Forwarding Metamorphosis: Fast Programmable Match

Forwarding Metamorphosis: Fast Programmable Match
Forwarding Metamorphosis: Fast Programmable
Match-Action Processing in Hardware for SDN
Pat Bosshart† , Glen Gibb‡ , Hun-Seok Kim† , George Varghese§ , Nick McKeown‡ ,
Martin Izzard† , Fernando Mujica† , Mark Horowitz‡
†
Texas Instruments
‡
Stanford University
§
Microsoft Research
[email protected] {grg, nickm, horowitz}@stanford.edu
[email protected] {hkim, izzard, fmujica}@ti.com
ABSTRACT
In Software Defined Networking (SDN) the control plane
is physically separate from the forwarding plane. Control
software programs the forwarding plane (e.g., switches and
routers) using an open interface, such as OpenFlow. This
paper aims to overcomes two limitations in current switching chips and the OpenFlow protocol: i) current hardware
switches are quite rigid, allowing “Match-Action” processing
on only a fixed set of fields, and ii) the OpenFlow specification only defines a limited repertoire of packet processing
actions. We propose the RMT (reconfigurable match tables) model, a new RISC-inspired pipelined architecture for
switching chips, and we identify the essential minimal set
of action primitives to specify how headers are processed in
hardware. RMT allows the forwarding plane to be changed
in the field without modifying hardware. As in OpenFlow,
the programmer can specify multiple match tables of arbitrary width and depth, subject only to an overall resource
limit, with each table configurable for matching on arbitrary
fields. However, RMT allows the programmer to modify all
header fields much more comprehensively than in OpenFlow.
Our paper describes the design of a 64 port by 10 Gb/s
switch chip implementing the RMT model. Our concrete
design demonstrates, contrary to concerns within the community, that flexible OpenFlow hardware switch implementations are feasible at almost no additional cost or power.
Categories and Subject Descriptors
C.2.1 [Computer-Communication Networks]: Network
Architecture and Design—Network Communications
General Terms
Design
Keywords
RMT model, Reconfigurable Match Tables, SDN
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1.
INTRODUCTION
To improve is to change; to be perfect is to change
often.
— Churchill
Good abstractions—such as virtual memory and timesharing—are paramount in computer systems because they
allow systems to deal with change and allow simplicity of
programming at the next higher layer. Networking has progressed because of key abstractions: TCP provides the abstraction of connected queues between endpoints, and IP
provides a simple datagram abstraction from an endpoint to
the network edge. However, routing and forwarding within
the network remain a confusing conglomerate of routing protocols (e.g., BGP, ICMP, MPLS) and forwarding behaviors
(e.g., routers, bridges, firewalls), and the control and forwarding planes remain intertwined inside closed, vertically
integrated boxes.
Software-defined networking (SDN) took a key step in abstracting network functions by separating the roles of the
control and forwarding planes via an open interface between
them (e.g., OpenFlow [27]). The control plane is lifted up
and out of the switch, placing it in external software. This
programmatic control of the forwarding plane allows network owners to add new functionality to their network, while
replicating the behavior of existing protocols. OpenFlow has
become quite well-known as an interface between the control plane and the forwarding plane based on the approach
known as “Match-Action”. Roughly, a subset of packet bytes
are matched against a table; the matched entry specifies a
corresponding action(s) that are applied to the packet.
One can imagine implementing Match-Action in software
on a general purpose CPU. But for the speeds we are interested in—about 1 Tb/s today—we need the parallelism
of dedicated hardware. Switching chips have remained two
orders of magnitude faster at switching than CPUs for a
decade, and an order of magnitude faster than network processors, and the trend is unlikely to change. We therefore
need to think through how to implement Match-Action in
hardware to exploit pipelining and parallelism, while living
within the constraints of on-chip table memories.
There is a natural tradeoff between programmability and
speed. Today, supporting new features frequently requires
replacing the hardware. If Match-Action hardware permitted (just) enough reconfiguration in the field so that new
types of packet processing could be supported at run-time,
then it would change how we think of programming the network. The real question here is whether it can be done at
reasonable cost without sacrificing speed.
Single Match Table: The simplest approach is to abstract matching semantics in what we call the SMT (Single
Match Table) model. In SMT, the controller tells the switch
to match any set of packet header fields against entries in
a single match table. SMT assumes that a parser locates
and extracts the correct header fields to match against the
table. For example, an Ethernet packet may have an optional MPLS tag, which means the IP header can be in two
different locations. The match is a binary exact match when
all fields are completely specified, and is ternary for matches
where some bits are switched off (wildcard entries). Superficially, the SMT abstraction is good for both programmers
(what could be simpler than a single match?) and implementers (SMT can be implemented using a wide Ternary
Content Addressable Memory (TCAM)). Note that the forwarding data plane abstraction has the most rigorous hardware implementation constraints because forwarding is often
required to operate at about 1 Tb/s.
A closer look, however, shows that the SMT model is
costly in use because of a classic problem. The table needs
to store every combination of headers; this is wasteful if the
header behaviors are orthogonal (the entries will have many
wildcard bits). It can be even more wasteful if one header
match affects another, for example if a match on the first
header determines a disjoint set of values to match on the
second header (e.g., in a virtual router [11]), requiring the
table to hold the Cartesian-product of both.
Multiple Match Tables: MMT (Multiple Match Tables) is a natural refinement of the SMT model. MMT goes
beyond SMT in an important way: it allows multiple smaller
match tables to be matched by a subset of packet fields. The
match tables are arranged into a pipeline of stages; processing at stage j can be made to depend on processing from
stage i < j by stage i modifying the packet headers or other
information passed to stage j. MMT is easy to implement
using a set of narrower tables in each stage; in fact, it is close
enough to the way existing switch chips are implemented to
make it easy to map onto existing pipelines [3, 14, 23, 28].
Google reports converting their entire private WAN to this
approach using merchant switch chips [13].
The OpenFlow specification transitioned to the MMT
model [31] but does not mandate the width, depth, or even
the number of tables, leaving implementors free to choose
their multiple tables as they wish. While a number of fields
have been standardized (e.g., IP and Ethernet fields), OpenFlow allows the introduction of new match fields through a
user-defined field facility.
Existing switch chips implement a small (4–8) number of
tables whose widths, depths, and execution order are set
when the chip is fabricated. But this severely limits flexibility. A chip used for a core router may require a very
large 32-bit IP longest matching table and a small 128 bit
ACL match table; a chip used for an L2 bridge may wish to
have a 48-bit destination MAC address match table and a
second 48-bit source MAC address learning table; an enterprise router may wish to have a smaller 32-bit IP prefix table
and a much larger ACL table as well as some MAC address
match tables. Fabricating separate chips for each use case is
inefficient, and so merchant switch chips tend to be designed
to support the superset of all common configurations, with a
set of fixed size tables arranged in a pre-determined pipeline
order. This creates a problem for network owners who want
to tune the table sizes to optimize for their network, or im-
plement new forwarding behaviors beyond those defined by
existing standards. In practice, MMT often translates to
fixed multiple match tables.
