Efficient Per-Flow Queueing in DRAM at OC-192 Line - carv-ics

Efficient Per-Flow Queueing in DRAM at OC-192 Line - carv-ics
Efficient Per-Flow Queueing in DRAM at OC-192
Line Rate using Out-of-Order Execution Techniques
Aristides A. Nikologiannis
Institute of Computer Science (ICS)
Foundation of Research & Technology – Hellas (FORTH)
P.O. Box 1385
Heraklio, Crete, GR-711-10 GREECE
Tel: +30 81 391 655, fax: +30 81 391 661
anikol@ics.forth.gr, anikol@ellemedia.com
Technical Report 279
Abstract
The explosive growth of Internet traffic has created an acute demand for networks of
ever increasing throughput. Besides raw throughput, modern mutlimedia
applications also demand Quality of Service (QoS) guarantees. Both of these
requirements result in a new generation of switches and routers, which use
specialized hardware to support high speeds and advanced QoS.
This thesis studies one of the subsystems of such switches/routers, namely
queue management in the ingress and egress line cards at OC-192 (10 Gbps) line
rate. The provision of QoS guarantee usually requires flow isolation, which is often
achieved using per-flow queueing. The implementation of a queue manager,
supporting thousands of flows and operating at such high speed, is challenging. We
study thoroughly this issue and we show that the queue manager implementation is
feasible by using advanced hardware techniques, similar to those employed in the
supercomputers of the 60's and in modern microprocessors. We use DRAM
technology for buffer memory in order to provide large buffer space. To effectively
deal with bank conflicts in the DRAM buffer, we have to use multiple pipelined
control processes, out-of-order execution and operand renaming techniques. To
avoid off-chip SRAM, we maintain the queue management pointers in the buffer
memory, using free buffer preallocation and free list bypassing. We have described
our architecture using behavioral Verilog (Hardware Description Language), at a
clock cycle accurate level, assuming Rambus DRAM, and we have partially verified
it by short simulation runs.
Table of Contents
1
Introduction..................................................................................... 7
1.1 Motivation ..............................................................................................................................7
1.1 Switch Evolution....................................................................................................................8
1.1.1
Switch Generations ........................................................................................................8
1.1.2
Switching Fabric Topologies .......................................................................................10
1.1.3
Queueing Architectures & Performance ......................................................................11
1.2 Related Work on High-Speed Switches .............................................................................14
1.3 Thesis Contribution.............................................................................................................15
1.4 Thesis Organization.............................................................................................................16
2
Ingress/Egress Interface Module Architecture .......................... 17
2.1 Ingress/Egress Module Functionality ................................................................................17
2.1.1
Ingress Module main Functions ...................................................................................17
2.1.2
Egress Module main Functions....................................................................................18
2.2 Ingress Module Chip Partitioning......................................................................................18
2.3 Datapath and Queue Management Chip ...........................................................................20
2.4 Header Protocol Processing Chip ......................................................................................20
2.4.1
Flow Classification ......................................................................................................20
2.4.2
Short Label Forwarding 1: ATM .................................................................................21
2.4.3
Short Label Forwarding 2: IP over ATM.....................................................................22
2.4.4
Short Label Forwarding 3: MPLS................................................................................23
2.5 Scheduling-Policing Chip....................................................................................................24
2.5.1
Basic Disciplines on Scheduling..................................................................................24
2.5.2
Scheduling Best-Effort Connections............................................................................25
2.5.3
Scheduling Guaranteed-Service Connections ..............................................................26
2.5.4
Leaky Bucket ...............................................................................................................26
2.5.5
Calendar Queue............................................................................................................28
2.5.6
Heap Management .......................................................................................................28
2.5.7
An Advanced Scheduler Architecture..........................................................................28
3
Datapath & Queue Management Chip Architecture ................ 31
3.1 Queue Management Data Structures.................................................................................31
3.1.1
Fragmentation loss .......................................................................................................32
3.1.2
Queue Management Operations...................................................................................32
3.2 Buffer Memory Technology................................................................................................33
3.2.1
DRAM versus SRAM ..................................................................................................33
3.2.2
Rambus DRAM Technology .......................................................................................34
3.2.3
Out-of-Order DRAM Accesses....................................................................................35
3.3 Multi-Queue Management Architecture at High-Speed (10Gbps) .................................35
3.3.1
Queue Management Architecture Overview................................................................35
3.3.2
Why Pipelined Queue Manager ...................................................................................37
3.3.3
Why Multiple Control Processes .................................................................................37
3.4 Queue Management Pipeline Dependencies......................................................................41
3.4.1
Successive Enqueue and Dequeue Operations for the same flow................................42
3.4.2
Successive Enqueue Operations of packet segments ...................................................42
3.4.3
Buffer Memory Module Dependencies........................................................................43
3.5 Pipeline Dependencies Handling ........................................................................................43
3.5.1
Operands Renaming (Tomasulo) [15, chapter 4], [16] ................................................43
3.5.2
Applying Operand Renaming Techniques to the Queue Management Architecture ...43
3.6 Queue Pointer Management & Architecture Modifications............................................46
3.6.1
Next-Pointers in the DRAM Buffer Memory ..............................................................46
3.6.2
Buffer Preallocation technique [29].............................................................................46
3.6.3
Link Throughput Saturation.........................................................................................48
3.6.4
Free List Bypassing technique [29]..............................................................................48
3.6.5
Per-memory bank Queueing Free List Organization ...................................................49
3.6.6
Free Buffer Cache ........................................................................................................50
3.7 The Overall Queue Management Architecture.................................................................50
4
Queue Management Micro-Architecture ................................... 53
4.1 Hardware Implementation of the QM Data Structures...................................................53
4.1.1
Queue Table.................................................................................................................53
4.1.2
Pending Write Table and Transit Buffer......................................................................54
4.1.3
Pending Read Table.....................................................................................................55
4.1.4
Free List Table and Free List Cache ............................................................................56
4.1.5
Control and data Buffer ...............................................................................................56
4.2 The Pipelined Control Processes Micro-Architecture......................................................57
4.2.1
Packet Fetching Process Micro-Architecture...............................................................57
4.2.2
Enqueue Operation Issuing Process Micro-Architecture .............................................59
4.2.3
Enqueue Execution Process Micro-Architecture .........................................................61
4.2.4
Handling Exceptional Cases during an Enqueue Operation ........................................65
4.2.5
Dequeue Operation Issuing Process Micro-Architecture.............................................66
4.2.6
Dequeue Operation Execution Process Micro-Architecture ........................................66
4.2.7
Handling Exceptional Cases during a Dequeue Operation ..........................................69
4.2.8
Queue Manager Interface Process Micro-Architecture................................................69
4.2.9
Resource Conflicts among Queue Management Processes..........................................72
4.2.10 Search Engines Architecture........................................................................................73
4.2.11 Free List Organization Alternatives.............................................................................77
4.3 Rambus Memory Technology ............................................................................................79
4.3.1
Read and Write Operations in a Pipelined Fashion .....................................................80
4.3.2
Rambus Memory Device Architecture ........................................................................81
4.3.3
Rambus Memory Module Architecture .......................................................................82
4.3.4
Rambus Memory Interface ..........................................................................................82
4.3.5
Rambus Memory Controller ........................................................................................83
5
Verilog Description & Simulation ............................................... 87
5.1
5.2
6
7
Hardware Implementation Cost ........................................................................................87
Verification ..........................................................................................................................89
Conclusions and Open Topics...................................................... 91
Appendix A .................................................................................... 93
7.1 Flow Classification ..............................................................................................................93
7.1.1
Recursive Flow Classification (RFC) ..........................................................................93
7.1.2
Flow Classification by using Hashing functions..........................................................94
7.2 IP Routing Lookup..............................................................................................................95
7.2.1
Multi-stage IP routing by using Small SRAM Blocks.................................................95
7.2.2
Two-stage IP routing by using Large DRAM Blocks..................................................96
8
Appendix B .................................................................................... 97
8.1
9
Block Diagrams of Queue Management Processes...........................................................97
References .................................................................................... 103
List of Figures
Chapter 1
Figure 1. 1 First generation switch architecture.....................................................................................8
Figure 1. 2 Second generation switch architecture ................................................................................9
Figure 1. 3 Third generation switch architecture ...................................................................................9
Figure 1. 4 Crossbars ...........................................................................................................................10
Figure 1. 5 Banyan...............................................................................................................................10
Figure 1. 6 Benes .................................................................................................................................11
Figure 1. 7 Benes constructure ............................................................................................................11
Figure 1. 8 Batcher-Banyan .................................................................................................................11
Figure 1. 9 Output Queueing
Figure 1. 10 Input Queueing.............................................................12
Figure 1. 11 Head of Line Blocking ....................................................................................................13
Figure 1. 12 Advanced Input Queueing ...............................................................................................13
Figure 1. 13 Switches with internal speedup .......................................................................................13
Chapter 2
Figure 2. 1 Ingress Module Chip Partitioning......................................................................................19
Figure 2. 2 ATM Translation Table.....................................................................................................22
Figure 2. 3 Ipv4, Ipv6 packet format ...................................................................................................23
Figure 2. 4 MPLS Hierarchical Network.............................................................................................23
Figure 2. 5. Rate-controlled Scheduler ................................................................................................26
Figure 2. 6 Leaky Bucket.....................................................................................................................27
Figure 2. 7. The two leaky buckets traffic shaping mechanism...........................................................27
Figure 2. 8 Calendar queue structure ...................................................................................................28
Figure 2. 9 A two stage scheduler........................................................................................................29
Chapter 3
Figure 3. 1 Queue Manager Data Structures........................................................................................31
Figure 3. 2 Enqueue Operation ............................................................................................................33
Figure 3. 3 Dequeue Operation............................................................................................................33
Figure 3. 4 Buffer Memory Throughput ..............................................................................................34
Figure 3. 5 Non-Interleaved versus Interleaved Transaction ...............................................................35
Figure 3. 6 Multi-Queue Management Block Diagram .......................................................................36
Figure 3. 7 . Incoming segment entry process .....................................................................................39
Figure 3. 8 Enqueue Issuing Process ...................................................................................................39
Figure 3. 9 Enqueue Execution Process...............................................................................................40
Figure 3. 10 Queue Management Interface Process.............................................................................41
Figure 3. 11 Successive Enqueue Operations of packet segments.......................................................42
Figure 3. 12 Pending Lists ...................................................................................................................43
Figure 3. 13 Segment list per packet arrival ........................................................................................44
Figure 3. 14 Operand renaming technique for successive enqueue operations....................................45
Figure 3. 15 Per-flow pending lists......................................................................................................45
Figure 3. 16 Operand renaming technique for successive enqueue operations....................................46
Figure 3. 17 No free buffer preallocation ............................................................................................47
Figure 3. 18 Buffer preallocation.........................................................................................................47
Figure 3. 19 Read and Write transactions of an enqueue and a dequeue operation at the same time slot
....................................................................................................................................................48
Figure 3. 20 Free List Bypassing (memory transactions) ....................................................................49
Figure 3. 21 Mutli-Queue Rambus Controller block diagram .............................................................52
Chapter 4
Figure 4. 1 Queue Table ......................................................................................................................53
Figure 4. 2 Pending Write Table..........................................................................................................54
Figure 4. 3 Pending Read Table ..........................................................................................................55
Figure 4. 4 Free List Table and Free Buffer Cache .............................................................................56
Figure 4. 5 The Control Buffer format ................................................................................................57
Figure 4. 6 Packet fetching process block diagram (mode 1) ..............................................................58
Figure 4. 7 Packet fetching process block diagram (mode 2) ..............................................................59
Figure 4. 8 Enqueue issuing process datapath (not-pending state) ......................................................60
Figure 4. 9 Enqueue issuing process datapath (pending state).............................................................61
Figure 4. 10 Enqueue Execution process (first stage)..........................................................................62
Figure 4. 11 Free buffer extraction ......................................................................................................63
Figure 4. 12 Second stage (execute an enqueue operation) .................................................................64
Figure 4. 13 Dequeue Operation Issuing Process ................................................................................66
Figure 4. 14 First Stage........................................................................................................................67
Figure 4. 15 Second stage....................................................................................................................68
Figure 4. 16 Detection circuit ..............................................................................................................71
Figure 4. 17 Ingress Module Output Access Conflict..........................................................................71
Figure 4. 18 Queue Management Interface Process in Pipeline Fashion.............................................72
Figure 4. 19 The Search Engine Block Diagram .................................................................................75
Figure 4. 20 Priority Alternatives ........................................................................................................76
Figure 4. 21 T1 and T2 Priority Encoder Cells....................................................................................76
Figure 4. 22 Two-bit Priority Encoder Figure 4. 23 Eight-bit Priority Encoder..............................77
Figure 4. 24 The Modified Priority Encoder .......................................................................................77
Figure 4. 25 Bitmap Free List organization.........................................................................................78
Figure 4. 26 Free list organization as a linked list ...............................................................................78
Figure 4. 27 Per-bank queueing Free List Organization......................................................................79
Figure 4. 28 Rambus Technology........................................................................................................79
Figure 4. 29 Read Transaction.............................................................................................................80
Figure 4. 30 Write Transaction............................................................................................................80
Figure 4. 31 Interleaved read and write transactions ...........................................................................81
Figure 4. 32 Transaction Insertion FSM ..............................................................................................84
Figure 4. 33 Read Transaction time-diagram ......................................................................................85
Figure 4. 34 Write Transaction time-diagram......................................................................................85
Chapter 5
Figure 5. 1 Datapath chip memory requirements.................................................................................88
Figure 5. 2 hardware complexity of our architecture...........................................................................88
Appendix A
Figure 2b. 1 Recursive Flow Classification.........................................................................................93
Figure 2b. 2 Flow Classification by Hashing.......................................................................................94
Figure 2b. 3 Multi-stage IP routing .....................................................................................................95
Figure 2b. 4 Two-stage IP routing.......................................................................................................96
Appendix B
Figure 8. 1 Block Diagram of Packet Entry Process ...........................................................................97
Figure 8. 2 Block Diagram of Enqueue Issue Process........................................................................98
Figure 8. 3 Block Diagram of Enqueue Execution Process (first stage)..............................................98
Figure 8. 4 Block Diagram of Enqueue Execution Process (second stage) .........................................99
Figure 8. 5 Block Diagram of Dequeue Issue Process.........................................................................99
Figure 8. 6 Block Diagram of Dequeue Execution Process (first stage)............................................100
Figure 8. 7 Block Diagram of Dequeue Execution Process (second stage) .......................................100
1 Introduction
1.1 Motivation
Within the past few years, there has been a rapid growth in network traffic. New
applications, particularly multimedia applications, have placed increased demands
on speed and Quality of Service (QoS) guarantees of networks infrastructure. These
requirements are expressed using the following Quality of Service (QoS) related
parameters:
?? Bandwidth - the rate at which an application's traffic must be carried by the
network
?? Latency - the delay that an application can tolerate in the delivery of a packet
of data
?? Jitter - the variation in latency
?? Loss - the percentage of lost data
Today, the most command network technologies are IP and ATM. IP technology
offers low-cost and flexible service on network resource distribution, but it offers no
QoS guarantees, at least in its traditional form. On the other hand, ATM technology
offers QoS guarantees by using admission control based on statistical properties of
policed connections, and by static sharing of network resources among these
connections. This works well for long-lived connections of limited burstness (voice,
video), but performs poorly for short-lived, highly bursty transmissions (datagrams).
