce RTT bias as much as possible at least over the low to typical range of RTTs that will interact in the intended deployment scenario (see the precise normative requirement wording in Section 4.3).¶
Motivation: The throughput of Classic congestion controls 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 congestion controls have never allowed a very low RTT path to exist because they induce a large queue. For instance, consider two paths with base RTT 1 ms and 100 ms. If a Classic congestion control induces a 100 ms queue, it turns these RTTs into 101 ms and 200 ms leading to a throughput ratio of about 2:1. Whereas if a scalable congestion control induces only a 1 ms 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 in Section 4.3.¶
Over higher than typical RTTs, L4S flows can use the same RTT bias as in current Classic congestion controls and still work satisfactorily. So, there is no additional requirement in Section 4.3 for high RTT L4S flows to remove RTT bias - they can but they don't have to.¶
One way for a Scalable congestion control to satisfy these requirements is to make its additive increase behave as if it were a standard Reno flow but over a larger RTT by using a virtual RTT (rtt_virt) that is a function of the actual RTT (rtt). Example functions might be:¶
These example functions are chosen so that, as the actual RTT reduces from high to low, the virtual RTT reduces less (see [I-D.briscoe-iccrg-prague-congestion-control] for details).¶
However, short RTT flows can more rapidly respond to changes in available capacity, whether due to other flows arriving and departing or radio capacity varying. So it would wrong to require short RTT flows to be as sluggish as long-RTT flows, which would unnecessarily under-utilize capacity and result in unnecessary overshoots and undershoots (instability). Therefore, rather than requiring strict RTT-independence, the wording says "as independent of RTT as possible without compromising stability or utilization". This allows shorter RTT flows to exploit their agility advantage.¶
Description: A scalable congestion control needs to remain responsive to congestion when typical RTTs over the public Internet are significantly smaller because they are no longer inflated by queuing delay (see the precise normative requirement wording in Section 4.3).¶
Motivation: As currently specified, the minimum congestion window of ECN-capable TCP (and its derivatives) is expected to be 2 sender maximum segment sizes (SMSS), or 1 SMSS after a retransmission timeout. Once the congestion window reaches this minimum, if there is further ECN-marking, TCP is meant to wait for a retransmission timeout before sending another segment (see section 6.1.2 of the ECN spec [RFC3168]). In practice, most known window-based congestion control algorithms become unresponsive to ECN congestion signals at this point. No matter how much ECN marking, the congestion window no longer reduces. Instead, the sender's lack of any further congestion response forces the queue to grow, overriding any AQM and increasing queuing delay (making the window large enough to become responsive again). This can result in a stable but deeper queue, or it might drive the queue to loss, then the retransmission timeout mechanism acts as a backstop.¶
Most window-based congestion controls for other transport protocols have a similar minimum window, albeit when measured in bytes for those that use smaller packets.¶
L4S mechanisms significantly reduce queueing delay so, over the same path, the RTT becomes lower. Then this problem becomes surprisingly common [sub-mss-prob]. 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 SMSS if the RTT is 6 ms or less, which is quite common when accessing a nearby data centre. So, any more than two such parallel TCP flows will become unresponsive to ECN and increase queuing delay.¶
Unless scalable congestion controls address the requirement in Section 4.3 from the start, they will frequently become unresponsive to ECN, negating the low latency benefit of L4S, for themselves and for others.¶
That would seem to imply that scalable congestion controllers ought to be required to be able work with a congestion window less than 1 SMSS. For instance, if an ECN-capable TCP gets an ECN-mark when it is already sitting at a window of 1 SMSS, RFC 3168 requires it to defer sending for a retransmission timeout. A less drastic but more complex mechanism can maintain a congestion window less than 1 SMSS (significantly less if necessary), as described in [Ahmed19]. Other approaches are likely to be feasible.¶
However, the requirement in Section 4.3 is worded as a "SHOULD" because it is believed that the existence of a minimum window is not all bad. When competing with an unresponsive flow, a minimum window naturally protects the flow from starvation by at least keeping some data flowing.¶