A second subtler problem is that switch chips offer only a
limited repertoire of actions corresponding to common processing behaviors, e.g., forwarding, dropping, decrementing
TTLs, pushing VLAN or MPLS headers, and GRE encapsulation. And to date, OpenFlow specifies only a subset of
these. This action set is not easily extensible, and is also not
very abstract. A more abstract set of actions would allow
any field to be modified, any state machine associated with
the packet to be updated, and the packet to be forwarded
to an arbitrary set of output ports.
Reconfigurable Match Tables: Thus in this paper,
we explore a refinement of the MMT model that we call
RMT (Reconfigurable Match Tables). Like MMT, ideal
RMT would allow a set of pipeline stages each with a match
table of arbitrary depth and width. RMT goes beyond MMT
by allowing the data plane to be reconfigured in the following four ways. First, field definitions can be altered and
new fields added; second, the number, topology, widths, and
depths of match tables can be specified, subject only to an
overall resource limit on the number of matched bits; third,
new actions may be defined, such as writing new congestion
fields; fourth, arbitrarily modified packets can be placed in
specified queue(s), for output at any subset of ports, with
a queuing discipline specified for each queue. This configuration should be managed by an SDN controller, but we do
not define the control protocol in this paper.
The benefits of RMT can be seen by considering new protocols proposed in the last few years, such as PBB [16],
VxLAN [22], NVGRE [19], STT [21], and OTV [20]. Each
protocol defines new header fields. Without an architecture
like RMT, new hardware would be required to match on and
process these protocols.
Note that RMT is perfectly compatible with (and even
partly implemented by) the current OpenFlow specification.
Individual chips can clearly allow an interface to reconfigure
the data plane. In fact, some existing chips, driven at least
in part by the need to address multiple market segments,
already have some flavors of reconfigurability that can be
expressed using ad hoc interfaces to the chip.
Many researchers have recognized the need for something
akin to RMT and have advocated for it. For example, the
IETF ForCES working group developed the definition of a
flexible data plane [17]; similarly, the forwarding abstraction
working group in ONF has worked on reconfigurability [30].
However, there has been understandable skepticism that the
RMT model is implementable at very high speeds. Without
a chip to provide an existence proof of RMT, it has seemed
fruitless to standardize the reconfiguration interface between
the controller and the data plane.
Intuitively, arbitrary reconfigurability at terabit speeds
seems an impossible mission. But what restricted form of
reconfigurability is feasible at these speeds? Does the restricted reconfigurability cover a large enough fraction of
the needs we alluded to earlier? Can one prove feasibility
via working silicon that embodies these ideas? How expensive is such an RMT chip compared to a fixed-table MMT
chip? These are the questions we address in this paper.
General purpose payload processing is not our goal.
SDN/OpenFlow (and our design) aim to identify the essential minimal set of primitives to process headers in hardware.
Think of it as a minimal instruction set like RISC, designed
to run really fast in heavily pipelined hardware. Our very
flexible design is cost-competitive with fixed designs—i.e.,
flexibility comes at almost no cost.
Paper Contributions: Our paper makes a concrete contribution to the debate of what forwarding abstractions are
practical at high speed, and the extent to which a forwarding
plane can be reconfigured by the control plane. Specifically,
we address the questions above as follows:
1. An architecture for RMT (§2): We describe an RMT
switch architecture that allows definition of arbitrary headers and header sequences, arbitrary matching of fields by
an arbitrary number of tables, arbitrary writing of packet
header fields (but not the packet body), and state update
per packet. Several restrictions are introduced to make the
architecture realizable. We outline how a desired configuration can be expressed by a parse graph to define headers,
and a table flow graph to express the match table topology.
2. Use cases (§3): We provide use cases that show how
the RMT model can be configured to implement forwarding
using Ethernet and IP headers, and support RCP [8].
3. Chip design and cost (§4–5): We show that the specific form of reconfigurability we advocate is indeed feasible and describe the implementation of a 640 Gb/s (64 ×
10 Gb/s) switch chip. Our architecture and implementation
study included significant detail in logic and circuit design,
floorplanning and layout, using techniques proven over the
design team’s long history of developing complex digital ICs.
An industry standard 28nm process was used. This work is
necessary to prove feasibility in meeting goals such as timing
and chip area (cost). We have not produced a complete design or actual silicon. Based on our investigation, we show
that the cost of reconfiguration is expected to be modest:
less than 20% beyond the cost of a fixed (non-reconfigurable)
version.
We make no claim that we are the first to advocate reconfigurable matching or that our proposed reconfiguration
functionality is the “right” one. We do claim that it is important to begin the conversation by making a concrete definition of the RMT model and showing it is feasible by exhibiting a chip, as we have attempted to do in this paper. While
chip design is not normally the province of SIGCOMM, our
chip design shows that a rather general form of the RMT
model is feasible and inexpensive. We show that the RMT
model is not only a good way to think about programming
the network, but also lends itself to direct expression in hardware using a configurable pipeline of match tables and action
processors.
2.
RMT ARCHITECTURE
We spoke of RMT as “allow(ing) a set of pipeline stages
. . . each with a match table of arbitrary depth and width
that matches on fields”. A logical deduction is that an RMT
switch consists of a parser, to enable matching on fields,
followed by an arbitrary number of match stages. Prudence
suggests that we include some kind of queuing to handle
congestion at the outputs.
Let’s look a little deeper. The parser must allow field definitions to be modified or added, implying a reconfigurable
parser. The parser output is a packet header vector, which is
a set of header fields such as IP dest, Ethernet dest, etc. In
addition, the packet header vector includes “metadata” fields
such as the input port on which the packet arrived and other
router state variables (e.g., current size of router queues).
The vector flows through a sequence of logical match stages,
each of which abstracts a logical unit of packet processing
(e.g., Ethernet or IP processing) in Figure 1a.
Each logical match stage allows the match table size to
be configured: for IP forwarding, for example, one might
want a match table of 256K 32-bit prefixes and for Ethernet
a match table of 64K 48-bit addresses. An input selector
picks the fields to be matched upon. Packet modifications
are done using a wide instruction (the VLIW—very long
instruction word—block in Figure 1c) that can operate on
all fields in the header vector concurrently.
More precisely, there is an action unit for each field F in
the header vector (Figure 1c), which can take up to three
input arguments, including fields in the header vector and
the action data results of the match, and rewrite F . Allowing each logical stage to rewrite every field may seem like
overkill, but it is useful when shifting headers; we show later
that action unit costs are small compared to match tables.
A logical MPLS stage may pop an MPLS header, shifting
subsequent MPLS headers forward, while a logical IP stage
may simply decrement TTL. Instructions also allow limited
state (e.g., counters) to be modified that may influence the
processing of subsequent packets.