Hence, the modification of the IP technology in order to provide QoS guarantees has
become an acute and challenging demand of today network designing.
In order to accomplish the increasing demands on network resources (bandwidth),
networking companies are called upon to design and manufacture the fastest
possible switches and routers. Line (port) speed is one parameter that must grow,
and the number of ports is the other such parameter. Port speed is in the OC-12 to
OC-48 (622 Mbps to 2.5 Gbps) range today, and will grow to OC-192 (10 Gbps) in
the next few years. The number of port is in the tens to hundreds range, and will
need to grow to thousands, soon.
In this thesis, we describe a switch-router architecture that can support the two
trends: rising bandwidth demand, and rising demand for QoS guarantees. We focus
on router mechanisms that can support differentiated service to different types of
traffic (data, voice video) using the same infrastructure. We describe effective
implementations of these mechanisms, such as per-flow queueing, by using
hardware in order to accomplish the high speed rates. We discuss the functionality
of a switch-router interface at 10 Gbps line rate, and we propose advanced
techniques for the queue management implementation at such high speed. Finally,
we implement a behavioral model of the queue management subsystem, at a clockcycle-accurate level, using the Verilog HDL and we estimate the hardware
complexity of such architecture in terms of gates, flip-flops and SRAM bit count.
Introduction
8
1.1 Switch Evolution
In this section we describe the evolution of switches in order to introduce and
classify our proposed switch architecture in this evolution. Section 1.2.1 describes
the three switch generations, section 1.2.2 shows the existing switch fabric
topologies and section 1.2.3 explains the need for queueing and the alternative
queueing structures.
1.1.1
Switch Generations
Switches are classified into three generations. Switches of the first generation
consist of a computer with attached line cards. The incoming packets are passing
through an I/O bus and stored in the main memory. The Central Processing Unit
(CPU) extracts the header field of each incoming packet in order to determine its
destination and to enqueue it at the appropriate output queue. Finally, the CPU
schedules the packet departures from each output queue. Each packet crosses the I/O
bus twice: from the input line card to main memory, and from the memory to the
output line card. There are three main bottlenecks in this architecture: CPU,
memory, I/O bus. The CPU centrally performs the routing and scheduling function
for the incoming packets of all the line cards. The limited operation rate of the CPU
makes this scheme have poor performance for high speed link rates. Additionally,
the bounded throughput of the main memory and the I/O bus makes this scheme not
scalable. First generation switches are easy to build and are suitable for low speed
line cards and small valence (number of ports).
Figure 1. 1 First generation switch architecture
The architecture of the second generation switches is presented in figure 1.2.
Routing-forwarding and buffering functions are embedded in the line cards. Each
line card incorporates two buffering queues: the input and the output queue. Each
incoming packet is buffered in the input queue of the line card. When it reaches the
head of the queue, the local routing process extracts the header field of the packet
and determines its output destination. Then the packet is forwarded through the I/O
bus to the corresponding output card and is buffered in its output queue. Access to
the common bus by each line card is arbitrated performed by central controller.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Switch Evolution
9
Second generation switches are more scalable than first generation ones because the
critical path of routing and buffering is performed locally in the line cards.
Additionally the traffic of each input line passes only once through the common bus.
The only bottleneck of this scheme is the I/O bus and its central arbiter because they
can only work properly for a limited number of interface cards.
Figure 1. 2 Second generation switch architecture
Figure 1. 3 Third generation switch architecture
The third generation of switches replaces the I/O bus with a switching fabric. The
routing and buffering functions are performed locally in the line cards. The
switching fabric transfers packets from the inputs to the outputs in parallel. The
central scheduler controls the line card access to the switching fabric and updates
their routing tables. This scheme is scalable to the number of supported line cards as
well as the line rate. The figure 1.3 presents the architecture of the third generation
switches. Implementing switching fabrics and their interface cards at high-speed is a
challenging issue. Section 2.2 presents switching fabrics and section 2.3 describes
queueing architectures.
Introduction
1.1.2
10
Switching Fabric Topologies
The main switching fabric topologies include Crossbars, Banyan, Benes, and
Batcher-Banyan networks.
The simplest switch fabric is a Crossbar. A NxN crossbar consists of N input, N
2
outputs, and N crosspoints, as shown in figure 1.4. Each crosspoint has a state bit; if
the (i,j) crosspoint state is on, the traffic of the input i is forwarded to the output j. It
2
requires a switching scheduler to set the N crosspoints in order to forward the
incoming packets. A crossbar is internally non-blocking but its cost grows
proportionally to N2 crosspoints [1,chapter 8]. A crossbar can operate well for
small-size switch fabrics with supported low link rates; it implies that crossbars are
not scalable.
Figure 1. 4 Crossbars
Figure 1. 5 Banyan
The simplest self-routing switch fabric topology is the Banyan. A NxN Banyan
switch fabric consists of log2N stages and N/2 elements per stage [1, chapter 8] for
2x2 elements. The routing in a banyan network is internally non-blocking, only if
the packets at the inputs are sorted to their destination outputs and gap replications
are eliminated.
An alternative of Banyan topology is the Benes topology that is presented in the
figure 1.6. Similar to the Banyan, the Benes networks are constructed recursively, as
shown in figure 1.7. The routing of input packets to the correct output lines requires
off-line evaluation because some paths can only be determined after some other
paths are entirely defined. Additionally, Benes networks are reconfigurably nonblocking; it means that when a flow is torn down or is set up, potentially all routing
paths may have to be reconfigured in order to avoid blocking. The routing
complexity of Benes networks is proportional to ? (log2 N -1/2) [1, chapter 8],
where N is the number of network inputs/ outputs.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Switch Evolution
11
Figure 1. 6 Benes
Figure 1. 7 Benes constructure
The combination of a Sorting and a Banyan network builds a strictly non-blocking
network where the routing paths are established on-line. Sorting networks rearrange
the incoming packets to put them in an increasing or a decreasing order of their
output destinations. Batcher is a sorting network which, combined with a Banyan
network, provides strictly non-blocking networks. Figure 1.8 shows a batcherbanyan network. The routing complexity of batcher-banyan networks is proportional
to ? /4 (log2? ) (log2? + 1) [1, chapter 8], where N is the number of network inputs/
outputs.
Figure 1. 8 Batcher-Banyan
1.1.3
Queueing Architectures & Performance
The main function of a switch is to forward traffic from one interface to another.
Each interface of a switch can send and receive traffic at a finite rate. If the rate at
which traffic is directed to an interface exceeds the rate at which the interface can
forward the traffic onward, then congestion occurs. Switches may handle this
condition by queueing traffic in the switch buffer memory until the congestion
subsides. There are two basic families of queueing architectures: input and output
queueing. The output queueing locates the buffer memories at the outputs as shown
in figure 1.9. When a packet arrives, it is immediately placed in a queue that is
Introduction
12
dedicated to its outgoing line, where it will wait until its departure from the switch.
This scheme achieves full throughput utilization but requires the fabric and memory
of an N x N switch run N 1 times as fast as the line rate. This implies that output
queueing is impractical for switches with high line rates, or with a large number of
ports. For example, consider a 32x32 output queueing switch operating at a line rate
of 10Gbps. If we use a 64-byte datapath, we require memory devices that can
perform a write and a read operation every 1.6 ns.
Figure 1. 9 Output Queueing
Figure 1. 10 Input Queueing
Another architecture, input queueing, locates the buffer memory at the inputs, as
shown in figure 1.10. When a packet arrives, it is immediately placed in its input
line queue and waits until it reaches the head of the queue. Then, it waits until the
scheduler of packet departures forwards it to the appropriate output. This scheme
requires a fabric that operates as fast as the input link rate and input link buffer
memory that provides throughput twice2 the line rate. For example, consider a 32x32
input queueing switch operating at a line rate of 10Gbps. In this case, the input line
buffer memory must provide a write and a read operation every 51.2 ns
(throughput=20 Gbps). However, input queueing suffers from head of line (HOL)
blocking: if the packet at the head of an input queue is destined to a busy output, the
subsequent packets in the same queue must wait (are blocked) even if they are
destined to non-busy outputs - see figure 1.11. HOL blocking reduces the packet
delivery rate through the switch by more than one third of the input link rate.
A modified scheme of the input queueing, which overcomes the head of line
blocking, is presented in figure 1.12. This scheme is called "Advanced Input
Queueing" or "Virtual Output Queueing". In this scheme each input maintains a
separate queue for each output; thus, each incoming packet is enqueued to the
corresponding queue of its destination output. Even if we can theoretically achieve
100% packet delivery rate through the switch by using advanced input queueing, we
can not achieve it practically because the scheduler of packet departures must
operate at least N times 3 as fast as the input link rate [2].
1
When all the packets of the N inputs are destined to the same output concurrently, then the fabric
must deliver N packets within a time interval and the memory must provide N times the throughput
of each line.
2
Write an incoming packet to the queue, and read a departing packet from the queue.
3
The number of input/output links of the switch is N
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Switch Evolution
13
Figure 1. 11 Head of Line Blocking
Figure 1. 12 Advanced Input Queueing
Another method that reduces the effects of HOL blocking is to provide some
internal speedup [3], [4] of the switch fabric. A switch with a speedup of S can
deliver S cells4 from each input and S cells to each output within a time slot5. If the
input link rate is a cell per time slot and the switching fabric can deliver S (S>1)
cells per time slot through the switching fabric, we can choose a value for S that
achieves delivery rates comparable or equal to link rate. It implies that the switching
fabric will operate at faster rates than the system input/output link rates. In [5] has
proved that a speedup of 2-1/N is both necessary and sufficient for a switch with
advanced full throughput utilization. Switches with internal speedup require
buffering at the inputs before switching as well as at the outputs after switching, as
shown in figure 1.13. Input buffering is required because multiple cells may arrive
for the same output and only S of them can be delivered; the remaining must be
buffered at the inputs until they are forwarded to the output. Output buffering is
required because the switching fabric feeds each output with cells at higher rate (due
to internal speedup) than the rates that the output transmits cells to the network.
Figure 1. 13 Switches with internal speedup
4
The modern high speed switches manipulate fixed-size cells; variable size packets are segmented to
fixed-size cells
5
The time slot is the time between cell arrivals at input ports of a switch
Introduction
1.2
14
Related Work on High-Speed Switches
Today, the Internet is facing two challenges simultaneously: the increase of line
capacity (bandwidth), and the introduction of quality of service guarantees. The only
standardized 10 Gbps interfaces today are the Synchronous Optical Network/
Synchronous Digital Hierarchy (SONET/ SDH) serial interfaces.
High-speed networking that provides quality of service guarantees requires flow
isolation. Network traffic is classified to separate flows at the edge network devices
(switches/routers) by using advanced traffic handling mechanisms. The most
significant classification mechanisms, which are currently used, include: 802.1p,
Differentiated Service (Diffserv), Integrated Service (Intserv), and ATM.
802.1p [6] is a traffic-classification mechanism for supporting QoS in most local
area networks (LANs). 802.1p defines a field in the layer-2 header of packets that
can carry one of eight priority values. LAN devices, such as switches, bridges and
hubs, treat the packets according to their priority values. Intserv and Diffserv
provide QoS to IP traffic. Intserv [7] classifies IP traffic into flows accordingly to
their source/destination address, source/destination port numbers and protocol. It
requires that switches maintain state information for each flow in order to allocate
network resources to each flow. Due to the significant number of active flows,
Intserv is not scalable. Diffserv [8] categorizes the flows into three classes: best
effort, controlled rate (like best-effort without congestion), and guaranteed service
(real time with delay bounds. Diffserv aggregates individual flows into aggregate
flows, which receive service that corresponds to a predefined class. Diffserv is more
scalable than Intserv. Both Intserv and Diffserv promise guarantees by reserving
network resources by means of admission control algorithms. ATM classifies its
traffic into the virtual circuits (VC), which are supported by one of the numerous
ATM services. These include constant-bit-rate (CBR), variable-bit-rate (VBR), and
unknown-bit-rate (UBR). ATM uses a low level signaling protocol to set up and tear
down ATM VCs.
Concluding, the above mechanisms are split into two main categories: per-flow
classification (Intserv, ATM) and aggregate classification (802.1p, Diffserv). Perflow classification [9], [10] examines some fields of the packet header and classifies
the packet into the corresponding flow. Packets of the same flow are organized in a
queue (per-flow queueing). Instead, aggregate classification look at some aggregate
identifier in the packet header. Due to the requirement of maintaining independent
state for each flow and applying processing for each flow, per-flow queueing may be
not practical in the case of supporting millions of flows simultaneously. The solution
is a moderate aggregation. Traditionally, flow classification (network processing in
general) has been performed using a conventional processor. Due to the limited
processor performance (processes a fixed number of instructions per second), it is
not a scalable solution. Today, flow classification mechanisms at high speed are
implemented in hardware by using hashing functions [11] and Content Addressable
Memories (CAM). Due to the flexibility of hardware implementation to operate at
high-speed makes the latter solution more scalable than the former.
High-speed networking requires the classification/routing functions as well as
buffering to be performed locally in the interface cards [1, chapter 8]. The
distribution of these functions makes the switch architecture more scalable to the
interface speed and to the number of supported line interfaces. Modern switches that
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Thesis Contribution
15
support high-speed interfaces use combined input and output queueing schemes and
internal speedup [12], [13] in order to achieve the output queueing performance and
the input queueing scalability. Input buffering is organized in a per-flow queueing
(Virtual Output Queueing) in order to eliminate the HOL blocking. Per-flow queues
share the same memory in each input/output interface module in order to maximize
the utilization efficiency of a fixed amount of buffer memory [12]. Furthermore, to
increase the number of buffer cells that can be integrated within a given silicon area,
the shared buffer memory is implemented in DRAM rather than SRAM technology
[1, chapter 9].
1.3
Thesis Contribution
In this thesis, we study the architecture of a high-speed switch-router. High valence,
high-speed switch-routers usually consist of a switching fabric, an ingress module
for each input link and an egress module for each output link. The implementation
of the switching fabric is challenging; however, it is not a topic of this thesis. This
thesis relates to the architecture of ingress and egress modules and concentrates on
the queue management subsystem. In our opinion, the provision of advanced QoS
guarantees requires true flow isolation that can only be achieved using per-flow
queueing in connection with a good scheduler. Per-flow queueing for many
thousands of flows, was considered an excessively expensive architecture up to a
few years ago. Modern, technology, however, provides the means to implement such
architectures within a fraction of an integrated chip (IC) [14]. This thesis studies the
implementation of such architectures at OC-192 (10Gbps) line rates. We show that,
although challenging, this implementation is feasible, using the advanced hardware
techniques that were developed for supercomputers in 60's and are used in high-end
microprocessors now a day.