By stating the requirement to go lower than 1 SMSS as a "SHOULD", while the requirement in RFC 3168 still stands as well, we shall be able to watch the choices of minimum window evolve in different scalable congestion controllers.¶
Description: When detecting loss, a scalable congestion control needs to be tolerant to reordering over an adaptive time interval, which scales with throughput, rather than counting only in fixed units of packets, which does not scale (see the precise normative requirement wording in Section 4.3).¶
Motivation: A primary purpose of L4S is scalable throughput (it's in the name). Scalability in all dimensions is, of course, also a goal of all IETF technology. The inverse linear congestion response in Section 4.3 is necessary, but not sufficient, to solve the congestion control scalability problem identified in [RFC3649]. As well as maintaining frequent ECN signals as rate scales, it is also important to ensure that a potentially false perception of loss does not limit throughput scaling.¶
End-systems cannot know whether a missing packet is due to loss or reordering, except in hindsight - if it appears later. So they can only deem that there has been a loss if a gap in the sequence space has not been filled, either after a certain number of subsequent packets has arrived (e.g. the 3 DupACK rule of standard TCP congestion control [RFC5681]) or after a certain amount of time (e.g. the RACK approach [RFC8985]).¶
As we attempt to scale packet rate over the years:¶
Therefore hosts have an incentive to detect loss in time units (so as not to fool themselves too often into detecting losses when there are none). And for hosts that are changing their congestion control implementation to L4S, there is no downside to including time-based loss detection code in the change (loss recovery implemented in hardware is an exception, covered later). Therefore requiring L4S hosts to detect loss in time-based units would not be a burden.¶
If the requirement in Section 4.3 were not placed on L4S hosts, even though it would be no burden on hosts to comply, all networks would face unnecessary uncertainty over whether some L4S hosts might be detecting loss by counting packets. Then all link technologies will have to unnecessarily keep reducing the time within which reordering occurs. That is not a problem for some link technologies, but it becomes increasingly challenging for other link technologies to continue to scale, particularly those relying on channel bonding for scaling, such as LTE, 5G and DOCSIS.¶
Given Internet paths traverse many link technologies, any scaling limit for these more challenging access link technologies would become a scaling limit for the Internet as a whole.¶
It might be asked how it helps to place this loss detection requirement only on L4S hosts, because networks will still face uncertainty over whether non-L4S flows are detecting loss by counting DupACKs. The answer is that those link technologies for which it is challenging to keep squeezing the reordering time will only need to do so for non-L4S traffic (which they can do because the L4S identifier is visible at the IP layer). Therefore, they can focus their processing and memory resources into scaling non-L4S (Classic) traffic. Then, the higher the proportion of L4S traffic, the less of a scaling challenge they will have.¶
To summarize, there is no reason for L4S hosts not to be part of the solution instead of part of the problem.¶
Requirement ("MUST") or recommendation ("SHOULD")? As explained above, this is a subtle interoperability issue between hosts and networks, which seems to need a "MUST". Unless networks can be certain that all L4S hosts follow the time-based approach, they still have to cater for the worst case - continually squeeze reordering into a smaller and smaller duration - just for hosts that might be using the counting approach. However, it was decided to express this as a recommendation, using "SHOULD". The main justification was that networks can still be fairly certain that L4S hosts will follow this recommendation, because following it offers only gain and no pain.¶
Details:¶
The speed of loss recovery is much more significant for short flows than long, therefore a good compromise is to 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 broadly the approach adopted by TCP RACK (Recent ACKnowledgements) [RFC8985]. However, RACK starts with the 3 DupACK approach, because the RTT estimate is not necessarily stable. As long as the initial window is paced, such initial use of 3 DupACK counting would amount to time-based loss detection and therefore would satisfy the time-based loss detection recommendation of Section 4.3. This is because pacing of the initial window would ensure that 3 DupACKs early in the connection would be spread over a small fraction of the round trip.¶