Control flow is realized by an additional output, nexttable-address, from each table match that provides the index of the next table to execute. For example, a match on
a specific Ethertype in Stage 1 could direct later processing
stages to do prefix matching on IP (routing), while a different Ethertype could specify exact matching on Ethernet
DAs (bridging). A packet’s fate is controlled by updating a
set of destination ports and queues; this can be used to drop
a packet, implement multicast, or apply specified QoS such
as a token bucket.
A recombination block is required at the end of the pipeline
to push header vector modifications back into the packet
(Figure 1a). Finally, the packet is placed in the specified
queues at the specified output ports and a configurable queuing discipline applied.
In summary, the ideal RMT of Figure 1a allows new fields
to be added by modifying the parser, new fields to be matched
by modifying match memories, new actions by modifying
stage instructions, and new queueing by modifying the queue
discipline for each queue. An ideal RMT can simulate existing devices such as a bridge, a router, or a firewall; and
can implement existing protocols, such as MPLS and ECN,
and protocols proposed in the literature, such as RCP [8]
that uses non-standard congestion fields. Most importantly,
it allows future data plane modifications without modifying
hardware.
2.1
Implementation Architecture at 640Gb/s
We advocate an implementation architecture shown in
Figure 1b that consists of a large number of physical pipeline
stages that a smaller number of logical RMT stages can be
mapped to, depending on the resource needs of each logical
stage. This implementation architecture is motivated by:
1. Factoring State: Router forwarding typically has several stages (e.g., forwarding, ACL), each of which uses a
separate table; combining these into one table produces the
cross-product of states. Stages are processed sequentially
with dependencies, so a physical pipeline is natural.
2. Flexible Resource Allocation Minimizing Resource Waste:
Logical Stage 1
Switch State
(metadata)
Packets
1
...
...
Payload
...
K
...
Configurable
Output
Queues
Recombine
VLIW
Action
New Header
...
Select
Match
Tables
Logical Stage N
State
Prog.
Parser
1
Input
Channels
...
Header
Packets
Statistics
K
Output
Channels
(a) RMT model as a sequence of logical Match-Action stages.
Physical
Stage 1
Physical
Stage 2
Physical
Stage M
Src 1
: Ingress logical
match tables
Match
Results
Action
Memory
OP code
(from inst
mem)
Packet
Header
Vector
Src 1
Src 2
Src 3
Action
Unit
Match
Tables
Src 3
...
Logical Stage 2
Packet
Header
Vector
...
Logical
Stage N
Very Wide
Header Bus
...
Action Input Selector (Crossbar)
Logical Stage 1
Action
Unit
Src 2
OP code
Ctrl
: Egress logical
match tables
VLIW Instruction Memory
(b) Flexible match table configuration.
(c) VLIW action architecture.
Figure 1: RMT model architecture.
A physical pipeline stage has some resources (e.g., CPU,
memory). The resources needed for a logical stage can vary
considerably. For example, a firewall may require all ACLs,
a core router may require only prefix matches, and an edge
router may require some of each. By flexibly allocating physical stages to logical stages, one can reconfigure the pipeline
to metamorphose from a firewall to a core router in the field.
The number of physical stages N should be large enough
so that a logical stage that uses few resource will waste at
most 1/N -th of the resources. Of course, increasing N will
increase overhead (wiring, power): in our chip design we
chose N = 32 as a compromise between reducing resource
wastage and hardware overhead.
3. Layout Optimality: As shown in Figure 1b, a logical
stage can be assigned more memory by assigning the logical
stage to multiple contiguous physical stages. An alternate
design is to assign each logical stage to a decoupled set of
memories via a crossbar [4]. While this design is more flexible (any memory bank can be allocated to any stage), worst
case wire delays between
a processing stage and memories
√
will grow at least as M , which in router chips that require
a large amount of memory M can be large. While these
delays can be ameliorated by pipelining, the ultimate chal-
lenge in such a design is wiring: unless the current match
and action widths (1280 bits) are reduced, running so many
wires between every stage and every memory may well be
impossible.
In sum, the advantage of Figure 1b is that it uses a tiled
architecture with short wires whose resources can be reconfigured with minimal waste. We acknowledge two disadvantages. First, having a larger number of physical stages seems
to inflate power requirements. Second, this implementation
architecture conflates processing and memory allocation. A
logical stage that wants more processing must be allocated
two physical stages, but then it gets twice as much memory
even though it may not need it. In practice, neither issue
is significant. Our chip design shows that the power used
by the stage processors is at most 10% of the overall power
usage. Second, in networking most use cases are dominated
by memory use, not processing.
2.2
Restrictions for Realizability
The physical pipeline stage architecture needs restrictions
to allow terabit-speed realization:
Match restrictions: The design must contain a fixed number of physical match stages with a fixed set of resources.
Our chip design provides 32 physical match stages at both
ingress and egress. Match-action processing at egress allows
more efficient processing of multicast packets by deferring
per-port modifications until after buffering.
Packet header limits: The packet header vector containing
the fields used for matching and action has to be limited.
Our chip design limit is 4Kb (512B) which allows processing
fairly complex headers.
Memory restrictions: Every physical match stage contains
table memory of identical size. Match tables of arbitrary
width and depth are approximated by mapping each logical match stage to multiple physical match stages or fractions thereof (see Fig. 1b). For example, if each physical
match stage allows only 1,000 prefix entries, a 2,000 IP logical match table is implemented in two stages (upper-left
rectangle of Fig. 1b). Likewise, a small Ethertype match table could occupy a small portion of a match stage’s memory.
Hash-based binary match in SRAM is 6× cheaper in area
than TCAM ternary match. Both are useful, so we provide
a fixed amount of SRAM and TCAM in each stage. Each
physical stage contains 106 1K × 112b SRAM blocks, used
for 80b wide hash tables (overhead bits are explained later)
and to store actions and statistics, and 16 2K × 40b TCAM
blocks. Blocks may be used in parallel for wider matches,
e.g., a 160b ACL lookup using four blocks. Total memory
across the 32 stages is 370 Mb SRAM and 40 Mb TCAM.
Action restrictions: The number and complexity of instructions in each stage must be limited for realizability.
In our design, each stage may execute one instruction per
field. Instructions are limited to simple arithmetic, logical,
and bit manipulation (see §4.3). These actions allow implementation of protocols like RCP [8], but don’t allow packet
encryption or regular expression processing on the packet
body.
Instructions can’t implement state machine functionality; they may only modify fields in the packet header vector, update counters in stateful tables, or direct packets to
ports/queues. The queuing system provides four levels of
hierarchy and 2K queues per port, allowing various combinations of deficit round robin, hierarchical fair queuing,
token buckets, and priorities. However, it cannot simulate
the sorting required for say WFQ.