We propose a chip partitioning for the ingress module that economizes on chip-tochip communication, so that pin count and power consumption are reduced. We use
modern DRAM “Rambus” technology for the buffer memory, which provides high
throughput and adequate buffer space. In order to effectively use DRAM buffer
memory, accesses have to be scheduled in the presence of bank conflicts. We use
multiple, pipelined control processes to achieve out-of-order execution of DRAM
accesses. The data dependencies among successive operations are handled using
Tomasulo’s dynamic scheduling techniques (operand renaming) [15], [16]. These
sophisticated techniques also handle variable time header processing in an efficient
manner. We propose a method of economizing off-chip memories and chip pins by
locating a fraction of the queue manager data structures in the buffer memory itself,
and using free list bypassing [29] and buffer pre-allocation [29]. Finally we describe
our architecture using behavioral Verilog 6, at a clock-cycle accurate level. We
estimate the complexity of the queue manager at 60 Kgates plus 80 Kflip-flops plus
4.2 Mbits SRAM for 64 Kflows.
6
Verilog is a hardware description language
Introduction
16
1.4 Thesis Organization
Chapter 2 describes the ingress and egress module functions and proposes an
effective chip partitioning of the ingress module. It also reviews advanced
techniques for network processing (classification, routing, scheduling, and policing),
which must be implemented in hardware at OC-192 rates. Chapter 3 presents the
architecture of our pipelined queue manager. More precisely, we describe the
manner that the out-of-order execution and operand renaming techniques are applied
to achieve high operation rates. Chapter 4 explains the queue manager and Rambus
memory controller micro-architecture. We analyze the queue management
operations in terms of memory accesses and hardware implementation at a clock
cycle accurate level. Chapter 5 explains the verification of our queue management
subsystem, while chapter 6 concludes and describes open topics.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
2 Ingress/Egress Interface Module Architecture
We quickly review high-speed switch architectures and propose a chip partitioning
for the ingerss/ egress interfaces. In these architectures, the main network processing
functions are performed in the input/output interface cards (third switch generation).
This chapter describes the functionality and the architecture of the ingress and egress
interface modules and proposes an effective implementation. Section 2.1 describes
the main functions that supported from both ingress/egress interface modules.
Section 2.2 describes an effective chip partitioning of the ingress module. Sections
2.3, 2.4, 2.5 describe the main functional blocks of each separate chip and reviews
implementation alternatives.
2.1 Ingress/Egress Module Functionality
High-speed switches usually have a switching fabric with internal speedup that
requires buffering at the inputs as well as at the outputs. Additionally, packet
reassembly at egress module requires buffering. Therefore, both ingress and egress
interface modules must provide adequate memory for buffering. The ingress module
performs network & link level processing to the incoming traffic, which include
framing, classification, routing and traffic policing. The provision of Quality of
Service guarantees requires flow isolation that can be achieved using per-flow
queueing in connection with a fair scheduler. Thus, both ingress and egress modules
support queueing and scheduling. The subsequent sections analyze more thoroughly
these functions.
2.1.1
Ingress Module main Functions
In order to construct a flexible scheme, the ingress module could support the
following physical interfaces: 1 STM64 / OC-192, or 4 STM-16 / OC-48, or 16
STM-4/OC-12, thus providing 10 Gbps aggregate throughput. We consider the serial
to parallel conversion of the input stream to be performed off-chip, because coping
with 10 GHz signals may require a technology other CMOS. There are a number of
network services that require packet classification, such as routing, access-control in
firewalls, policy-based routing, provision of differentiated qualities of service, and
traffic billing. During the arrival of a new packet, the header field of the packet is
extracted and examined in order to be classified in a flow and receive the
appropriate type of service. Routing is performed on the incoming packets in order
to determine the packets' destination and to assign the proper output port. In order to
construct a more flexible and efficient architecture, the ingress module should be
configurable to support and route many types of traffic, such as IP (IPv4 and Ipv6),
ATM, IP over ATM, and MPLS. Thus the routing function must accommodate
large, off-chip routing tables (for IP longest prefix matching) and smaller, on-chip
translation tables (for fixed-size label translation).
In order to manage packets that belong to different flows, we have to organize them
in queues by using per-flow queueing. Since, thousands of flows may be active
simultaneously, queue management must be able to handle thousands of queues at
high-speed. Per-flow queueing can be implemented in the same way as advanced
Ingress/Egress Interface Module Architecture
18
input queueing1, which implies that the queue manager must operate at least2 twice
as fast as the input link rate (enqueue an incoming packet and dequeue a departing
packet).
All the queues in the ingress module dynamically share the space of a single buffer
memory, thus efficiently utilizing this buffer space. The shared buffer memory is
organized into fixed-size blocks, because this simplifies memory management; thus,
variable size packets are segmented into fixed-size segments.
The provision of quality of service guarantees also requires traffic shaping and
scheduling of packet departures, and in some cases may also require policing of the
incoming stream. Traffic shaping is required in order to conform the incoming
traffic to its traffic descriptor parameters3. The most commonly used traffic shaping
mechanism is the leaky bucket. In order to police the traffic of the majority of the
system flows, we use a leaky bucket for each flow. When thousands of flows are
supported, shaping becomes expensive due to the need for thousands of leaky
buckets. In addition, a good scheduler is required in order to service the different
flows according to their service class as well as to service flows of the same class
with fairness. The scheduler must keep state information for all the system queues,
which makes it, too, expensive.
2.1.2
Egress Module main Functions
The egress module is quite simpler compared to the ingress module, because it does
not perform flow classification or routing. It supports four main functions: buffering,
queueing, reassembly and scheduling. As mentioned before, high-speed switches use
internal speedup. Internal speedup requires a combination of input and output
buffering. Buffering in the outputs is required because the switching fabric may feed
the egress module with packets at higher rate than the eventual packet departure rate.
Thus, departing packets are accumulated at the egress module and require buffering.
Buffering is also required due to the segments' reassembly4 in order to form the
initial packet at the outputs.
In addition, queueing and scheduling is required in order to separate flows in the
outputs and provide different service to them. Policing/Shaping is not required
because it is already performed in the ingress modules. Finally, the egress module
supports 1 STM64 / OC-192, or 4 STM-16/ OC-48, or 16 STM-4 / OC-12 physical
interfaces, i.e. demultiplexing into the lower-rate links.
Because the supported functions in the egress module are a subset of the supported
functions in the ingress module we will focus on the architecture of the ingress
module.
2.2 Ingress Module Chip Partitioning
The ingress module has considerable complexity, and thus its implementation as a
single chip would be problematic, even using modern VLSI technology. This
1
An extreme of per-flow queueing is the advanced input queueing, while the other extreme is to keep
multiple flows per output.
2
It will operate at higher rates due to the internal speedup
3
A traffic descriptor is a set of parameters that describes the behavior of a data source.
4
Packets are segmented into fixed-size segments at the ingress module.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
19
Ingress Module Chip Partitioning
complexity is due to both of the number of functions to be performed (routing,
buffering, scheduling), and the number and size of memories required
(routing/classification tables, buffer memory, queueing data structures, scheduling
parameters and state).
We assumed a partitioning of the ingress module into three chips plus the off-chip
memories, as shown in figure 2.1. This partitioning reduces the chip-to-chip
communication throughput, so as to reduce both pin count and power consumption.
Packet bodies account for the majority of bits under manipulation, when compared
to packet headers. Thus, packet bodies are kept inside a single chip until they are
buffered in (necessarily off-chip) memory and when entering or leaving the ingress
module.
Figure 2. 1 Ingress Module Chip Partitioning
More precisely, we assume that reception of the incoming traffic is performed by
means of the SONET (serial) interface protocol at 10 Gbps line rate (OC-192). The
framing block extracts the encapsulated packets or cells from the SONET frames.
Next, packet headers are extracted and forwarded to the header protocol processing
chip for routing and classification. The incoming packets must wait until the routing
and classification functions have been completed, and then, the queue manager
stores them to the buffer memory and links them to the appropriate queue. As soon
as the scheduler permits the departure of a buffered packet, the queue manager in
conjunction with the memory (Rambus) controller retrieve the packet and forward it
toward the switching fabric.
Header processing represents a considerable amount of work that only
communicates with the rest of the module though packet headers and flow
identifiers; thus we assumed that it is placed in a separate chip. The same is true for
scheduling, which only communicates with the other chips through narrow words:
flow identifiers. The scheduling chip may also include traffic shaping/ policing
Ingress/Egress Interface Module Architecture
20
functions and may receive flow control backpressure from the switching fabric. The
flow control backpressure informs about congested outputs, so that packets to other
outputs are preferentially scheduled. The following sections describe the architecture
of the ingress module chips in more detail.
2.3 Datapath and Queue Management Chip
This chip accommodates the datapath and the memory management subsystem. The
incoming packets are included in SONET frames; the farming block extracts and
verifies the included packets from the SONET frames. Next, the header extraction
block extracts the selected header fields of each packet and forwards them to the
header protocol processing chip for routing and classification. After the header
extraction, the packets bodies have to wait until header processing has identified the
flow to which they belong; subsequently, queue management must identify a buffer
address in the shared memory; then, DRAM memory has to become available (see
section 3.2.3). During this waiting period, packet bodies are kept in a memory which
we call transit buffer, and are not moved from processing stage to processing stage,
so as to avoid additional power consumption.
The datapath consists of a 16-byte data bus in order to write or read a 16-byte data
block per 10 ns (100 MHz clock), which is the transfer granularity of our buffer
memory (Rambus - see section 3.2), in order to handle the input line rate of 10
Gbps. The queue manager manipulates fixed-size segments in order to simplify the
memory management. The queue manager must perform two operations per time
interval: write an incoming segment to the buffer memory and read a buffered
segment from the memory in order to forwarded it to the switching fabric. Thus the
queue manager operates at a rate twice as fast as the line rate (2 x 10 Gbps = 20
Gbps), plus relevant internal speedup. The architecture of the datapath and queue
management chip will be described more thoroughly in chapter 3.
2.4 Header Protocol Processing Chip
The provision of QoS guarantees requires intelligent allocation of the network
resources to the submitted traffic. For example, under congestion, a network device
might choose to buffer the traffic that is latency-tolerant and immediately forward
the traffic that is latency-intolerant to the next network device. In this example, the
interface capacity is a resource that is granted to the latency-intolerant traffic, while
the device memory is a resource that is granted to the latency- tolerant traffic. The
intelligent allocation of the network resources requires the classification of the
incoming traffic into separate flows in order to handle each flow differently.
2.4.1
Flow Classification
Routers classify packets in order to determine which flow they belong to, and to
decide what service they should receive. Classification may be based on an arbitrary
number of fields in the packet’s header. There are a number of network services that
require packet classification, such as routing, access-control in firewalls, policy
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Header Protocol Processing Chip
21
based routing, provision of differentiated qualities of service, and traffic billing. In
each case, it is necessary to determine which flow an arriving packet belongs to so
as to determine whether to forward it, what class of service it should receive, or how
much should be charged for transporting it. A flow is identified by a combination of
several fields in the packet's header: the source and destination Network-layer
address (32-bits each), source and destination Transport-layer port numbers (16-bits
each for TCP and UDP), Type-of-service (TOS) field (8-bits), Protocol Field (8bits), and Transport-layer protocol flags (8-bits), amounting to a total of 120 bits.
The routing function requires the examination of the destination address field in the
packet header. Because of the hierarchical structure of the IP destination address, the
longest prefix matching algorithms can perform routing effectively. Instead, when
multiple fields of the packet header must be processed in order to classify a packet,
the longest prefix matching algorithms are not applicable because the overall set of
examined fields does not have a hierarchical structure. In the latter case fixed-length
lookup functions must be applied. The following sections perform schemes for
routing and flow classifications that can be implemented in hardware.
The main flow classification mechanisms implemented in hardware are the
Recursive Flow Classification (RFC) [17] and the Hashing Flow Classification
(HFC) [18]. The main idea of RFC is the recursive mapping of the packet
classification identifiers to the corresponding flow identifiers. The RFC algorithm
can classify 30 million packets per second, using SDRAM at 125 MHz clock rate in
0.20 µm CMOS technology. The main idea of HFC is to use hashing functions in
order to classify packets. Hashing functions can be implemented effectively by using
Content Addressable Memories (CAM) [19] in 0.18 µm CMOS technology, today.
HFC promises up to 100 million packet classifications per second, using. More
information about these classification mechanisms can be found in appendix A.
Two representative implementations in hardware of IP routing function are the [18]
and [20]. In both of them, the longest prefix matching lookup is based on a tree
representation of the routing table, where the tree is searched from shorter prefixes
to longer. Both of them could be pipelined and could provide packet routing rates of
one routing lookup per memory access. The difference of these routing
implementations are that [18] uses SRAM for the routing table and is split into tree
to five pipeline stages, while [20] uses DRAM and is split into two pipeline stages.
Both of them are described more precisely in the appendix A.
2.4.2
Short Label Forwarding 1: ATM
The provision of quality of service guarantees is an inherent feature of ATM
networks. ATM networks are fundamentally connection oriented, which means that
a connection must be set up across the ATM network prior to any data transfer. Each
connection determines a separate flow in the network. The network provides
different type of service to different flows. Two types of connections exist: virtual
paths, which are identified by the virtual path identifiers (VPI), and virtual circuits,
which are identified by the combination of a VPI and a virtual channel identifier
(VCI). The VCIs and VPIs have only local significance across a particular link and
are remapped, as appropriate, at each switch.
The basic routing operation of an ATM switch is straightforward: the ATM cell is
received across a link on a known VPI/VCI value. The switch looks up the
connection value in a local translation table to determine the outgoing port (or ports)
Ingress/Egress Interface Module Architecture
22
of the connection and the new VPI/VCI value of the connection on that link. The
switch then transmits the cell on that outgoing link with the appropriate connection
identifiers. Because of the local significance of VPI/VCI label across a particular
link, these values are remapped, as necessary, at each switch. Each ATM switch
maintains a routing table that it updates whenever a connection is set up or torn
down. The table has one entry per connection. The entry has the following format:
incoming link, incoming VPI/VCI label, outgoing link, and outgoing VPI/VCI label,
as the figure 2.2 illustrates. Note that Line In, VP In, VC In fields are the index of
the translation table.
Figure 2. 2 ATM Translation Table
2.4.3
Short Label Forwarding 2: IP over ATM
The provision of QoS guarantees to the widespread IP networks creates the demand
for mapping IP technology onto ATM technology in order to exploit the inherent
QoS features of the ATM network. Mapping IP onto ATM has proved to be
challenging task. Software or hardware solutions have been proposed.