As mentioned above, hardware implementations of loss recovery using DupACK counting exist (e.g. some implementations of RoCEv2 for RDMA). For low latency, these implementations can change their congestion control to implement L4S, because the congestion control (as distinct from loss recovery) is implemented in software. But they cannot easily satisfy this loss recovery requirement. However, it is believed they do not need to, because such implementations are believed to solely exist in controlled environments, where the network technology keeps reordering extremely low anyway. This is why controlled environments with hardly any reordering are excluded from the scope of the normative recommendation in Section 4.3.¶
Detecting loss in time units also prevents the ACK-splitting attacks described in [Savage-TCP].¶
Description: This item concerns TCP and its derivatives (e.g. SCTP) as well as RTP/RTCP [RFC6679]. The original specification of ECN for TCP precluded the use of ECN on control packets and retransmissions. Similarly RFC 6679 precludes the use of ECT on RTCP datagrams, in case the path changes after it has been checked for ECN traversal. 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 other transports, e.g. SCTP, RTCP.¶
Motivation (TCP): 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. ECN++ [I-D.ietf-tcpm-generalized-ecn] 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 itself satisfies the accurate feedback requirement in Section 4.2 for using a scalable congestion control).¶
Motivation (RTCP): L4S experiments in general will need to observe the rule in the RTP ECN spec [RFC6679] that precludes ECT on RTCP datagrams. Nonetheless, as ECN usage becomes more widespread, it would be useful to conduct specific experiments with ECN-capable RTCP to gather data on whether such caution is necessary.¶
Description: It would improve performance if scalable congestion controls did not limit their congestion window increase to the standard additive increase of 1 SMSS 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 [RFC8257], DCTCP uses the traditional 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. TCP Cubic [RFC8312] was designed to solve this problem, but as flow rates have continued to increase, the delay accelerating into available capacity has become prohibitive. See, for instance, the examples in Section 5.1 of the L4S architecture [I-D.ietf-tsvwg-l4s-arch]. Even when out of its Reno-compatibility mode, every 8x scaling of Cubic's flow rate leads to 2x more acceleration delay.¶
In the steady state, DCTCP induces about 2 ECN marks per round trip, so it is possible to quickly detect when these signals have disappeared and seek available capacity more rapidly, while minimizing the impact on other flows (Classic and scalable) [LinuxPacedChirping]. Alternatively, approaches such as Adaptive Acceleration (A2DTCP [A2DTCP]) have been proposed to address this problem in data centres, which might be deployable over the public Internet.¶
Description: It would improve performance if scalable congestion controls converged (reached their steady-state share of the capacity) faster than Classic congestion controls or at least no slower. This affects 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 a Classic congestion control 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 a regular loss-based TCP slow start is even less favourable [Paced-Chirping] due to the shallow ECN marking threshold needed for L4S. It is exacerbated by the typically greater mismatch between the link rate of the sending host and typical Internet access bottlenecks. This problem is detrimental in general, but would particularly harm the performance of short flows relative to Classic congestion controls.¶
This appendix is informative, not normative. As explained in Section 2, there is insufficient space in the IP header (v4 or v6) to fully accommodate every requirement. So the choice of L4S identifier involves tradeoffs. This appendix records the pros and cons of the choice that was made.¶
Non-normative recap of the chosen codepoint scheme:¶
Packets with ECT(1) and conditionally packets with CE signify L4S semantics as an alternative to the semantics of Classic ECN [RFC3168], specifically:¶
Cons:¶
Classifying all CE packets into the L4S queue risks any CE packets that were originally ECT(0) being incorrectly classified as L4S. If there were delay in the Classic queue, these incorrectly classified CE packets would arrive early, which is a form of reordering. Reordering within a microflow can cause TCP senders (and senders of similar transports) to retransmit spuriously. However, the risk of spurious retransmissions would be extremely low for the following reasons:¶
Even then, when a few packets are delivered early, it takes very unusual conditions to cause a spurious retransmission, in contrast to when some packets are delivered late. The first bottleneck has to apply CE-marks to at least N contiguous packets and the second bottleneck has to inject an uninterrupted sequence of at least N of these packets between two packets earlier in the stream (where N is the reordering window that the transport protocol allows before it considers a packet is lost).¶