In our chip, each stage contains over 200 action units:
one for each field in the packet header vector. Over 7,000
action units are contained in the chip, but these consume a
small area in comparison to memory (< 10%). The action
unit processors are simple, specifically architected to avoid
costly to implement instructions, and require less than 100
gates per bit.
How should such an RMT architecture be configured?
Two pieces of information are required: a parse graph that
expresses permissible header sequences, and a table flow
graph that expresses the set of match tables and the control flow between them (see Figure 2 and §4.4). Ideally, a
compiler performs the mapping from these graphs to the
appropriate switch configuration. We have not yet designed
such a compiler.
3.
EXAMPLE USE CASES
To give a high level sense of how to use an RMT chip, we
will take a look at two use cases.
Example 1: L2/L3 switch. First, we need to configure
the parser, the match tables and the action tables. For our
first example, Figure 2a shows the parse graph, table flow
graph, and memory allocation for our L2/L3 switch. The
Parse Graph and Table Flow Graph tell the parser to extract
and place four fields (Ethertype, IP DA, L2 SA, L2 DA)
on the wide header bus. The Table Flow Graph tells us
which fields should be read from the wide header bus and
matched in the tables. The Memory Allocation tells us how
the four logical tables are mapped to the physical memory
stages. In our example, the Ethertype table naturally falls
into Stage 1, with the remaining three tables spread across
all physical stages to maximize their size. Most hash table
RAM is split between L2 SA and DA, with 1.2 million entries
for each. We devote the TCAM entries in all 32 stages to
hold 1 million IP DA prefixes. Finally, we need to store the
VLIW action primitives to be executed following a match
(e.g. egress port(s), decrement TTL, rewrite L2 SA/DA).
These require 30% of the stage’s RAM memory, leaving the
rest for L2 SA/DA. If enabled, packet and byte counters
would also consume RAM, halving the L2 table sizes.
Once configured, the control plane can start populating
each of the tables, for example by adding IP DA forwarding
entries.
Example 2: RCP and ACL support. Our second use
case adds Rate Control Protocol (RCP) support [8] and an
ACL for simple firewalling. RCP minimizes flow completion
times by having switches explicitly indicate the fair-share
rate to flows, thereby avoiding the need to use TCP slowstart. The fair-share rate is stamped into an RCP header
by each switch. Figure 2b shows the new parse graph, table
flow graph and memory allocation.
To support RCP, the packet’s current rate and estimated
RTT fields are extracted and placed in the header vector
by the parser. An egress RCP table in stage 32 updates
the RCP rate in outgoing packets—the min action selects
the smaller of the packet’s current rate and the link’s fairshare rate. (Fair-share rates are calculated periodically by
the control plane.)
A stateful table (§4.4) accumulates data required to calculate the fair-share rate. The stateful table is instantiated in
stage 32 and accumulates the byte and RTT sums for each
destination port.
We also create 20K ACL entries (120b wide) from the last
two stages of TCAM, reducing the L3 table to 960K prefixes,
along with RAM entries to hold the associated actions (e.g.,
drop, log).
In practice, the user should not be concerned with the
low-level configuration details, and would rely on a compiler
to generate the switch configuration from the parse graph
and table flow graph.
4.
CHIP DESIGN
Thus far, we have used a logical abstraction of an RMT
forwarding plane which is convenient for network users, and
a physical architecture to realize RMT. We now describe
implementation design details.
We chose a 1GHz operating frequency for the switch chip
because at 64 ports × 10 Gb/s and an aggregate throughput
of 960M packets/s, a single pipeline can process all input
port data, serving all ports, whereas at lower frequencies
we would have to use multiple such pipelines, at additional
area expense. A block diagram of the switch IC is shown in
Figure 3. Note that this closely resembles the RMT archi-
Parse Graph
Table Flow Graph
Ethernet
Memory Allocation
Legend
Tables
{Ethertype}
Ethertype
{Src Port, Src MAC}
RAM
END
IPv4
IP route
Src MAC
Action: Set src/dst MAC,
decrement IP TTL
Dst MAC
Action: Send
to controller
{Src/Dst IP,
IP Proto,
Src/Dst Port}
{RCP}
Action: Set
output port
TCAM
END
{Dst MAC}
{Dst IP}
Stage:
1
2
…
Table Flow Graph
Logical flow
Forward to buffer
Drop packet
32
(a) L2/L3 switch.
Parse Graph
Table Flow Graph
Ethernet
Memory Allocation
Ethertype
IP route
UDP
Action: Send
to controller
Dst MAC
Action: Set output port,
insert OMPLS header (opt.)
Action: Set
queue ID
ACL
TCAM
TCP
Src MAC
Action: Set src/dst MAC, decrement
IP TTL, insert OMPLS header (opt.),
set src/dst IP (opt.)
RCP
RAM
IPv4
RCP
Action: Clear
output port
Action: Update
RCP rate
Stage:
1
…
2
30
31
32
(b) RCP and ACL support.
Figure 2: Switch configuration examples.
queues
Ingress
Parsers
Match
Stage
1
Match
Stage
...
packet
pointer
(dequeue)
Ingress
Deparser
Egress processing
Egress
Parsers
32
Output Ch. 1
Match
Stage
1
Match
Stage
...
32
Egress
Deparser
...
...
Input Ch. 64
...
Input Ch. 1
...
packet
pointer
(enqueue)
Ingress processing
Output Ch. 64
Common data buffer packet
packet
data
data
Figure 3: Switch chip architecture.
tectural diagram of Figure 1a.
Input signals are received by 64 channels of 10Gb SerDes
(serializer-deserializer) IO modules. 40G channels are made
by ganging together groups of four 10G ports. After passing through modules which perform low level signalling and
MAC functions like CRC generation/checking, input data is
processed by the parsers. We use 16 ingress parser blocks
instead of the single logical parser shown in Figure 1a because our programmable parser design can handle 40Gb of
bandwidth, either four 10G channels or a single 40G one.
Parsers accept packets where individual fields are in variable locations, and output a fixed 4 kb packet header vector,
where each parsed field is assigned a fixed location. The location is static, but configurable. Multiple copies of fields
(for example, multiple MPLS tags or inner and outer IP
fields) are assigned unique locations in the packet header
vector.
The input parser results are multiplexed into a single
stream to feed the match pipeline, consisting of 32 sequential match stages. A large shared buffer provides storage to
accommodate queuing delays due to output port oversubscription; storage is allocated to channels as required. Deparsers recombine data from the packet header vector back
into each packet before storage in the common data buffer.
A queuing system is associated with the common data
buffer. The data buffer stores packet data, while pointers
to that data are kept in 2K queues per port. Each channel
in turn requests data from the common data buffer using
a configurable queuing policy. Next is an egress parser, an
egress match pipeline consisting of 32 match stages, and a
deparser, after which packet data is directed to the appropriate output port and driven off chip by 64 SerDes output
channels.