IP packets are segmented into ATM cells in order to be transmitted in the ATM
network. Routing IP traffic over ATM networks in software demands cell
reassembly in order to form the original packet and then to route the packet in the
normal way. Instead, the hardware oriented routing solutions avoids the re-assembly
overhead. Connectionless ATM [21] and Wormhole IP over (Connectionless) ATM
[22] propose IP routing over ATM without reassembling the packet's cells. The main
idea of both solutions is that the routing information of an incoming packet is
contained in the first cell; this applies to both IP versions (IPv4 and IPv6), as figure
2.3 shows. The ATM switch knows, by context, which cell on an ATM virtual
channel is the first in an AAL5 sequence. By extracting the IP destination address
from the first cell it can perform IP routing in hardware. When an outgoing link and
a VP/VC label are assigned, then the first and all subsequent cells are forwarded to
the next hop. The modification of the ATM switches in order to support the routing
function is proposed in [21]. Instead, [22] proposes single-input, single output
wormhole IP routers, which function as VP/VC translation filter and interoperate
with existing ATM switches and networks. The VP/VC label is locally assigned in
order to be achieved cell’s multiplexing in the outgoing links.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Header Protocol Processing Chip
23
Figure 2. 3 Ipv4, Ipv6 packet format
2.4.4
Short Label Forwarding 3: MPLS
Multiprotocol Label Switching (MPLS) was introduced in order to improve the
price/performance of network layer routing, to improve the scalability of the
network layer, and to provide greater flexibility in the delivery of (new) routing
services. The innovation of this technology is that it introduces the technique of
label swapping in the routing function. The main issue for the MPLS working group
is the integration of the label swapping forwarding with the existing IP network
layer routing.
Label swapping allows packet forwarding to be based on an exact match of a short
label, rather than the longest prefix match algorithms currently applied to IP routing.
Label swapping is a very powerful technique that is already applied to ATM
networks. It simplifies and increases the speed of the forwarding function. More
precisely, MPLS uses the conventional IP protocols (OSPF and BGP) in order to
build the routing tables, but uses the ATM fixed-size label forwarding paradigm.
MPLS allows hierarchical operation, which means that it can be used for routing at
multiple levels. Figure 2.4 illustrates an example of how MPLS may operate in a
hierarchy. The routers R1, R2, R3, R8, R9, R10 are domain boundary routers, while
R4, R5, R6, R7 are domain internal routers. In this example, there are two levels of
routing taking place: the OSPF for internal routers and the BGP for the domain
boundary routers. When the IP packet traverses the domain 2, it will contain two
labels, encoded as a label stack. The higher level label would be encapsulated inside
a header specifying a lower level label used within domain 2.
Figure 2. 4 MPLS Hierarchical Network
Ingress/Egress Interface Module Architecture
24
2.5 Scheduling-Policing Chip
This chip accommodates a scheduler and a traffic policer/shaper. The discipline and
the implementation of the scheduler and the policer/shaper are described in the
subsequent sections.
2.5.1
Basic Disciplines on Scheduling
Scheduling discipline has four parameters: the number of priority levels, the type of
service (work-conserving or non-work-conserving), the degree of aggregation, and
the service order within a priority level.
If a scheduler supports priorities, then it serves a packet from a priority level only if
there are no packets waiting for service in an upper priority level. With such a
scheme, connections that require QoS and are intolerant of delays can be serviced
with higher priority than others. However, a priority scheme allows a misbehaving
user at higher priority level to increase the delay and decrease the available
bandwidth for connections at all lower priority levels. An extreme case of this is
starvation, where the scheduler never serves a packet of a lower priority level
because there is always something to send from a higher priority level. In an
integrated services network, at least three priority levels are desirable: a higher
priority level for urgent messages, usually for network control; a medium priority
level for guaranteed service traffic; and a low priority level for best-effort traffic.
There are two types of scheduling service: the work conserving and the non-workconserving disciplines. A work-conserving scheduler is idle only when there is no
packet awaiting service. In contrast, a non-work-conserving scheduler is idle even if
it has packets to serve, so as to shape the outgoing traffic, in order to reduce the
traffic burstiness and the delay jitter. The work-conserving discipline is more
suitable for best-effort traffic (IP traffic) and the non-work-conserving discipline is
better applied to guaranteed-service traffic (voice – video). The new integrated
network systems need schedulers that will serve both types of traffic. Note that a
non-work-conserving scheduler does not necessary have to waste bandwidth when it
has no eligible packets to serve: it can simply serve best-effort packets to use up the
otherwise idle link. Implementation approaches for the non-work-conserving
scheduler are reviewed in section 2.5.6.
An important decision on the scheduler design is the degree of aggregation. The
scheduler aggregates individual connections in order to simplify the manipulation of
them, especially in the case that the supported connections are many (thousands of
connections). Routing and connection admission protocols require from the network
switches to advertise their current state to the rest of the network, allowing a source
to select a path that is likely to have sufficient resources. The larger the state kept in
the scheduler, the more there is to advertise, which costs bandwidth. Additionally,
the smaller the amount of scheduler states the easier it is to implement it. On the
other hand, if the scheduler uses a great degree of aggregation, it can not
differentiate the aggregated connections in order to give them different bandwidth
and delay bounds. Additionally, connections that belong to the same class are not
protected from each other. Because the scheduler cannot distinguish among
connections in the same class, the misbehavior of a connection in the class affects
the whole. By concluding the above considerations, an intermediate aggregation
provides an effective scheduling scheme.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Scheduling-Policing Chip
25
The final parameter in designing scheduling discipline is the order in which the
scheduler serves packets from connections at a given priority level. There are three
fundamental choices: FCFS, weighted Round Robin, and out-of-order according to a
per packet service tag. Servicing the packets in the same order as the order of their
arrivals is easy to implement but it is not a flexible and fair decision. The roundrobin scheme is a fair solution with easy implementation. Finally, out-of-order
packet service needs a significant overhead for the packet tags during packet arrivals
and requires specialized hardware data structures, such as sorted linked lists, to
support out of order service. By using out of order service we accomplish to provide
differentiated service to the different connections of the same priority.
2.5.2
Scheduling Best-Effort Connections
The main goal in scheduling best-effort connections is fairness in order to achieve
each connection the same amount of throughput and tolerate the same delay with the
other connections of the same priority. We present some best-effort scheduling.
An ideal work-conserving scheduling discipline is called Generalized Processor
Sharing (GPS). GPS serves packets from different queues by visiting each nonempty queue in turn and serving an infinitesimally small amount of data from each
queue, so that, in any finite time interval, it visits every logical queue.
The simplest emulation of GPS is the round-robin, which serves entire packets,
instead of infinitesimal amount from each non-empty queue; the packets of each
connection are temporarily buffered in a separate queue for each connection. The
round-robin policy serves each queue in a cyclic order and gives the same
bandwidth to all queues. In the case that the queues take different amount of
bandwidth from each other then the scheduler assigns a weight to each queue and
serves them proportion to their weight. This scheduling discipline with weights is
called weighted round-robin (WRR). The main restriction of both algorithms is that
they manipulate fixed-size packets (for fairness).
Deficit round-robin (DRR) modifies weighted Round Robin scheduling to allow it to
handle variable packet size. A DRR scheduler associates each connection with a
deficit counter initialized to 0. The scheduler visits all the queues in turn and tries to
serve one quantum worth of bits from each visited connection. The packet at the
head of the queue is serviced if it is no larger than the quantum size. If it is larger,
the quantum is added to the connection’s deficit counter. If the schedule visits a
connection where the sum of the connection deficit counter and the quantum is
larger than or equal to the size of the packet at the head of the queue, then the packet
at the head of the queue is serviced, and the deficit counter is reduced by the packet
size.
Weighted fair queueing (WFQ) is another approximation of GPS scheduling. The
main idea of WFQ is to compute the time a packet would complete service if a GPS
server serviced it. In other words, the WFQ simulates the GPS algorithm and assigns
the result of the simulation to each packet as a tag. The packets are serviced in
decreasing order of their tags by using a heap – priority queue.
Ingress/Egress Interface Module Architecture
2.5.3
26
Scheduling Guaranteed-Service Connections
The weighted fair queueing scheduling can be applied to provide connection
performance guarantees. A variant of WFQ is the virtual clock scheduling. A virtual
clock scheduler stamps packets with a tag, and packets are serviced in order of their
tags, as in WFQ. However, the tags are not computed to emulate GPS scheduling,
but to emulate time-division multiplexing.
Another variant of WFQ is the earliest-due-date scheduling. Similar to WFQ, it
assigns a tag to each packet that called deadlines and serves them in order of their
deadlines. The scheduler set a packet’s deadline to the time at which it should be
sent had it been received according to the connection’s contract, that is slower than
its peak rate. To be noted here that during a connection’s set up, each source
negotiates a service contract with the scheduler and a traffic descriptor is
determined. The traffic descriptor will be explained in the leaky bucket section.
Figure 2. 5. Rate-controlled Scheduler
Another scheme of guaranteed-service scheduling is the rate-controlled. The ratecontrolled scheduling can give connection bandwidth, delay, and delay-jitter bounds.
It has two components: a regulator and a scheduler, as the figure 2.5 shows. The
incoming packets are placed in the regulator, which uses one of many algorithms to
determine the packet’s eligibility time stamps. When packet becomes eligible, it is
placed in the scheduler, which arbitrates among eligible packets. By delaying
packets in the regulator, we can shape the flow of incoming packets to obey any
constraint. The scheduler can service packets in a first-come-first-served order or
serve them using WFQ. The service properties of a rate-controlled scheduler depend
on the choice of the regulator and scheduler.
2.5.4
Leaky Bucket
Traffic shaping is a mechanism that alters the traffic characteristics of a stream of
packets/cells in order to make them conform to a traffic descriptor. A traffic
descriptor is a set of parameters that describes the behavior of a data source. There
are three main parameters that describe the data source traffic: the average rate (a),
the peak rate (p), and the burst size (bs). Shaping the data source traffic to the above
traffic parameters means that the data source can send packets at the long-term
average rate (a) or it can send bursts of size (bs packets) at the peak rate (p). Traffic
shaping is performed at the entrance nodes of the network and the devices that shape
the incoming traffic are called regulators.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Scheduling-Policing Chip
27
Figure 2. 6 Leaky Bucket
The leaky bucket is such a regulator. The leaky bucket has a pool of tokens - a
token bucket. The leaky bucket accumulates fixed-size tokens in the bucket. An
incoming packet can be transmitted only if the bucket has enough tokens. Otherwise,
the packet waits in a buffer until the bucket has enough tokens for the length of the
packet. The figure 2.6 illustrates the leaky bucket operation. As the figure 2.6 shows,
the regulator adds tokens to the bucket at the average rate (a). On a packet departure,
the leaky bucket removes the appropriate number of tokens. If we consider that the
incoming packets have fixed-size or are segmented into fixed-size units and that for
a packet departure one token is removed from the bucket, then the size of the bucket
corresponds to burst size (bs). By replenishing tokens in the bucket at the average
rate (a) and permitting the departure of (bs) contiguous packets we control the two
of the three traffic parameters: the average rate and the burst size. In order to control
the peak rate, a second leaky bucket must be introduced. If the token replenishment
interval corresponds to the peak rate, and the token bucket size is set to one token,
then the second leaky bucket is a peak rate regulator. The second leaky bucket is
located prior to the first leaky bucket in order to insert traffic, which is conforming
to peak rate. This leaky bucket does not have a buffer, but instead of dropping the
non-conformant packets it marks them and transmits them to the next leaky bucket.
In case of buffer overflow the marked packets are dropped. If the next leaky bucket
does not have a buffer to keep the non-conforming packets, it is called policer. A
policer drops the non-conforming or marked packets. Figure 2.7 shows the two
leaky buckets. A leaky bucket can be implemented as a calendar queue, see section
2.5.5.
Figure 2. 7. The two leaky buckets traffic shaping mechanism
Ingress/Egress Interface Module Architecture
2.5.5
28
Calendar Queue
Figure 2. 8 Calendar queue structure
A calendar queue consists of a clock and an array of pointers to lists of packets, as
figure 2.8 shows. Each pointer corresponds to an array slot points to the list of
packets that will be serviced during this slot. The “initial” duration of slot equals to
the calendar queue’s clock period. However, due to the variability in the number of
the packets in each list, the time slot duration is variable. When all the packets of a
slot’s list are serviced, we move to the next slot. The pointer of the next slot indexes
to the corresponding list of packets. A packet is inserted to the proper slot after the
scheduler assigns a slot tag to it. The size of a calendar queue in slots must be
greater than a limited value in order to avoid conflicts; a packet that must be
serviced during a slot in the current round may be linked in the same list with a
packet that must be serviced at the next round. The calendar queue size is estimated
as follows: the number of slots times the calendar queue clock period must be
greater than the period of the slowest connection’s traffic rate.
2.5.6
Heap Management
The heap data structure is an efficient way to organize a priority queue. A heap is
organized as a balanced tree, completely filled on all levels except possibly the
lowest, which is filled from left up to a point. Each node to the tree corresponds to
an element, and stores the id and priority values of that element. All the nodes have
the same number of children, which defines the heap’s degree and each node has a
higher priority than all its children. The heap can provide the element with the
highest priority in O(1) time; due to the heap property, this element will stored at the
root. However, inserting a new element to the heap or decreasing the priority of an
element in the heap may cost the time for traversing the heap (O(N) complexity,
where N is the number of elements in the heap). An implementation approach in
hardware is described in [23], while the pipelined version is presented in [24]. The
pipelined heap management can achieve O(1) operation rates with latency still in the
order of O(log N).
2.5.7
An Advanced Scheduler Architecture
Designing a scheduler for thousands of flows (e.g. 64K flows) at high-speed
(10Gbps) is a challenging task. We propose a multi-stage scheduling architecture,
which combines disciplines of both work-conserving and non-work-conserving
schedulers. Figure 2.9 shows such a two-stage scheduler.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Scheduling-Policing Chip
29
The first stage distinguishes the guaranteed-service and best effort flows. It also
aggregates the flows into sets in order to minimize the control state information that
the scheduler must keep and advertise to the remaining network. For example, flows
for the same output and of the same priority may belong to the same set. In a scheme
where there are 64 outputs and four priority levels, the 64k flows may be aggregated
into 256 sets of flows. In the case of guaranteed service flows, the rate-controlled
scheduling discipline performs well. As mentioned above, a rate-controlled
scheduler consists of a regulator and a scheduler. More precisely, there must be a
leaky bucket per flow, which implies that a few thousand leaky buckets must be
supported. After the shaping of incoming traffic, an earliest-due-date scheduler
assigns a time stamp to the conforming packets. These time stamps, then, are
inserted in a priority queue (heap). The heap prepares an eligible packet to be
forwarded to the second stage of scheduling. On the other hand, a weighted roundrobin scheduler must service the best effort set of flows.
The second stage of scheduling has to serve the aggregated flows of the previous
stage. The sets of guaranteed flows have higher priority than the sets of best effort
flows. A weighted round robin scheduler is quite simple and works efficiently in that
stage of scheduling. The main goal of this stage is fairness. This stage also receives
flow control information from the switching fabric, which informs for the state of
the output links traffic and for traffic congestion. This stage stalls the service of the
congested sets of flows. Finally, a weighted round robin scheduling with priorities
would perform well, if it was applied to the second stage of scheduling.
Figure 2. 9 A two stage scheduler
3 Datapath & Queue Management Chip Architecture
Our proposed architecture keeps the incoming packets in a shared buffer memory
and the queue manager organizes them in logical queues. The queue manager does
not operate on the packets themselves, but on the pointers assigned to them. The
manipulation of pointers requires appropriate data structures to store them. In
section 3.1 we show the queue manager operations and the main data structures.