It is extremely unlikely that the above set of 5 eventualities that are each unusual in themselves would all happen simultaneously. But, even if they did, the consequences would hardly be dire: the odd spurious fast retransmission. Whenever the traffic source (a Classic congestion control) mistakes the reordering of a string of CE marks for a loss, one might think that it will reduce its congestion window as well as emitting a spurious retransmission. However, it would have already reduced its congestion window when the CE markings arrived early. If it is using ABE [RFC8511], it might reduce cwnd a little more for a loss than for a CE mark. But it will revert that reduction once it detects that the retransmission was spurious.¶
In conclusion, the impact of early reordering on spurious retransmissions due to CE being ambiguous will generally be vanishingly small.¶
With this scheme, when a source receives ECN feedback, it is not explicitly clear which type of AQM generated the CE markings. This is not a problem for Classic ECN sources that send ECT(0) packets, because an L4S AQM will recognize the ECT(0) packets as Classic and apply the appropriate Classic ECN marking behaviour.¶
However, in the absence of explicit disambiguation of the CE markings, an L4S source needs to use heuristic techniques to work out which type of congestion response to apply (see Appendix A.1.5). Otherwise, if long-running Classic flow(s) are sharing a Classic ECN AQM bottleneck with long-running L4S flow(s), which then apply an L4S response to Classic CE signals, the L4S flows would outcompete the Classic flow(s). Experiments have shown that L4S flows can take about 20 times more capacity share than equivalent Classic flows. Nonetheless, as link capacity reduces (e.g. to 4 Mb/s), the inequality reduces. So Classic flows always make progress and are not starved.¶
When L4S was first proposed (in 2015, 14 years after the Classic ECN spec [RFC3168] was published), it was believed that Classic ECN AQMs had failed to be deployed, because research measurements had found little or no evidence of CE marking. In subsequent years Classic ECN was included in per-flow-queuing (FQ) deployments, however an FQ scheduler stops an L4S flow outcompeting Classic, because it enforces equality between flow rates. It is not known whether there have been any non-FQ deployments of Classic ECN AQMs in the subsequent years, or whether there will be in future.¶
An algorithm for detecting a Classic ECN AQM as soon as a flow stabilizes after start-up has been proposed [ecn-fallback] (see Appendix A.1.5 for a brief summary). Testbed evaluations of v2 of the algorithm have shown detection is reasonably good for Classic ECN AQMs, in a wide range of circumstances. However, although it can correctly detect an L4S ECN AQM in many circumstances, its is often incorrect at low link capacities and/or high RTTs. Although this is the safe way round, there is a danger that it will discourage use of the algorithm.¶
Pros:¶
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 as a nonce for integrity protection of congestion notifications in future. The ECN Nonce RFC [RFC3540] has been reclassified as historic, partly because other ways (that do not consume a codepoint in the IP header) have been developed to protect feedback integrity of TCP and other transports [RFC8311]. 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 identifier 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 the L4S service identifier defined in Section 3.¶
Thanks to Richard Scheffenegger, John Leslie, David Taeht, 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, Tom Henderson, David Black, Gorry Fairhurst, Brian Carpenter, Jake Holland, Rod Grimes, Richard Scheffenegger, Sebastian Moeller, Neal Cardwell, Praveen Balasubramanian, Reza Marandian Hagh, Pete Heist, Stuart Cheshire, Vidhi Goel, Mirja Kuehlewind, Ermin Sakic and Martin Duke for providing help and reviewing this draft and thanks to Ingemar Johansson for reviewing and providing substantial text. Thanks also to the area reviewers: Valery Smyslov, Ines Robles. Thanks to Sebastian Moeller for identifying the interaction with VPN anti-replay and to Jonathan Morton for identifying the attack based on this. Particular thanks to tsvwg chairs Gorry Fairhurst, David Black and Wes Eddy for patiently helping this and the other L4S drafts through the IETF process. 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). The contribution of Koen De Schepper was also part-funded by the 5Growth and DAEMON EU H2020 projects. Bob Briscoe was also funded partly by the Research Council of Norway through the TimeIn project, partly by CableLabs and partly by the Comcast Innovation Fund. The views expressed here are solely those of the authors.¶