While a separate 32-stage egress processing pipeline seems
like overkill, we show that egress and ingress pipelines share
the same match tables so the costs are minimal. Further,
egress processing allows a multicast packet to be customized
(say for its congestion bit or MAC destination) by port without storing several different packet copies in the buffer. We
now describe each of the major components in the design.
4.1
Configurable Parser
The parser accepts the incoming packet data and produces the 4K bit packet header vector as its output. Parsing
is directed by a user-supplied parse graph (e.g., Figure 2),
converted by an offline algorithm into entries in a 256 entry
× 40b TCAM, which matches 32b of incoming packet data
and 8b of parser state. Note that this parser TCAM is completely separate from the match TCAMs used in each stage.
When, for example, the 16-bits of the Ethertype arrives, the
CAM matches on 16 of say 32 arriving bits (wildcarding the
rest), updating state indicating the next header type (e.g.,
VLAN or IP) to direct further parsing.
More generally, the result of a TCAM match triggers an
action, which updates the parser state, shifts the incoming
data a specified number of bytes, and directs the outputting
of one or more fields from positions in the input packet to
fixed positions in the packet header vector. This loop repeats to parse each packet, as shown in Figure 4. The loop
was optimized by pulling critical update data, such as input shift count and next parser state, out of the RAM into
TCAM output prioritization logic. The parser’s single cycle loop matches fields at 32 Gb/s, translating to a much
higher throughput since not all fields need matching by the
parser. A 40 Gb/s packet stream is easily supported by a
single parser instance.
Parser
Header
Identification
State &
header
data
TCAM
Field
Extraction
Next
state
Match index
Field
locations
Fields
Result
Header data
Packet Header
Vector
To
Match
Engine
Action
RAM
Each match stage has 106 RAM blocks of 1K entries ×
112b. The fraction of the RAM blocks assigned to match,
action, and statistics memory is configurable. Exact match
tables are implemented as Cuckoo hash tables [10, 26, 32]
with (at least) four ways of 1K entries each, each way requiring one RAM block. Reads are deterministically performed
in one cycle, with all ways accessed in parallel. Each match
stage also has 16 TCAM blocks of 2K entries × 40b that
can be combined to make wider or deeper tables.
Associated with each match table RAM entry is a pointer
to action memory and an action size, a pointer to instruction memory, and a next table address. The action memory
contains the arguments (e.g., next hop information to be
used in output encapsulations), and the instructions specify the function to be performed (e.g., Add Header). Action
memory, like the match memory, is made of 8 narrower units
consisting of 1K words × 112 bits, yielding 96b data each
(along with field valid bits and memory ECC—error correcting code—bits). Action memory is allocated from the 106
RAM blocks, while action instructions are held in a separate
dedicated memory.
As in OpenFlow, our chip stores packet and byte statistics counters for each flow table entry. Full 64b versions of
these counters are contained in off-chip DRAM, with limited resolution counters on-chip using the 1K RAM blocks
and applying the LR(T) algorithm [33] to provide acceptable
DRAM update rates. One word of statistics memory configurably holds the counters for either two or three flow entries,
allowing tradeoffs of statistics memory cost vs DRAM update rate. Each counter increment requires a read and write
memory operation, but in the 1GHz pipeline only one operation is available per packet, so a second memory port is
synthesized by adding one memory bank1 .
4.3
Figure 4: Programmable parser model.
4.2
Configurable Match Memories
Each match stage contains two 640b wide match units, a
TCAM for ternary matches, and an SRAM based hash table for exact matches. The bitwidth of the SRAM hash unit
is aggregated from eight 80b subunits, while the ternary
table is composed of 16 40b TCAM subunits. These subunits can be run separately, in groups for wider widths, or
ganged together into deeper tables. An input crossbar supplies match data to each subunit, which selects fields from
the 4Kb packet header vector. As described earlier (Figure 1b), tables in adjacent match stages can be combined to
make larger tables. In the limit, all 32 stages can create a
single table.
Further, the ingress and egress match pipelines of Figure 3 are actually the same physical block, shared at a fine
grain between the ingress and egress threads as shown in Figure 1b. To make this work, first the packet header vector is
shared between input and output vectors; each field in the
vector is configured as being owned by either the ingress or
egress thread. Second, the corresponding function units for
each field are allocated in the same way to ingress or egress.
Lastly, each memory block is allocated to either ingress or
egress. No contention issues arise since each field and memory block is owned exclusively by either egress or ingress.
Configurable Action Engine
A separate processing unit is provided for each packet
header field (see Figure 1c), so that all may be modified
concurrently. There are 64, 96, and 64 words of 8, 16, and
32b respectively in the packet header vector, with an associated valid bit for each. The units of smaller words can
be combined to execute a larger field instruction, e.g., two
8b units can merge to operate on their data as a single 16b
field. Each VLIW contains individual instruction fields for
each field word.
OpenFlow specifies simple actions, such as setting a field
to a value, and complex operations, such as PBB encapsulate
and inner-to-outer or outer-to-inner TTL copies, where the
outer and inner fields may be one of a number of choices.
Complex modifications can be subroutines at low speeds, but
must be flattened into single-cycle operations at our 1 GHz
clock rate using a carefully chosen instruction set.
Table 1 is a subset of our action instructions. Depositbyte enables depositing an arbitrary field from anywhere in
a source word to anywhere in a background word. Rotmask-merge independently byte rotates two sources, then
merges them according to a byte mask, useful in performing
IPv6 to IPv4 address translation [18]. Bitmasked-set is useful for selective metadata updates; it requires three sources:
the two sources to be merged and a bit mask. Move, like
other operators, will only move a source to a destination
if the source is valid, i.e., if that field exists in the packet.
1
S. Iyer. Memoir Systems. Private communication, Dec.
2010.
Time
Another generic optional conditionalization is destination
valid. The cond-move and cond-mux instructions are useful for inner-to-outer and outer-to-inner field copies, where
inner and outer fields are packet dependent. For example,
an inner-to-outer TTL copy to an MPLS tag may take the
TTL from an inner MPLS tag if it exists, or else from the IP
header. Shift, rotate, and field length values generally come
from the instruction. One source operand selects fields from
the packet header vector, while the second source selects
from either the packet header vector or the action word.
Category
logical
shadd/sub
arith
deposit-byte
rot-mask-merge
bitmasked-set
move
cond-move
cond-mux
Description
and, or, xor, not, . . .
signed or unsigned shift
inc, dec, min, max
any length, source & dest offset
IPv4 ↔ IPv6 translation uses
S1 &S2 | S1 &S3 ; metadata uses
if VS1 S1 → D
if VS2 &VS1 S1 → D
if VS2 S2 → D else if VS1 S1 → D
Table 1: Partial action instruction set.
(Si means source i; Vx means x is valid.)
A complex action, such as PBB, GRE, or VXLAN encapsulation, can be compiled into a single VLIW instruction and
thereafter considered a primitive. The flexible data plane
processing allows operations which would otherwise require
implementation with network processors, FPGAs, or software, at much higher cost and power at 640Gb/s.