Section 3.2 discusses buffer memory technology. Sections 3.3 to 3.5 present the
pipelined architecture of our queue manager, the pipeline dependencies, and
effective techniques for handling them. Section 3.6 proposes an effective scheme of
queue pointer management and section 3.7 presents the final queue management
architecture.
3.1 Queue Management Data Structures
The queue manager should manipulate fixed-size units, in order to operate at highspeed. Fixed-size units simplify hardware, thus reducing its cost and increasing its
speed. Also, implementing multiple queues inside a shared buffer memory, in
hardware, is almost impossible unless all memory allocation is done in multiples of
a fixed-size block unit. Moreover, efficiently scheduling the traffic over a switching
fabric is very hard unless all traffic sources start and finish their transmissions in
synchrony, thus implying that they all use a common-size unit of transmission. The
fixed-size unit of traffic must be relatively small, so as to reduce delay for highpriority traffic. In our implementation the queue manager manipulates 64-byte 1
units, since this size is close to the ATM cell size. For this purpose, variable size
packets are fragmented into 64-byte segments. We underline that the queue manager
manipulates fixed-size segments instead of variable size packets.
Figure 3. 1 Queue Manager Data Structures
1
by placing the queue management pointers in the buffer memory, as proposed in section 3.6.1, the
unit size is reduced to 60-bytes
Datapath & Queue Management Chip Architecture
32
We define as buffer, the memory space required to store one packet segment in the
buffer memory (i.e. 64-bytes). The queue manager handles the buffers as units,
when performing queue operations. For this purpose one pointer, the next pointer, is
associated to each buffer, and the queue manager executes instructions that only use
such pointers as arguments. Note here, that the memory contains two types of
buffers: free and occupied buffers. The free buffers do not store any segment, while
the occupied buffers store segments. The queue manager organizes the free buffers
in a single linked list, called Free List, by linking their associated pointers, as shown
in the figure 3.1. The occupied buffers are organized in queues by linking their
associated pointers in linked lists (see figure 3.1). Occupied buffers that store
segments of the same flow are organized in the same queue. Each pointer in a list
indicates the address of the next buffer. Apart from the next pointers, the queue
manager needs two additional pointers: one pointing to the head and one pointing to
the tail of each list. The head and tail pairs of all the system queues are maintained
in the Queue table, while the head/tail pair of the Free List is maintained in the Free
List head/ tail register 2. Additionally, it is necessary to store one bit per list, the
Empty bit, to indicate whether the corresponding list is empty or not.
3.1.1
Fragmentation loss
The segmentation of the variable-size incoming packets into fixed-size segments
introduces fragmentation loss3. Consider the extreme case where the segment size is
60-bytes and the incoming packet size is 61-bytes. The incoming packet will be
segmented into two 60-bytes segments; this situation augments the incoming traffic
to 1.96 times the input traffic. The normal IP packet size usually varies from 40
bytes TCP acknowledgement to 1500 bytes Ethernet Maximum Transfer Unit MTU
[27][28]. Shorter and longer packets are also possible, but are seldom seen in core
networks. In the case of 40 bytes TCP acknowledgement the fragmentation loss
augments the input traffic by 50%, while in the case of 1500 bytes Ethernet packet
the fragmentation loss is 0. The fragmentation loss of the ATM traffic is 0.13, while
the fragmentation loss of an average packet size of 270 bytes [22], [28], [30]
augments the input traffic by 20%.
3.1.2
Queue Management Operations
In this section we describe the two main queue management operations: enqueue
and dequeue. The enqueue operation is illustrated in figure 3.2: consider a buffer
memory that contains eight buffers; three of them are occupied and belong to queue
Q1, while the remaining five buffers belong to the free list (figure 3.2, left). When a
new segment arrives and the header processor defines that it belongs to Q1, the
queue manager must enqueue it there. A free buffer must be extracted from the free
list. As we mentioned above, the queue manager does not operate on memory
buffers themselves, but on their associated pointers. We extract a pointer to a free
buffer by reading the head pointer of the Free List (head pointer = 3). The next step
2
in order to handle free buffers more efficiently, we organize them in per-memory bank lists, as
proposed in section 3.6.5.
? packet _ size ?
? segment _ size ? * segment _ size
?
?
3
fragmentation loss = 1 packet _ size
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Buffer Memory Technology
33
is to store the segment to the extracted free buffer and to link this buffer to the
queue. Writing the free buffer pointer to the associated next pointer field of the Q1
tail, the free buffer is linked to the queue. Finally, the head of the free list and the
tail of the queue must be updated. Figure 3.2 (right) shows the state of queue
manager data structures after the enqueue.
Similar to the enqueue operation, we describe the dequeue operation by using the
example in figure 3.3. Consider a memory of eight buffers; four of them are
occupied and belong to queue Q1 and the remaining belong to free list (figure 3.3
left). When the scheduler decides to forward a segment from Q1, the queue manager
must perform a dequeue operation. The first step is to read the pointer to the buffer
at the head of Q1 (head pointer=0). The next step is to retrieve the buffered segment
body from the memory and to link the corresponding buffer to free list. Writing the
buffer’s pointer to the next pointer field of the free list tail performs this linking.
Finally, the tail pointer of the free list and the head pointer of Q1 must be updated.
Figure 3.2 (right) shows the state of the data structures after the dequeue.
Figure 3. 2 Enqueue Operation
Figure 3. 3 Dequeue Operation
3.2 Buffer Memory Technology
3.2.1
DRAM versus SRAM
A crucial design decision at such high rates is the choice of buffer memory
technology. SRAM provides high-throughput but limited capacity, while DRAM
offers comparable throughput4 and significantly higher capacity per unit cost. In
order to increase the number of buffers that can be integrated within ingress/ egress
4
Memory Chip throughput is a matter of I/O interface rather than storage core; SRAM and DRAM
both use similar I/O interface techniques, today. Of course, they differ in latency
Datapath & Queue Management Chip Architecture
34
system, we assumed DRAM rather than SRAM technology [25]. Among DRAM
technologies, we chose Rambus [26] over DDR or SDRAM, because Rambus offers
higher throughput while using less pins.
3.2.2
Rambus DRAM Technology
Rambus technology [25] provides 12.8Gbps peak throughput per memory device
(RDRAM) by using a 2-byte wide datapath at 400 with double clocking i.e.
800Mbps/ pin. A “RIMM” module packages up to 16 devices and provides
128Mbytes total capacity (accommodates 2 million of 64-byte buffers). Each
memory device is partitioned into 16 banks (the RIMM module contains 256 banks
totally) in order to provide interleaving, i.e. in order to allow multiple parallel
accesses. Up to four accesses to different banks can be in progress, simultaneously.
The memory transfer granularity is 16-byte blocks.
The access row latency is about 50ns, while successive accesses to the same or
adjacent banks may be performed every 80ns to 100ns due to the bank precharging
period. Note that access latency refers to the time interval between the insertion of a
new read or write command at the Rambus channel and the response5 of a Rambus
memory device with the data or the loading of the writing data to the Rambus
channel. In other words, this latency is the Rambus core latency. However, any
external system accesses (communicates with) the Rambus core through the memory
controller and the Rambus interface cell (RAC). The memory controller provides
the protocol for performing read and write transactions to the Rambus memory
channel (Rambus core), while the RAC interfaces the core logic of a CMOS ASIC
memory controller to the high-speed Rambus Channel. The Rambus controller and
the RAC interface introduce additional latency to a memory transaction. Hence, the
total latency, which an external device tolerates/receives when accessing the
Rambus memory, is about 100 ns.
Figure 3. 4 Buffer Memory Throughput
The queue manager must operate at least twice as fast as the link rate because it has
to buffer the incoming packets in the buffer memory and to retrieve the buffered
packets from the memory in order to forward them toward the switching fabric,
simultaneously. Assuming the OC-192 physical interface, the queue manager must
handle 12 Gbps input traffic and 12 Gbps output traffic due to fragmentation loss. In
order to support the 24 Gbps throughput, two RIMM memory modules are required.
The total 25.6 Gbps provided throughput by the two RIMM modules allowing a
5
During a read operation a memory device loads the data results to the rambus channel after passing
a constant delay from the corresponding read command insertion
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
35
Multi-Queue Management Architecture at High-Speed (10Gbps)
moderate internal speedup of 1 Gbps, as shown in figure 3.4.
3.2.3
Out-of-Order DRAM Accesses
When a memory transaction tries to access a currently busy bank (a bank that has
not yet been precharged), as opposed to an available bank, we say that a bank
conflict has occurred. This conflict causes the new transaction to be delayed until the
bank becomes available, thus reducing memory utilization. When random accesses
are made to an interleaved DRAM, some bank conflict will inevitable occur, as
illustrated in left part of the figure 3.5, where a memory consisting of 4 banks (A, B,
C, D) is assumed. If the bank cycle time is 3 time units, we can access the same
bank every 3 time units. Therefore, the second and the successive transactions would
be delayed two time slots.
In order to reduce the number of bank conflicts, thus increasing memory utilization,
we rearrange the order of memory accesses, as in the simple example of the right
part of the figure 3.5. The reordering of memory accesses implies out-of-order
enqueue and dequeue operations, which requires some control hardware complexity
(section 5.2), but is quite beneficial to performance.
Figure 3. 5 Non-Interleaved versus Interleaved Transaction
3.3 Multi-Queue Management Architecture at High-Speed
(10Gbps)
In this section we will present the architecture of the datapath and queue
management chip. This architecture implements dynamic scheduling of memory
accesses (out-of-order memory transactions) in order to maximize buffer memory
utilization and increase queue management performance. Section 3.3.1 describes the
overall architecture and presents the way to achieve dynamic scheduling. Section
3.3.2 identifies the reasons for pipelining this architecture in order to achieve high
operation rates (an enqueueing and a dequeueing per time slot). Section 3.3.3
describes the pipelined architecture consisting of multiple control processes. Section
3.3.4 presents the main data dependencies of the pipelined architecture, and section
3.3.5 describes how we handle these dependencies.
3.3.1
Queue Management Architecture Overview
The datapath and queue management chip architecture achieves high-operation rates
by keeping the majority of incoming traffic to this single chip boundaries and by
using the buffer memory (DRAM) throughput effectively. The first is achieved by
Datapath & Queue Management Chip Architecture
36
transmitting only a small fraction of the incoming traffic (only some fields of each
packet header) to the header-processing chip, which requires the header processor to
operate at much lower rate than the input link rate. The second is achieved by using
out-of-order memory access techniques.
The datapath consists of three parts as shown in figure 3.6. In the first part, the
incoming traffic is temporary stored in the “transit buffer”. In the second part,
packets are moved from the transit buffer to the buffer memory. In the third part, the
traffic is moved from the buffer memory to the switching fabric. In the first part, the
incoming packets are kept in the transit buffer until the header processor classifies
them to the proper flow. An incoming packet has also to wait in the transit buffer
until the queue manager identifies a free buffer and the memory is available for
accessing. Concluding, by keeping the incoming packets in the transit buffer, we
avoid packet movements from processing stage to processing stage, so as to avoid
additional throughput and power consumption.
Figure 3. 6 Multi-Queue Management Block Diagram
Our architecture supports dynamic scheduling of memory transactions by using two
operation tables: the pending write table (PWT), and the pending read table (PRT),
as shown in figure 3.6. We divide the enqueue and the dequeue operations into two
phases: the issue and the execution phase. During the issue phase, the enqueue or
dequeue operation selects its arguments and stores the corresponding write or read
transaction to a PWT entry or PRT entry, respectively. The arguments of an enqueue
operation are two: the address of a free buffer in the memory and the queue tail
pointer. The enqueue operation writes the packet body to the free buffer and links
the associated pointer of this buffer to the queue tail. The dequeue operation has
only one argument: the address of the departing buffer. The dequeue operation
retrieves the packet body from the departing buffer and links the associated pointer
of this buffer to the free list tail. During the execution phase, a search engine
examines the pending write and read transactions of the PWT and PRT, in parallel,
and selects an eligible (non-conflicting) write and an eligible read transaction to
perform. The parallel search in the PWT and PRT provides the flexibility of
reordering the memory transactions. The implementation of such a parallel search
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
37
Multi-Queue Management Architecture at High-Speed (10Gbps)
engine is a challenging issue, and is described in section 4.2.10.
3.3.2
Why Pipelined Queue Manager
As we said (section 3.2.2), the queue manager must operate at 25.6 Gbps, i.e it must
enqueue or dequeue one segment every 40 ns6 (time slot). We examine the
performance of a non-pipelined queue manager by estimating the latency of an
enqueue and a dequeue operation. In our evaluation, we consider that the queue
manager data structures are kept in fast on-chip SRAMs, which implies that the data
structure updates can be performed during the time slot of writing or reading a
packet to/from the buffer memory.
During a packet arrival, some header fields of the packet are extracted and
forwarded to the header processor for routing and classification, and, concurrently,
the packet is kept to the transit buffer. The packet has to wait in the transit buffer
because the header processing may last more than a time slot and the dynamic
scheduler of memory transactions may delay the packet writing to the memory for
some time slots. Hence, the latency of an enqueue operation may last many time
slots. Similarly to the enqueue operation, the dequeue operation may last some time
slots because of two reasons: due to the dynamic scheduling of memory transactions
and due to the DRAM memory access latency. More precisely, a read operation may
not be performed immediately because the corresponding memory bank is not
available due to the memory precharging period. Additionally, we remind that the
Rambus memory read access latency is about 100ns (2.5 time slots).
Since the latency of an enqueue or dequeue operation is longer than a time slot,
pipelining is needed to achieve the desired operation rate. The only parameter we
know for this pipeline is the pipeline stage length; the length of the queue
management pipeline stage must be equal to a time slot in order to insert an enqueue
and a dequeue operation per time slot.
3.3.3
Why Multiple Control Processes
Given that the enqueue and dequeue operations are two independent operations, the
queue management pipelining will consist of two independent pipelined processes.
Additionally, both enqueue and dequeue operation is subdivided into two phases in
order to support out-of-order memory accesses. The first phase accumulates pending
transactions to the operation tables and the second phase services these pending
transactions in an order that utilizes the memory throughput more effectively. This
implies that each of two phases can implement a separate process, which works
independently of the other. Accessing common resources7 performs the
communication of these processes. Concluding until now, we have four independent
pipelined processes (2 for enqueue operation and 2 for dequeue operation) that
implement the queue management pipeline.
As mentioned in the section 3.3.1, the datapath from the input to the buffer memory
is split into two stages: from the input to the transit buffer and from the transit buffer
to the memory. These two stages operate separately: the first stage accumulates
6
40 ns time slot is the time interval for reading or writing an 64-byte segment from/to Rambus buffer
memory
7
In the case of enqueue operation, the common resource is the Pending Write Table. In the case of
dequeue operation, the common resource is the Pending Read Table
Datapath & Queue Management Chip Architecture
38
incoming traffic to the transit buffer, while the second retrieves the buffered packets
and loads them to the buffer memory. Hence, both stages can implement two
independent, pipelined processes. Note that the second process in embedded in the
second enqueue operation process.
In order to interface the queue management operation processes to the buffer
memory (Rambus) controller, we introduce an additional process. This process
inserts write and read transactions to the memory controller and receives the
memory data responses in order to forward retrieved packets to the switching fabric.
In conclusion, the pipelined queue management architecture is consisting of six
parallel and fully pipelined control processes. Three of them are dedicated to
manipulate the incoming traffic (enqueue operation). The first, which we call
“packet entry” process, buffers the incoming packets to the transit buffer. The
second, which we call “enqueue issuing” process, issues enqueue operations and
keeps their arguments and the corresponding write transactions in the pending write
table. The third, which we call “enqueue execution” process, selects an eligible write
transaction from the PWT, and executes a pending enqueue operation by transferring
(writing) the buffered packet body from the transit buffer to the memory.
Additionally, two processes are dedicated to manipulate the outgoing traffic
(dequeue operation). The first, which we call “dequeue issuing” process, issues
dequeue operations and keeps their arguments and the corresponding read
transaction in the pending read table. The second, which we call “dequeue
execution” process, selects an eligible read transaction from the PRT, and executes a
pending dequeue operation by retrieving (reading) the buffered packet body from
the memory and forwarding it to the switching fabric. Finally, we call the process,
which interfaces the queue management processes to the memory controller, as
“queue management interface” process. All the queue management processes will be
described more thoroughly in the subsequent sections. In order to simplify our
description, we consider fixed-size (i.e. 64-byte7) packets.
Packet Entry process
This process receives the segments of a packet and stores them to the transit buffer.
The transit buffer is a memory organized as a set of 64-byte blocks. Each block is
identified by its address, which we call transit_id. Upon a packet arrival the entry
process extracts the transit_id of a free block and stores the packet, as the figure 3.7
shows. Note that the transit_id of the packet accompanies the transmitted header
fields of the packet to header processor in order to identify them among others.
7
we consider 64-byte packets in order to avoid segmentation; because the segment size is 64-byte,
the packets and the segments are identical quantities of traffic.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Multi-Queue Management Architecture at High-Speed (10Gbps)
39
Figure 3. 7 . Incoming segment entry process
Enqueue issuing process
The initiator of an enqueue operation is the header processor. Whenever the header
processor assigns an incoming packet to a flow, it sends a command to the queue
manager in order to enqueue the packet in the corresponding queue. The enqueueing
command of the header processor has the following format: the transit_id of the
waiting packet and the queue_id of the flow. When the enqueue issuing process
receives an enqueue command, it identifies the address of a free buffer and the
queue tail pointer. The free buffer is required to keep the packet in the memory,
while the queue tail pointer is required to link the buffer to the queue. As soon as,
the process accomplishes these arguments, it stores them to a Pending Write Table
entry. Note that each entry in the PWT is associated to an entry in the transit buffer
by assigning the same transit_it to both. The PWT entry keeps control information,
while the transit buffer keeps the packet body of a pending enqueue operation. The
acquisition of enqueue operation arguments and their keeping to the PWT entry is
shown in the figure 3.8. The figure 3.8a shows the state of the transit buffer and
PWT before enqueue issuing, while the figure 3.8.b shows the state of these blocks
after enqueue issuing. We use an additional flag, the “ready flag” for each entry in
the PWT to indicate whether the pending enqueue operations has accomplished its
arguments or not.
Figure 3. 8 Enqueue Issuing Process
Datapath & Queue Management Chip Architecture
40
Enqueue execution process
This process traces the valid and ready pending enqueue operations in the PWT, and
selects a write transaction that will not cause bank conflict in the memory. When it
selects an eligible write transaction, sends the corresponding write command to the
memory controller8. The write command consists of the write address (free buffer
address) and the data (the packet body that is buffered in the PWT). During the write
transaction period (40ns), this process links the writing buffer to the queue by
writing its associated pointer to the next pointer field of the queue tail buffer.
Additionally, it has to updates the queue management data structures. More
precisely, it updates the new queue tail pointer with the address of the writing buffer,
and updates the new free list head pointer. The new free list head pointer is
accomplished by accessing the next pointer field of the writing buffer. The functions
of the packet writing and the data structures update are presented in the figure 3.9
(next page).
Dequeue issuing process
The initiator of an dequeue operation is the scheduler of packet departures.
Whenever the scheduler decides to forward packets from a flow, it sends a command
to the queue manager in order to dequeue a packet of the corresponding queue. The
dequeueing command of the scheduler has the following format: queue_id of the
servicing flow. When the dequeue issuing process receives a dequeue command, it
identifies the queue head pointer, which is the address of the departing buffer. Then,
it stores the address of the departing buffer to the PRT, in order to perform a read
transaction later. The PRT keeps control information for pending dequeue
operations. Similar to PWT, we use an additional flag, the “ready flag”, which
indicates whether the pending dequeue operations has accomplished its arguments or
not.
Figure 3. 9 Enqueue Execution Process
8
the write command is sent to the memory controller via the queue management interface process.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Queue Management Pipeline Dependencies
41
Dequeue executing process
This process traces the valid and ready pending dequeues in the PRT, and selects a
read transaction that will not cause bank conflict in the memory. When it selects an
eligible read transaction, it sends the corresponding read command to the memory
controller9. The read command consists of the read address (departing buffer
address). After the read transaction is send to the controller, this process links the
departing buffer to the free list by writing its associated pointer to the next pointer
field of the free list tail buffer. Additionally, it has to updates the queue management
data structures. More precisely, it updates the new free list tail pointer with the
address of the departing buffer, and updates the new queue head pointer. The new
queue head pointer is accomplished by accessing the next pointer field of the
departing buffer.
Queue Management Interface Process
This process converts the queue management write and read transactions to a form
that is compatible to the memory controller. It also receives the data responses from
the memory and forwards them to the switching fabric. The operation of this process
is illustrated in the figure 3.10.
Figure 3. 10 Queue Management Interface Process
3.4 Queue Management Pipeline Dependencies
By designing a system in a pipelined fashion, its throughput is increased but the
control became much more complicated. Due to the parallel execution of multiple
operations, dependencies among successive operations may occur. In the following
sections we will present the dependencies of the pipelined queue manager.
9
the read command is sent to the memory controller via the queue management interface process.
Datapath & Queue Management Chip Architecture
3.4.1
42
Successive Enqueue and Dequeue Operations for the same flow
A major effect of pipelining is to change the relative timing of queue manager
operations by overlapping their execution. This introduces data hazards. Data
hazards occur when the pipeline changes the order of read/write accesses to the
queue manager data structures so that the order differs from the order seen by
sequentially executing operations on an unpipelined architecture. More precisely, an
enqueue operation reads the queue tail address during its issuing phase, and later, it
updates the new queue tail address during its execution phase. Since the time
interval between the operation issuing and execution may last more than one time
slot, the queue tail address may be non-updated/pending during this interval. When a
newly issued enqueue operation finds the queue tail address as pending, data
dependence occurs. Similar to the enqueue operation, successive dequeue operations
of the same flow may be dependent because the latter operation tries to read the
queue head address while the former has not updated it yet. These data dependencies
introduce stalls in the pipeline and no farther operations can be issued until the data
dependencies are removed. This condition decreases the pipeline’s performance and
must be eliminated.
3.4.2
Successive Enqueue Operations of packet segments
Because of the queue manager manipulation on fixed size units, a variable size
packet enqueueing is split into multiple fixed size segment enqueueing. Instead of
issuing an enqueue operation per packet arrival, many enqueue operations must be
issued. The initiator of an enqueue operation can be whether the header processor or
the queue manager. In the case that the header processor initiates the enqueue
operations for all the packet segments, it should keep data structures for the
incoming segments as the figure 3.13 shows. This approach increases the
complexity of the header processor chip and introduces tasks such as the keeping of
data structures, which suit better to the queue manager functionality. Alternatively,
another approach is that the header processor initiates the enqueue operation for the
first segment of a packet and the queue manager initiates the enqueue operations for
the remaining segments of a packet. It implies that the queue manager keeps data
structures for the packet segments in order to initiate the proper number of enqueue
operations. Since all these enqueue operations are performed to the same queue, data
dependencies among the successive operations occur. These data dependencies are
equivalent to the data dependencies explained in section 3.4.1 and must be handled
with the same way.
Figure 3. 11 Successive Enqueue Operations of packet segments
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Pipeline Dependencies Handling
43
3.4.3
Buffer Memory Module Dependencies
The enqueue and dequeue operations are two parallel and independent operations.
The enqueue operation writes a segment to the buffer memory, while the dequeue
operation reads a segment from the buffer memory. If both operations try to access
simultaneously the same RIMM module, then a memory module conflict occurs. In
that case one of the two operations must be delayed, which implies that the achieved
operation rate will be inferior to the required (26Gbps). We propose techniques,
which handle this issue effectively and provide full memory throughput utilization.
3.5 Pipeline Dependencies Handling
3.5.1
Operands Renaming (Tomasulo) [15, chapter 4], [16]
The data dependencies occurred by the successive enqueue/dequeue operations to
the same flow are overcome by using operand-renaming techniques. The operand
renaming techniques are originated in the Tomasulo dynamic scheduling processors.
This technique assigns an identifier to each issued operation. In the case that a newly
issued operation cannot achieve a resource because a predecessor operation has got
it but has not updated it yet, the latter operation acquires the identifier of the
operation that will update the resource value. As soon as, the former operation
updates the resource the latter acquires the updated resource’s value. An interesting
issue is how a pending operation learns that the expected resource is available. The
Tomasulo dynamic scheduling technique uses a communication bus that informs any
pending operation. This solution will be expensive and infeasible at high speed.
Instead, we propose a technique that organizes the dependent operations in a
pending list. As soon as, an operation updates the resource, it informs the next
pending operation in the list with the resource value. Additionally, when successive
operations access the same resource, only the last one is actually used to update the
resource. Each intermediate operation updates directly its next pending operation
instead of the resource. In other words the expected operand is not accomplished by
the original resource but is forwarded by the operation that produces it. It is known
that forwarding/bypassing techniques eliminate pipeline stalls and improve the
pipeline performance.
Figure 3. 12 Pending Lists
3.5.2
Applying Operand Renaming Techniques to the Queue Management
Architecture
As we mentioned above, dependent enqueue or dequeue operations are organized in
Datapath & Queue Management Chip Architecture
44
pending lists. Each operation in the list forwards the required resource values to the
next pending operation in the list. The pending lists for the enqueue operations are
kept in the Pending Write Table, while the pending lists for the dequeue operations
are kept in the Pending Read Table. There are so many pending lists in the operation
tables as the number of active flows in the system. Additionally, the pending
enqueue operations are categorized into two types: those initiated by the header
processor and those initiated by the queue manager (see 3.4.2 section). We remind
that the header processor initiates enqueue operations per packet10, while the queue
manager initiates enqueue operations per segment11. Due to the existence of two
types of pending enqueue operations, we keep two pending lists per active flow: per
packet pending list and per segment pending list, as figure 3.12 shows.
We explain how we can construct the enqueue pending lists in the PWT by using a
simple example. The figure 3.13a shows an incoming packet consisting of four
segments to wait the packet entry process to assign a transit_id to each segment and
to store them in the transit buffer. The entry process organizes the packet segments
in a pending list as figure 3.13b shows. We use two pointers in order to keep the
pending lists in the PWT: a pointer to the next segment, the “next segment pointer”,
and a pointer to the last segment, the “last segment pointer”. The next segment
pointer indicates the transit_id of the next packet segment (or alternatively the
transit_id of the corresponding pending enqueue operation in the PWT). The last
segment pointer indicates the transit_id of the last packet segment. Concluding, the
packet entry process organizes a segment list per packet arrival.
Figure 3. 13 Segment list per packet arrival
Organizing dependent successive packet enqueue operations in a pending list
requires merging of different segment lists (per packet) into a single pending
segment list. This linking operation is performed during the enqueue issuing process
because this process checks for dependencies among successive operations. We
explain the linking operation by using the example of the figure 3.14a. We consider
that an enqueue operation issuing has acquired the tail pointer of the Q1 (Q1: queue
identifier) from the Queue table and has not updated it yet. It modified the state of
the Q1 tail pointer to pending and it stores its transit_id to the Q1 tail pointer field
(operand renaming). If the header processor assigns a successive packet to the Q1
flow, the corresponding issued enqueue operation will find the Q1 state as pending
but it will know that the Q1 tail pointer value keeps the transit_id of the last enqueue
1
0 The header processor initiates the enqueue operation for the first segment of each packet
1 The queue manager initiates enqueue operations for the remaining segments of each packet
1
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Pipeline Dependencies Handling
operation, which has acquired the Q1 tail pointer. Then, the newly enqueue
operation is linked to the corresponding pending list in the PWT, as figure 3.14b
shows. Additionally, the newly issued enqueue operation leaves its transit_id to the
Q1 tail pointer field (in Queue table) in order to indicate to a successive enqueue
operation of the same flow that it was the last that accessed this field.
Figure 3. 14 Operand renaming technique for successive enqueue operations
Organizing dependent dequeue operation in a pending list in the PRT is presented in
the figure 3.15. Since the scheduler of segment departures initiates dequeue
operations per segment, the pending list requires only a pointer that indicates the
transit_id of the next pending dequeue operation. The figure 3.16 shows the function
of linking a newly issued pending dequeue operation. The figure 3.16a shows the
state of the pending list before the linking of the new dequeue operation, while the
figure 3.16b shows this state after the linking.
Figure 3. 15 Per-flow pending lists
Datapath & Queue Management Chip Architecture
46
Figure 3. 16 Operand renaming technique for successive enqueue operations
3.6 Queue Pointer Management & Architecture Modifications
3.6.1
Next-Pointers in the DRAM Buffer Memory
Economizing on the off-chip memories, we locate the next pointers in the buffer
memory. This idea could be accomplished, if a small fraction of each buffer was
dedicated to store the next pointer field. The next pointer size is 22 bits in order to
address the 222 (4 million) buffers accommodated in the two RIMM modules of
buffer memory; thus a next pointer field size of 32 bits is adequate and decreases the
segment size from 64 bytes to 60 bytes. Locating the next pointers in the buffer
memory also economizes on the chip pins count. However, this achievement comes
at the expense of increasing the number of memory accesses and the latency of the
queue manager operations. The following paragraphs address the drawbacks, which
caused by the next pointers locating in buffer memory.
3.6.2
Buffer Preallocation technique [29]
If the next pointers were located in a separate memory, an enqueue operation would
require an access to the buffer memory: the writing of the segment body in the
buffer memory. The linking of this segment to the proper queue is performed in
parallel with segment body writing. If the next pointers were located in the buffer
memory, an enqueue operation would require memory writes at two different
addresses: write the data field of an arriving segment, and write the next pointer
field of the previous segment on the queue, as figure 3.17 shows.
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Queue Pointer Management & Architecture Modifications
Figure 3. 17 No free buffer preallocation
We overcome this drawback by using the buffer preallocation technique. In that
technique, each queue reserves one free buffer, which is the buffer to be used next;
thus ,a segment is always enqueued into the reserved buffer of the target queue, and
a newly extracted buffer is reserved for the next time. A pointer to the newly
reserved buffer is written into the current memory buffer, along with the data of the
arriving segment. By using the buffer preallocation technique we reduce the two
memory writes to one. The figure 3.18 shows how this technique achieves reduction
the required memory accesses per enqueue operation.