4.4
Other Features
Reducing Latency by Dependency Analysis: One
can easily ensure correctness by requiring that physical match
stage I processes a packet header vector P only after stage
I − 1 completely finishes processing P . But this is overkill
in many cases and can severely increase latency. Key to
reducing latency is to identify three types of dependencies
between match tables in successive stages: match dependencies, action dependencies and successor dependencies [2],
each described below.
Match dependencies occur when a match stage modifies
a packet header field and a subsequent stage matches upon
that field. No overlap of execution is possible in this case
(Figure 5a). Action dependencies occur when a match stage
modifies a packet header field and a subsequent stage uses
that field as an input to an action, for example, if one stage
sets a TTL field and the next stage decrements the TTL.
Partial execution overlap is possible (Figure 5b). Successor
dependencies occur when the execution of a match stage is
predicated on the result of execution of a prior stage; the
prior stage execution may cause the successor stage to be
skipped. The successor stage can be run concurrently with
its predecessor (Figure 5c) if it is run speculatively and predication is resolved before side effects are committed. Tables
with no dependencies between them can also be run concurrently. The pipeline delays between successive stages are
statically configurable between the three options of Figure 5,
individually for the ingress and egress threads.
Multicast and ECMP: Multicast processing is split between ingress and egress. Ingress processing writes an output port bit vector field to specify outputs, and optionally,
Stage 1
Stage 2
Match Action
Match Action
(a) Match dependency.
Stage 1
Stage 2
Match Action
Match Action
(b) Action dependency.
Stage 1
Stage 2
Match Action
Match Action
(c) No dependency or successor dependency.
Figure 5: Match stage dependencies.
a tag for later matching and the number of copies routed to
each port. A single copy of each multicast packet is stored in
the data buffer, with multiple pointers placed in the queues.
Copies are created when the packet is injected into the egress
pipeline, where tables may match on the tag, the output
port, and a packet copy count to allow per-port modifications. ECMP and uECMP processing is similar.
Meters and Stateful Tables: Meters measure and classify flow rates of matching table entries, and can be used to
modify or drop packets. Meters are but one example of stateful tables, where an action modifies state that is visible to
subsequent packets and can be used to modify them. Stateful counters that can be arbitrarily incremented and reset.
They can be used to implement, for example, GRE sequence
numbers (that are incremented in each encapsulated packet)
and OAM [15, 25].
Consistent and atomic updates: To allow consistent
updates [34], version information is contained in table entries, and a version ID flows through the pipeline with each
packet, qualifying table matches by version compatibility.
5.
EVALUATION
We characterize the cost of configurability in terms of the
increased area and power of our design relative to a conventional less programmable switch chip. Our comparison
culminates in a comparison of total chip area and power in
Section 5.5. To get there, we consider the contributors to
cost by considering the parser, the match stages and the
action processing in turn.
5.1
Programmable Parser Costs
Programmability comes at a cost. A conventional parser
is optimized for one parse graph, whereas a programmable
parser must handle any supported parse graph. Cost is evaluated by comparing synthesis results for conventional and
programmable designs. Total gate count is shown in Figure 6 for conventional parsers implementing several parse
graphs and a programmable parser. We assume parser aggregate throughput of 640 Gb/s by combining 16 instances
of a 40 Gb/s parser running at 1 GHz. The result module
in all designs contains the 4Kb packet header vector during parsing. The programmable parser uses a 256 × 40 bit
TCAM and a 256 × 128 bit action RAM.
Parser gate count is dominated by logic for populating the
parser header vector. The conventional design requires 2.9–
3.0 million gates, depending upon the parse graph, while the
programmable design requires 5.6 million gates, of which 1.6
Gates (×106)
Result
Hdr. Ident./Field Extract.
Action RAM
TCAM
Co
n
ve
nt
Si ion
m al:
pl
Co
e
nv
En ent
te ion
rp a
ris l:
Co
e
nv
Co en
re tio
ro na
ut l:
Co
er
D nve
at nt
a io
ce na
nt l:
C
Se o
er
rv nv
ice en
t
pr io
ov na
Co ider l:
nv
Co en
m tio
po na
sit l:
e
Pr
og
ra
m
m
ab
le
6
5
4
3
2
1
0
Figure 6: Total gate count of parsers providing 640 Gb/s
aggregate throughput.
million is contributed by the added TCAM and action RAM
modules. From these results, the cost of parser programmability is approximately 2 (5.6/3.0 = 1.87 ≈ 2).
Despite doubling the parser gate count, the parser accounts for less than 1% of the chip area, so the cost of making the parser programmable is not a concern.
5.2
Memory Costs
There are several costs to memories the reader may be
concerned about. First, there is the cost of the memory technology itself (hash table, TCAMs) versus standard SRAM
memory and the cost of breaking up the memory into smaller
blocks which can be reconfigured; second, there is the cost of
additional data needed in each match table entry to specify
actions and keep statistics; and third, there is the cost of internal fragmentation such as when an Ethernet Destination
address of 48 bits is placed in a 112-bit wide memory. We
treat each overhead in turn and specifically point out the
(small) additional cost for programmability. In what follows we refer to an entry in a match table (such as a 48-bit
Ethernet DA) as a flow entry.
5.2.1
Memory Technology Costs
Exact matching: We use cuckoo hashing for exact matching because its fill algorithm provides high occupancy, typically above 95% for 4-way hashtables. Cuckoo hashtables
resolve fill conflicts by recursively evicting conflicting entries
to other locations. Additionally, while our memory system is
built up out of 1K by 112 bit RAM blocks for configurability,
one might expect an area penalty vs using larger, more efficient memory units. However, using 1K RAM blocks incurs
an area penalty of only about 14% relative to the densest
SRAM modules available for this technology.
Wildcard matching: We use large amounts of TCAM
on chip to directly support wildcard matching such as prefix
matching and ACLs. TCAM is traditionally thought to be
infeasible due to power and area concerns. However, TCAM
operating power has been reduced by about 5× by newer
TCAM circuit design techniques [1]. Thus, in the worst
case, at the maximum packet rate with minimum packet size
on all channels, TCAM power is one of a handful of major
contributors to total chip power; at more typical mixtures
of long and short packets, TCAM power reduces to a small
percentage of the total.
Next, while a TCAM typically has an area 6–7× that
of an equivalent bitcount SRAM, both ternary and binary
flow entries have other bits associated with them, including
action memory, statistics counters, and instruction, action
data, and next table pointers. For example, with 32 bits of
IP prefix, 48 bits of statistics counter, and 16 bits of action
memory (say for specifying next hops), the TCAM portion
is only 1/3 of the memory bitcount and so the area penalty
for TCAM drops to around 3×.