Figure 3. 18 Buffer preallocation
Datapath & Queue Management Chip Architecture
3.6.3
48
Link Throughput Saturation
Another drawback of the next pointers placing in the buffer memory is the
dequeueing latency increment. If the next pointers were located in a separate
memory, the queue’s head update would cost the off-chip memory access delay and
it would be performed in parallel with the segment’s reading from the buffer
memory. Instead, if the next pointers were located in the buffer memory, the queue’s
head update would cost the Rambus memory access latency (80ns). The dequeue
operation latency increment has negative effect to the queue manager performance.
Successive dequeue operations for the same flow are data dependent because they
expect the head pointer from the same queue; thus, each one will stay pending at
least the Rambus memory access latency. We overcome this dequeue operation’s
latency increment by interleaving dequeue operations for different flows. Due to the
Rambus memory access latency, which lasts four time slots, we can interleave
dequeue operations from four different flows without losing memory throughput. It
implies that if there were four active flows in the system they yield full link
throughput saturation. Less than four active flows receive one fourth of the link
throughput. This problem could be handle only if the next pointer were placed in a
separate memory.
3.6.4
Free List Bypassing technique [29]
If the next pointers were located in a separate memory, during an enqueue or a
dequeue operation, the free list update is performed in parallel with the segment
writing or reading to/from the buffer memory. This parallelism could be achieved
because the free list update requires accesses to the next pointers, which are located
in a separate off-chip memory. By locating the next pointers in the buffer memory,
the free list update requires accesses to the buffer memory, which implies that
during an enqueue or a dequeue operation the buffer memory accesses are increased.
Figure 3. 19 Read and Write transactions of an enqueue and a dequeue operation at
the same time slot
More precisely, in the case of an enqueue operation that uses the buffer preallocation
technique, two memory accesses must be performed to the buffer memory: the
writing of the segment body to the reserved buffer of the target queue, and the
reading of the next pointer field of the newly extracted buffer in order to accomplish
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49
Queue Pointer Management & Architecture Modifications
the address of the next free buffer in the free list. Another issue, which rises now, is
that the two buffer memory accesses must be directed to different RIMM modules in
order to avoid module memory conflicts. It implies that the newly extracted free
buffer and the reserved buffer must belong to different memory modules. This issue
is addressed in the section 3.6.5. In the case of a dequeue operation, two memory
accesses must be performed: the reading of the departing buffer body and the
linking12 of this buffer to the free list. The figures 3.19 shows the required buffer
memory accesses when an enqueue operation to the Q2 flow and a dequeue
operation to Q1 flow take place.
As mentioned in section 3.4.1, the queue manger must perform an enqueue and a
dequeue operation per time slot. By locating the next pointers in the buffer memory,
four memory accesses must be performed to the buffer memory per time slot. It
implies that the required memory throughput is at least twice the provided. Using the
free list bypassing technique could reduce this memory throughput requirement. In
this technique, rather than dequeueing a departing buffer from an output queue and
enqueueing that buffer into the free list, and rather than extracting a buffer from the
free list and enqueueing it into another queue upon arrival, we combine the two
operations: the buffer into which an arriving segment is placed, is precisely the
buffer from which a segment is departing during the same time slot. Therefore, there
is not free list operation, which implies that the required memory throughput equals
to the provided throughput. The figure 3.20 shows the reduction to the number of
memory access by using the free list bypassing technique.
Figure 3. 20 Free List Bypassing (memory transactions)
3.6.5
Per-memory bank Queueing Free List Organization
As referred before, during an enqueue operation, a writing access to the reserved
buffer and a reading access to the newly extracted free buffer must be performed at
the same time slot. In order to perform both memory accesses simultaneously, these
accesses must be performed to different RIMM modules. However, if the free list
was organized as a single queue, the above requirement might not be achieved
1
2This linking is performed by writing the buffer address to the next pointer field of the free list tail
buffer
Datapath & Queue Management Chip Architecture
50
because the free buffer at the head of the free list may belong to the same RIMM
module with the reserved buffer of the targeted queue. In order to ensure that the
extracted free buffer and the targeted queue reserved buffer belong to different
RIMM modules, we organize the free list in a per-bank queueing scheme. This
organization also ensures that the extracted buffer does not belong to a busy bank.
The only expense of the above organization is the requirement for keeping the head
and tail pointers of the 512 13 different bank-queues in the Free List table.
However, the per-bank queueing organization complicates the free list update for a
dequeue operation. During a dequeue operation the departing buffer must be
enqueued to the proper free list bank-queue. The issue is that this queue is located in
the same memory module as the departing buffer, which means that during a
dequeue operation both memory accesses (read and write) must be performed to the
same memory module. Both accesses cannot be performed simultaneously; thus, one
of them must be delayed. Since the reading access to the departing buffer has higher
priority than the free list update, the free list update is delayed for later.
3.6.6
Free Buffer Cache
In order to handle the occurred drawback of a dequeue operation due to the free list
organization we use caching techniques. During a dequeue operation, instead of
enqueueing the departing buffer to the free list, we keep its address in a cache; we
call this cache “free buffer cache”. The free buffer cache is a pool of “isolated” 14
buffers. The introduction of the free list cache modifies the free list bypassing
technique in our system. When an enqueue and a dequeue operation are performed
simultaneously, the enqueue operation extracts a free buffer from the free buffer
cache and the dequeue operation inserts the departing buffer to the free buffer cache.
This modification requires the free buffer cache to have at least one buffer, in order
to be extracted during the enqueue operation. During the dequeue operation a new
buffer is inserted to the cache. However, in the case of successive dequeue
operations without enqueue operations, the free buffer cache may be overflowed.
We overcome this overflow by enqueueing the isolated buffers in the proper free list
queue when no enqueue operation is performed.
The packet entry process handles the enqueueing operations of the cached buffers
(independent buffers that is kept in the cache). When there is no incoming segment
waiting to be stored in the transit buffer, the packet entry process issues a write
operation in the PWT, which is originated by an isolated buffer enqueueing to the
free list.
3.7 The Overall Queue Management Architecture
Throughout this chapter, we described the basic structures of the queue management
architecture. It uses the Rambus DRAM technology for the buffer memory, which
requires out-of-order memory accesses for full throughput utilization. We introduced
dynamic scheduling techniques for reordering memory accesses and designed the
architecture in a pipelined fashion in order to improve the operation rate15. Next, we
1
3both RIMM modules contain 512 banks
1
4Isolated buffer is a buffer that is not linked to the free list
5 An enqueue and a dequeue operation per time slot
1
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51
The Overall Queue Management Architecture
detected the pipeline dependencies and we handled them by using operand-renaming
techniques. We placed a portion of the queue management data structures, the next
pointers, in the buffer memory in order to economize on the external memories and
on the pins count. However, the location of the next pointers in the buffer memory
increases the number of memory accesses per enqueue or dequeue operation and
enlarges the latency of these operations. We handled this drawback by using buffer
preallocation and free list bypassing techniques. In order to make our architecture
more flexible, we organized the free list in a per-bank queueing scheme and we used
caching techniques. The entire architecture is illustrated in the figure 3.21. A brief
description of the presented architecture is given in the following paragraphs.
When a segmented packet arrives at the input, it is kept in the transit buffer and a
transit_id is assigned to each packet segment. During packet buffering the
corresponding entries in the Pending Write Table are allocated; the packet segments
are organized in a list by writing the next segment and last segment fields of the
allocated entries in the PWT. Concurrently, the packet header fields, which are
located in the first 64-byte packet segment, along with the transit_id of the first
packet segment are transmitted to the header protocol processing chip for packet
routing and classification.
As soon as the header processor assigns an incoming packet to a flow, it issues an
enqueue operation to the flow by sending the transit_id of the packet (the first packet
segment) and the flow_id of the targeted flow. Then the queue manager extracts the
tail pointer of the corresponding queue, by accessing the tail field in the Queue table,
and stores it to the write address field of the PWT entry indexed by the transit_id of
the first packet segment. If the queue tail field in the Queue table is pending, it keeps
the transit_id of the last enqueue operation that has acquired the queue tail pointer
value. In that case, the newly issued enqueue operation is linked to the pending list
of the corresponding flow in the PWT.
In order to support out-of-order memory transactions, a search engine traces the
pending enqueue (write) operations in the PWT, in parallel, and selects an enqueue
operation that will not cause a memory bank conflict. Then, it extracts a free buffer
in order to reserve it for the next enqueue operation (buffer preallocation). The free
buffer can be extracted by two sources: the free list, and the free buffer cache. If a
dequeue operation is performed at the same time slot with the enqueue operation
then the free buffer is extracted by the free buffer cache due to the free list bypassing
technique. Otherwise, if only an enqueue operation is performed at the current time
slot, the free buffer will be extracted from the free list. As soon as, the free buffer
address is accomplished, the queue manager sends a write transaction to the memory
controller; this transaction writes the segment body to the queue tail (reserved)
buffer and writes the address of the newly extracted free buffer to the next pointer
field of the same buffer. Additionally, the queue manager writes the free buffer
address to the queue tail field in the Queue table, in order to update the new queue
tail pointer.
On the other hand, as soon as the scheduler decides to forward a segment from an
active flow, it issues a dequeue operation by sending the flow_id the the serviced
flow. Then the queue manager acquires the corresponding queue head pointer, by
accessing the Queue table. As soon as, the queue head pointer is accomplished, it is
stored in read address field of a PRT entry. If the queue head field in the Queue table
is pending, the newly issued dequeue operation accomplishes the transit_id of the
Datapath & Queue Management Chip Architecture
52
last operation that accessed this field (operand renaming) and is linked to the
corresponding pending list in the PRT.
Similar to the enqueue operation, a search engine traces the pending dequeue/read
operations in the PRT and selects a dequeue operation that it will not cause a
memory bank conflict. Then, the queue manager sends a read transaction to the
memory controller. This read transaction will retrieve the segment body of the
departing buffer and the next pointer value, which is kept in the departing buffer and
indicates the next buffer of this queue. As soon as the buffer memory responds with
the data of the departing buffer, the segment body is forward to the switching fabric,
while the next pointer field updates the head pointer field of the corresponding
queue in the Queue table. Additionally, the address of the departing buffer is stored
in the free buffer cache.
The queue management architecture, which is described above, achieves highoperation rates (an enqueue and a dequeue operation per time slot <time slot =
40ns>) by using advanced pipelining and by applying dynamic scheduling
techniques originated in the supercomputers in 60’s. The detailed micro-architecture
of the queue management block is performed in the chapter 4.
Figure 3. 21 Mutli-Queue Rambus Controller block diagram
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
4 Queue Management Micro-Architecture
4.1 Hardware Implementation of the QM Data Structures
The queue management subsystem uses data structures, such as the Queue table, the
next pointers, the free list, the free buffer cache, the operation tables, such as the
pending write table and the pending read table, and the read-write control/data
buffers in order to manipulate the incoming and the outgoing traffic at 10Gbps link
rate. This section describes in detail the queue management memory blocks.
4.1.1
Queue Table
The queue management subsystem supports up to 64K active flows. Due to per-flow
queueing organization the head and tail pointers of the 64K queues are kept in the
Queue table, as figure 4.1 shows. This table is split into two separate tables: the
Head table and the Tail table. This isolation yields parallel accessing to both queue
head and tail pointers. Since, the enqueue and the dequeue operations are two
independent and parallel operations, they may perform concurrent accesses to the
Queue Table. The Head table, apart from the head pointer field, keeps two state
flags: the Pending Head (PH) flag and the Almost Ready (AR) flag. The PH flag
indicates whether the state of an entry (the queue head pointer state) is updated or
pending. The AR flag is set only if the PH flag is set. The AR flag indicates that the
dequeue operation, which makes this head table entry pending, is in execution and it
is not presented in the Pending Read Table. Because of the Rambus memory high
latency and the location of the next pointers in the buffer memory, a read operation
may last at most four time slots (see section 4.2.8) until it responds with the results
for updating the queue head pointer. If a newly inserted dequeue operation finds the
AR flag to be 1, it should be handled differently than the original case (AR=0), see
section 4.2.6.
Figure 4. 1 Queue Table
Similar to the Head table, the Tail table consists of the tail field and its state flag: the
Pending Tail (PT) flag. The Tail table does not need an Almost Ready flag because
the update of the tail table entries (queue tail pointers) is performed during the
Queue Management Micro-Architecture
54
enqueue execution phase. The queue tail update does not require a buffer memory
access, because the new tail pointer is the pointer of the extracted free buffer from
the free list and is currently available. The size of the supported system queues is
64K flows; thus the Head and Tail tables maintain 64K entries. The Head table
width size is 24 bits: 1bit for the PH flag, 1bit for the AR flag and 22 bits for the
head pointer1. Respectively, the Tail table width size is 23 bits: 1bit for PT flag, and
22 bits for the tail pointer. Additionally, an empty (E) flag is required for each entry
in the Queue table in order to indicate whether a queue is empty or not.
4.1.2
Pending Write Table and Transit Buffer
The Pending Write Table (PWT) keeps control information for the pending write
transactions which are originated either by an enqueue or by a free list update
operation. The PWT is split into four isolated tables, as figure 4.2 shows: the state
table, the operand table, and the two pending list tables (the next segment table and
the last segment table). The state table contains three flags: Valid (V), Busy (B), and
Ready ® flag. The Valid flag indicates whether the corresponding entry is used or
not. The Busy flag indicates whether the corresponding write operation, which is
kept in the PWT entry, is in execution or it is waiting execution. The Ready flag
indicates whether the corresponding write operation has accomplished its operands
and it is ready for execution.
Figure 4. 2 Pending Write Table
The operand table contains three fields: the write address, the flow identifier and the
Free List Update (FLU) flag. The write address field keeps the address of the buffer
to which a waiting segment will be written. The flow_id field keeps the identifier of
the queue to which the waiting segment will be enqueued. The FLU flag indicates
whether the pending write operation originated by an enqueue operation or by a free
list update operation. The next segment table has two fields: the next segment field
and the next segment (NS) flag. The next segment field keeps the transit_id of the
next segment in the pending list, while the NS flag indicates whether the next
segment field has a valid value. The last segment table has two fields: the last
segment field and the last segment (LS) flag. The last segment field keeps the
We remind that both RIMM modules contains 222 ( 4 million) 64-byte buffers
1
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Hardware Implementation of the QM Data Structures
55
transit_id of the last segment in the pending list, while the LS flag indicates whether
this PWT entry keeps the enqueue operation, which is the last that has accessed the
queue tail field in the Tail table.
We define the last enqueue operation that has accessed the queue tail pointer as the
“server” of the corresponding pending list in the PWT. The name server means that
this enqueue operation will assist a successive enqueue operation for the same flow
to be linked in the pending list. Note that there is only one server for each pending
list. If a pending enqueue operation is the “server” of the pending list, it has to
update the Tail table when it is executed. As mentioned above, only the last
operation in a pending list updates the tail pointer fields in the Queue table. The
intermediate pending operations update only the next operation in the pending list.