While 3× is significant, given that 32b (IPv4) or 128b
(IPv6) longest prefix matching and ACLs are major use
cases in all existing routers, devoting significant resources
to TCAM to allow say 1M IPv4 prefixes or 300K ACLs
seems useful. While we could have used just SRAM with
special purpose LPM algorithms instead as in [6], achieving
the single-cycle latency of TCAMs for a 32 or 128 bit LPM
is difficult or impossible. Nevertheless, deciding the ratio
of ternary to binary table capacity (our chip proposes a 1:2
ratio) is an important implementation decision with significant cost implications, for which currently there is little real
world feedback.
5.2.2
Costs of Action Specification
Besides the flow entry, each match table RAM entry also
has a pointer to action memory (13b), an action size (5b),
a pointer to instruction memory (5b for 32 instructions),
and a next table address (9b). These extra bits represent
approximately 35% overhead for the narrowest flow entries.
There are also bits for version and error correction but these
are common to any match table design so we ignore them.
In addition to overhead bits in a flow entry, other memories are required for storing actions and statistics. These
add to the total overhead—the ratio of total bits required to
just the match field bits— but both of these extra costs can
sometimes be reduced. We will show how in some cases it
is possible to reduce flow entry overhead bits. Furthermore,
applications require varying amounts of action memory, and
sometimes statistics are not needed, so these memory costs
can be reduced or eliminated.
Given the variable configuration of memory blocks between match, action, and statistics, we use a few configuration examples to see how the bookkeeping overhead varies
compared to non-configurable fixed allocations.
In the first configuration, shown in Figure 7a and Table 2,
32 memory blocks are used for match memory in a stage,
implementing 32K 80b wide exact match flow entries. Another 16K 80b ternary entries are in the TCAM modules.
All flow entries have action entries of the same size, requiring 48 memories for actions. Statistics consume 24 memory
banks, along with a spare bank for multiporting the statistics
memory. An approximately equal portion of action memory
is allocated for each match memory, and might be considered
a base case with a minimum amount of flow table capacity.
Excluding the 24 banks used for ternary actions and statistics in case a, 40% of the banks used for binary operations
are match tables, indicating a 2.5× overhead. Compounding this with the 35% bit overhead in the match tables, the
total binary overhead is 3.375×, the ratio of total bitcount
to match data bitcount. In other words, only a third of the
RAM bits can be used for flow entries.
Cases a2 and a3 of Table 2 change the match width to
160 and 320 bits respectively, reducing action and statistics
requirements, and yielding increased match capacity.
Configuration b of Table 2 further increases the binary
a)
: Binary match
: Stats
b)
c)
: Binary action
: Stats or binary match
: Ternary action
: Unused
Figure 7: Match stage unit memory map examples.
and ternary flow table match widths to 640 bits, as shown in
Figure 7b, (memory width is shown horizontally), reducing
the number of flow entries by a factor of 8 compared to
the base case, along with the required action and statistics
capacity. While such a wide match may be rare (say for
an entire header match), we see that with 8× wider flow
entries than the base case above, 80 banks can be used for
exact match, 75% of memory capacity, 2.5× higher table
capacity than the base case.
Configuration c1 of table 2, shown in Figure 7c, takes advantage of a use case where the number of individual actions
is limited. For example, a data center address virtualization application requires a large number of flow entries, but
match entries may point to one of only 1000 possible destination top of rack switches. Thus 4K of action memory would
suffice. That would allow 62 memory blocks for match and
40 for statistics, almost doubling the number of exact match
entries from the base case. If statistics were not desired, 102
memory blocks could be used for match as shown in table
entry c2 , 96% of the total memory capacity.
Case
a1
a2
a3
b
c1
c2
MatchWidth
80
160
320
640
80
80
Match
32
52
72
80
62
102
Action
48
34
22
12
4
4
Stats
25
18
12
7
40
0
Relative
1.000×
1.625×
2.250×
2.500×
1.900×
3.250×
Table 2: Memory unit allocation and relative exact match
capacity.
In short, flow entry density is greatly increased if actions
or statistics are reduced or eliminated. If the user of the
chip deems these complex actions and statistics necessary,
then it is unfair to blame the chip configurability options
for this overhead. The only fundamental book-keeping costs
that can be directly attributed to programmability are the
instruction pointer (5b) and the next table address (9b),
which is around 15%.
Costs can be further reduced in tables with fixed behaviors. A fixed-function table uses the same instruction and
next-table pointers for all entries, e.g., L2 dest MAC table;
static values may be configured for these attributes, allowing
14 bits of instruction and next-table pointer to be reclaimed
for match. More generally, a configurable width field in the
flow entry can optionally provide LSBs for action, instruction, or next-table addresses, allowing a reduced number of
different instructions or actions, or addressing a small array
for next table, while reclaiming as many instruction, nexttable, and action address bits as possible given the complexity of the function.
Next, tables can provide an action value as an immediate
constant rather than as a pointer to action memory for small
constants, saving pointer and action memory bits.
A simple mechanism enables these optimizations: match
table field field boundaries can be flexibly configured, allowing a range of table configurations with arbitrary sizes
for each field, subject to a total bitwidth constraint. Tables
with fixed or almost fixed functions can be efficiently implemented with almost no penalty compared to their fixed
counterparts.
5.2.3
Crossbar Costs
A crossbar within each stage selects the match table inputs from the header vector. A total of 1280 output bits
(640b for each of TCAM and hash table) are selected from
the 4Kb input vector. Each output bit is driven by a 224
input multiplexor, made from a binary tree of and-or-invert
AOI22 gates (with the logic function AB + CD), and costing one 0.65µm2 per mux input. Total crossbar area is
1280 × 224 × 0.65µm2 × 32 stages ≈ 6 mm2 . Area computation for the action unit data input muxes is similar.
5.3
Fragmentation Costs
A final overhead is internal fragmentation or packing costs.
Clearly, a 48-bit Ethernet Destination Address placed in
an 112 b wide memory wastes more than half the memory. In comparison, a fixed-function Ethernet bridge contains custom 48-bit wide RAM. Thus this cost is squarely
attributable to programmability and our choice of 112b wide
RAMs. One could reduce this overhead for Ethernet by
choosing 48b as the base RAM width, but how can a chip
designed for general purpose use (and future protocols) predict future match identifier widths?
Fortunately, even this overhead is reduced through another architectural trick that allows sets of flow entries to
be packed together without impairing the match function.
For example, the standard TCP 5-tuple is 104 bits wide.
Three of these entries can be packed into four memory units
of width 448 b, rather than separately requiring each to
consume two memory units. Or, with low entry overhead
equivalent to a simple pass/fail, 4 of these can be packed
into 4 words, due to amortization of ECC bits over wider
data. Fundamentally, this is possible efficiently because under the covers, longer matches are constructed from trees
of smaller 8-bit matches; flexibility only slightly complicates
this logic.
The combination of these two techniques, variable data
packing into a data word (to reduce action specification
costs) and variable flow entry packing into multiple data
words (to reduce fragmentation costs), assures efficient memory utilization over a wide range of configurations. In summary, while conventional switches have highly efficient implementations for specific configurations of tables, this architecture can approach that efficiency, not only for those
specific table configurations, but for a wide range of other
configurations as well.
5.4
Costs of Action Programmability
Besides the costs of the instruction RAMs and action
memories, there are around 7000 processor datapaths ranging in width from 8 to 32 bits. Fortunately, because they use
a simple RISC instruction set, their combined area consumes
only 7% of the chip.
5.5
Area and Power Costs
This switch is designed with large match table capacity,
so match and action memories contribute substantially to
chip area estimates as shown in Table 3. The first item,
which includes IO, data buffer, CPU, etc., occupies a similar
area in conventional switches. As can be seen, the VLIW
action engine and parser/deparser contributions to area are
relatively small.
We suggested earlier that the match stage unit RAMs
suffered a 14% area penalty compared to the best possible
RAMs. Given this penalty to the match stage SRAM (not
TCAM) area, and some allowance for additional bitcount vs
conventional switches (15%), excess memory area is about
8% of the chip total. If excess logic in the parser and action
engine add another 6.2%, a 14.2% area cost results, justifying the earlier claim of a less than 15% cost differential.
Section
Area
IO, buffer, queue, CPU, etc 37.0%
Match memory & logic
54.3%
VLIW action engine
7.4%
Parser + deparser
1.3%
Total extra cost:
Cost
0.0%
8.0%
5.5%
0.7%
14.2%
Table 3: Estimated chip area profile.
Section
Power
I/O
26.0%
Memory leakage 43.7%
Logic leakage
7.3%
RAM active
2.7%
TCAM active
3.5%
Logic active
16.8%
Total extra cost:
Cost
0.0%
4.0%
2.5%
0.4%
0.0%
5.5%
12.4%
Table 4: Estimated chip power profile.
Estimated switch power is detailed in Table 4 under worst
case operating conditions (temperature, chip process), 100%
traffic with a mix of half min and half max (1.5 kB) size
packets, and all match and action tables filled to capacity.
Input/output power is equivalent to a conventional switch.
Memory leakage power is proportional to memory bitcount,
so if this programmable switch can be implemented with
equivalent bitcount to a conventional switch, power will be
comparable. The remaining items, totalling 30%, will be
less in a conventional switch because of the reduced functionality in its match-action pipeline. We estimate that our
programmable chip dissipates 12.4% more power than a conventional switch, but is performing much more substantial
packet manipulation.
The overall competitive evaluation with conventional switches suggests that equivalent functions can be performed
by this switch with equivalent memory bitcounts. This in
turn drives parity in dominant aspects of chip cost and
power. The additional power and area costs borne by the
programmable solution are quite small given the more comprehensive functionality of the switch.
6.
RELATED WORK
Flexible processing is achievable via many mechanisms.
Software running on a processor is a common choice. Our
design performance exceeds that of CPUs by two orders of
magnitude [7], and GPUs and NPUs by one order [5,9,12,29].
Modern FPGAs, such as the Xilinx Virtex-7 [35], can forward traffic at nearly 1 Tb/s. Unfortunately, FPGAs offer
lower total memory capacity, simulate TCAMs poorly, consume more power, and are significantly more expensive. The
largest Virtex-7 device available today, the Virtex-7 690T,
offers 62Mb of total memory which is roughly 10% of our
chip capacity. The TCAMs from just two match stages
would consume the majority of lookup-up tables (LUTs)
that are used to implement user-logic. The volume list price
exceeds $10,000, which is an order of magnitude above the
expected price of our chip. These factors together rule out
FPGAs as a solution.
Related to NPUs is PLUG [6], which provides a number of
general processing cores, paired with memories and routing
resources. Processing is decomposed into data flow graphs,
and the flow graph is distributed across the chip. PLUG
focuses mainly on implementing lookups, and not on parsing
or packet editing.
The Intel FM6000 64 port × 10Gb/s switch chip [24] contains a programmable parser built from 32 stages with a
TCAM inside each stage. It also includes a two-stage matchaction engine, with each stage containing 12 blocks of 1K ×
36b TCAM. This represents a small fraction of total table
capacity, with other tables in a fixed pipeline.
The latest OpenFlow [31] specification provides an MMT
abstraction and partly implements an RMT model. But its
action capability is still limited, and it is not certain that a
standard for functionally complete actions is on the way or
even possible.
7.
CONCLUSIONS
Ideally, a switch or router should last for many years.
Dealing with a changing world requires programmability that
allows software upgrades to add new functions and new protocols in the field. Network processors (NPUs) were introduced to support this vision, but neither NPUs or GPUs
have come close to achieving the speeds of fixed function
switches using ASICs; nor have we seen a case study of reprogramming an NPU-based router such as Cisco’s CRS-1 to
add a new protocol. Likewise FPGAs, which only recently
approached ASIC forwarding speeds, remain prohibitively
expensive.
Our chip design resurrects this ancient vision of programmability, expressed in the RMT model, within the constraints
of what is possible on a real chip. New fields can be added,
lookup tables can be reconfigured, new header processing
added, all through software reconfiguration. While our chip
cannot do regular expressions, or manipulate packet bodies,
a box built from this chip could metamorphose from an Ethernet chip on Tuesday to a firewall on Wednesday and to a
completely new device on Thursday, all by the right software
upgrades. The challenge is to do this today at a capacity
approaching a terabit. The chip design we propose has surprising specs: it contains 7,000 processor datapaths, 370 Mb
of SRAM, and 40 Mb of TCAM, across 32 processing stages.
In terms of ideas, we single out the RMT model as a powerful way to map the programmer’s desired forwarding behavior onto a pipeline built from a flexible parser, a configurable arrangement of logical match stages with memories
of arbitrary width and depth, and flexible packet editing.
These abstractions require new algorithms to be efficiently
implemented at terabit speeds. Our use of memory blocks
that can be ganged within or across stages is key to realizing
the vision of reconfigurable match tables; our large scale use
of TCAM greatly increases matching flexibility; and finally,
our use of fully-parallel VLIW instructions is key to packet
editing. Our design suggests that this greatly increased flexibility comes at an additional cost of less than 15% in area
and power consumption of the chip. Ultimately, the implementation challenge which we have addressed that may
not be apparent to SIGCOMM audiences is producing an
architecture which is possible to wire efficiently on chip.
While the OpenFlow specification hints at RMT and several researchers [17] have actively pursued this dream, RMT
models remained theoretical without an existence proof of
a chip design that works at terabit speeds. Our paper contributes a concrete RMT proposal and a proof of its feasibility. Clearly it is possible to go further and remove some
of the restrictions we have imposed for implementation reasons. But now the conversation can begin.
Acknowledgements
We would like to thank Nick Feamster for feedback on an
early draft, Aditya Akella for shepherding the final version,
and the anonymous reviewers for their useful comments.
8.
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