The Pending Write Table contains 128 entries. The width of the state table is 3 bits,
the width of the operand table is 39 bits (22 for write address, 16 for flow_id and 1
for FLU flag), the width of next segment table is 8 bits (7 for transit_id and 1 for the
NS flag), and the last segment table width is also 8 bits (7 for transit_id and 1 for the
LS flag).
The Transit Buffer consists of 128 entries corresponding to 128 entries of PWT. The
size of each entry is 64-bytes (it stores a 64-byte segment). The 128 entries are kept
in a 16-bytes –width on-chip memory. The memory width-size is selected in order to
achieve 16 bytes memory access granularity. The 16-byte granularity achieves
writing and reading transaction rates of 12.8 Gbps (16 bytes / 10ns clock cycle).
Figure 4. 3 Pending Read Table
4.1.3
Pending Read Table
The Pending Read Table keeps control information for the pending read
transactions, which are originated by the dequeue operations. The PRT is split into
three separate tables, as figure 4.3 shows: the state table, the operand table, and the
pending list table (the next segment table). The state table contains three 1-bit flags:
Valid, Busy, Ready; these flags have the same functionality as the corresponding
flags in the state table in the PWT but they are referred to dequeue operations. The
operand table has two fields: the read address field and the flow_id field. The read
Queue Management Micro-Architecture
56
address field keeps the address of the departing buffer, while the flow_id field
contains the identifier of the serviced flow. The next segment table has identical
format with its counterpart in the PWT but it is referred to the pending list of
dequeue operations. The Pending Read Table maintains 128 entries. This length is
independent of the PWT length but it is a normal size for keeping information of a
significant number of pending operations. The width size for state table is 3 bits, for
the operand table is 39-bits (22-bits for the read address, and 16-bits for the
flow_id), and 8-bits for the next segment table (7 for transit_id and 1 for NS flag.
4.1.4
Free List Table and Free List Cache
The Free List table keeps the head and tail pointers of the 512 free buffer queues.
As mentioned in the section 3.6.5, the free list is organized in a per-bank queuing
scheme. The length of this table is 512 entries (two RIMM modules contain 512
banks). Each entry has three fields, as figure 4.4a shows: the head pointer, the tail
pointer and the Empty (E) flag that indicates whether a queue of free buffers is
empty or not. The width-size of each entry is 25 bits (1-bit for the empty flag, 22bits for head pointer, and 22-bits for tail pointer).
The Free List Cache (FLC) is a pool of free/unlinked/independent buffers. All these
buffers are originated by dequeue operations. The FLC is a FIFO on-chip memory.
Two pointers are required – the write and the read pointer- in order to point the
location of inserting a new free buffer or the location for extracting a free buffer
to/from the FLC, as the figure 4.4b shows. We implement the free list bypassing
technique using the FLC, see section 3.6.6. The length of FLC can be up to some
tens of entries and the width of an entry must be 22-bits as the address-size of a
memory buffer.
Figure 4. 4 Free List Table and Free Buffer Cache
4.1.5
Control and data Buffer
Due to the location of the next pointers in the buffer memory, a read transaction
retrieves the segment body and reads the associated next pointer field. The
accomplishing of the next pointer value implies the requirement to update the queue
management data structures (Queue table or Free List table) or a pending read
transaction in the PRT. So, each read transaction must be accompanied with
additional control information related to the update destination. This information is
kept in a control buffer that accompanies each read or write transaction from its
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57
The Pipelined Control Processes Micro-Architecture
generation until its completion. Because there are two memory RIMM modules, two
memory accesses may take place at each time slot; it implies that two control buffers
may be active, simultaneously.
Each control buffer consists of seven fields: the transaction type (R/W), the Valid
flag, the source of the transaction (table and the transit_id of the table entry that
keeps the transaction), the update destination (table and the identifier of the table
entry), and the transaction address. Note that, the destination update fields are valid
in the case of a read transaction. The control buffer size is 51 bits: 1-bit for the Valid
flag,1-bit for the operation type, 2-bits for the destination table, 16-bits for the
destination table entry (for the worst case of updating the Queue table), 2-bits for the
source table, 7-bits for the source table entry –PWT or PRT-, and 22-bits for the
operation address. The format of the control buffer is illustrated in the figure 4.5.
Figure 4. 5 The Control Buffer format
In the case of a write transaction, we move the segment body from the transit buffer
to a local buffer near to the memory controller. The memory controller splits the
segment body writing into four phases because it pipelines the memory transactions.
By keeping the segment body in a local buffer near to the memory controller, we
ensure that the segment body will be available for accessing from the memory
controller. We call this local buffer as “Data Buffer” and its size equals to the
segment body size (64-bytes).
4.2 The Pipelined Control Processes Micro-Architecture
As mentioned above, the queue manager architecture is composed of six parallel and
fully pipelined processes: packet fetching, enqueue issuing, enqueue execution,
dequeue issuing, dequeue execution and queue manager interface process. The
pipeline stage length for all the processes equals to the time slot duration in order to
perform an enqueue and a dequeue operation per time slot (40ns). The queue
manager architecture operates with a clock speed of 100MHz, the frequency
required to support the 12.8Gbps rate over 16-byte data path. This clock frequency,
which is conservative for the 0.18-micron technology, simplified our logicpartitioning and pipelining tasks; one access to an on-chip memory plus several
levels of combinational logic fit within a clock cycle with relative easy. Each
pipeline stage has latency quadruple the clock cycle period; we will refer to each
clock cycle in a pipeline stage as the first, second, third, and forth cycle. The
following sections describe the six –pipelined processes thoroughly. ?
4.2.1
Packet Fetching Process Micro-Architecture
This process works in two modes. The first mode initiates an enqueue operation by
temporarily storing an incoming segment to the Transit Buffer and allocating the
corresponding entry to the pending write table. It also organizes the incoming
?
The transactions that take place at the first, second, third, or fourth cycle, are represented in the
figures of this chapter with black, red, blue, and green colored arrows, respectively
Queue Management Micro-Architecture
58
segments of a packet into a single linked list. The information for the lists is kept in
the next and last segment fields of the pending list table, which are parts of the
PWT. The second mode initiates and simultaneously issues a free list update
operation by allocating an entry in the PWT and achieving the proper operands for
this task.
The description of the first mode is illustrated in the figure 4.6. The inputs of this
process are the transit_id of a free entry in the Transit Buffer and an incoming
segment body. During the first cycle the first 16-bytes-part of the segment body is
stored to the Transit Buffer indexed by the transit_id. The PWT entry indexed by the
transit_id is also allocated by setting the valid flag to 1. If the incoming segment was
the first of a packet it writes its transit_id in the head and middle registers. The head
register stores the transit_id of the first segment of a packet, while the middle
register stores the transit_id of each intermediate segment. The information in the
head and middle registers is used by the subsequent segments of the same packet in
order to be linked in the pending list, see the section 3.5.2. If the incoming segment
were the packet intermediate segment, it would achieve the transit_id of the previous
segment from the middle register and then it would be linked to the pending list by
writing its transit_id to the previous segment next segment field. If the incoming
segment were the packet last segment, it would achieve the transit_id of the previous
segment and the transit_id of the first segment from the middle and head register
respectively. Then, it would be linked to the list by updating the next segment field
of the previous segment, and it would update the last segment field of the first
segment. By this way we organize the incoming segments into pending lists during
their arrivals.
Figure 4. 6 Packet fetching process block diagram (mode 1)
During the second cycle, the second 16-bytes-part of the segment body is stored to
the Transit buffer entry indexed by the transit_id augmented by one. Additionally, a
search engine searches for the next free entry in the PWT and holds its transit_id.
During the third and forth cycle, the third and forth 16-bytes-part of the segment
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
59
The Pipelined Control Processes Micro-Architecture
body is stored to the Transit Buffer entry indexed by the current transit_id
augmented by two and by three, correspondingly.
The description of the second mode is illustrated in the figure 4.7. The inputs of the
packet fetching process for this mode are the transit_id of the current free entry in
the Transit Buffer and the address of the free buffer from the Free List Cache
(FIFO). During the first cycle, the address of the free buffer is accomplished and it is
kept in a register. During the second cycle, the address of the tail buffer at the free
list is acquired by accessing the Free List Table. Similar to the first mode, in this
cycle a search engine accomplishes the next free entry in the PWT. During the third
cycle the free list update task operands are written to the proper fields in the PWT.
More precisely, the fields of the PWT entry indexed by the current transit_id is
updated as follow: the flags Valid, Busy, Ready are set to 1; the address of the tail
free buffer is written to the write address field; the flow_id field is not updated
because it is not used; the fields next and last segment are set to zero in order to
indicate that this entry does not belong to a pending list; the FLU flag is set to 1 in
order to indicate that this operation is a free list update operation. Additionally, the
address of the free buffer is written to the 32 most significant bits of the Transit
Buffer entry indexed by the current transit_id. Finally, the value of the tail pointer of
the free list Table is updated with the free buffer address. The forth cycle is idle. The
latency of the packet fetching process equals to a time slot, so its pipeline has only
one stage. A block diagram of this process is presented in the Appendix B.
Figure 4. 7 Packet fetching process block diagram (mode 2)
4.2.2
Enqueue Operation Issuing Process Micro-Architecture
The main task of this process is to issues an enqueue operation and to collect its
operands. An enqueue operation issuing begins when the header processor assigns a
flow_id to the first segment of a packet. The inputs of this process are the flow_id
and the transit_id of the first segment. During the first cycle, we access the Tail table
entry indexed by the flow_id, in order to accomplish the queue tail pointer and its
state. Depending on the queue tail pointer state, we follow two different datapaths.
Queue Management Micro-Architecture
60
The datapath in the case that the state is not pending is presented in the figure 4.8.
During the second cycle, the proper queue tail pointer is written to the write address
field at the PWT entry indexed by the input transit_id and the Ready flag is set to 1.
The flow_id field is also updated by the input flow_id, while the FLU flag is set to 0
to indicate that this entry keeps an enqueue operation. During the third cycle, the
transit_id of the newly issued enqueue operation is written to the corresponding
queue tail pointer field (operand renaming), while the PT flag is set to 1 (it indicates
pending state). This access informs a successive enqueue operation of the same flow
that the queue tail state is pending and the tail field keeps the transit_id of the last
enqueue operation that has accessed this field. During the fourth cycle the LS flag in
the last segment table is set to 1 in order to indicate to the kept operation of this
entry that it was the last that has accessed the queue tail pointer, and it has to update
the new queue tail pointer or to forward this value to the successive pending
enqueue operation.
Figure 4. 8 Enqueue issuing process datapath (not-pending state)
The datapath in the case that the queue tail state is pending state is presented in the
figure 4.9. If the state of the tail pointer is pending then the queue tail pointer field
keeps the transit_id of the last enqueue operation that has accessed the queue tail.
The entry of PWT that stores this enqueue operation keeps information for the
pending list to which it belongs. During the second cycle, we access the last segment
field of this entry in order to acquire the transit_id of the last entry in the pending
list. Accomplishing this information, we link the current inserted enqueue operation
to the tail of the corresponding pending list during the third cycle. The linking is
performed by writing the transit_id of the newly issued enqueue operation to the
next pointer field of the last entry of the pending list. During the same (third) cycle,
the transit_id of the newly issued enqueue operation is written to the queue tail
pointer field of the Queue table, and the queue tail state remains pending. Now, the
“server” of this pending list will be the newly issued enqueue operation. The Queue
table knows this information, but we have to inform the enqueue operation itself that
it is the server by setting the LS flag of the PWT entry that keeps the newly issued
enqueue operation to 1 and by resetting this flag of the previous server enqueue
operation. The former of these tasks is performed in the third cycle while the latter
in the forth cycle. The latency of this process equals to a time slot, so its pipeline has
only one stage. A block diagram of this process is presented in the Appendix B.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
The Pipelined Control Processes Micro-Architecture
61
Figure 4. 9 Enqueue issuing process datapath (pending state)
4.2.3
Enqueue Execution Process Micro-Architecture
The main task of this process is to select an eligible (non-conflicting) write
transaction from the PWT in order to execute it. The write transaction may be
originated either by an enqueue operation or by a free list update operation. The
information for the origination is kept to the FLU flag in the PWT. The operation
selecting is performed at run time in order to dynamically schedule the eligible write
operations and to avoid memory bank conflicts. In order to design a more flexible
scheme, this process tries to extract two eligible write operations that they are
directed to different memory modules. Simultaneously, the counterpart process of
dequeue operations –we analyze it later- tries to extract two eligible read operations
that they are directed to different memory modules. By means of a combinational
logic circuit, we select a write and a read operation, which are directed to different
memory modules to utilize the provided memory throughput more effectively. The
following paragraphs analyze this process in the level of the clock cycle accuracy.
This process pipeline consists of two stages.
First Stage of the pipelined process
This paragraph describes the transactions that take place in the first stage of this
process. Figure 4.10 shows the transactions of this stage. During the first cycle, two
parallel search engines seek two write operations that are eligible for execution and
are directed to different memory modules. The search engines respond with at most
two transit_ids, which correspond to two write operations. Note that, the operation
eligibility means that the operation does not cause memory bank conflict. During the
second cycle both the extracted transit_ids are stored to the two share-accessed
registers, which we call them as “write_tr_id1” and “write_tr_id2”. During the same
cycle, the counterpart dequeue process stores the extracted read_tr_ids to the two
share-accessed registers, “read_tr_id1” and “read_tr_id2”. Simultaneously, the write
address, the flow_id and the FLU flag of the PWT entry, indexed by the
write_tr_id1, are acquired and then they are stored to a local register. The
corresponding fields of the enqueue operation, which is kept in the PWT entry
indexed by the write_tr_id2, are acquired and stored to another local register during
the third cycle. A combinational logic circuit, which we call it as “dynamic
Queue Management Micro-Architecture
62
scheduler”, has four inputs: the write_tr_id1, the write_tr_id2, the read_tr_id1, and
the read_tr_id2; the dynamic scheduler chooses the transit_ids of a read and a write
transaction that they are directed to different memory modules. At the end of the
third cycle the transit_id of the selected write transaction is available. During the
forth cycle, we collect the remaining information related to the selected write
transaction from the PWT. More precisely, we learn if this write operation belongs
to a pending list, if it has successive operations in the list or it is the last one, or
alternatively, if it is the “server” of the pending list. This information is
accomplished by accessing the next and last fields of the corresponding entry. The
results of the access are kept to some temporary registers. Simultaneously, the Busy
flag of this entry is set to 1 in order to indicate that this operation is in execution.
Finally, during the forth cycle of the first stage, a free buffer is extracted; Note that
the buffer must not belong to a busy bank.
Figure 4. 10 Enqueue Execution process (first stage)
The operation of extracting a new free buffer is presented in the figure 4.11. There
are two sources for extracting a free buffer: the Free List and the Free Buffer Cache.
The Free List keeps buffers organized in queues, while the Free Buffer Cache keeps
independent (unlinked) buffers. The choice of the free buffer source is dependent on
the dynamic scheduler results. If a write transaction –originate from an enqueue
operation- and a read transaction – originate from a dequeue operation- were
selected to be performed at the next time slot, concurrently, the free buffer would be
extracted from the Free Buffer Cache due to the free list bypassing technique.
Otherwise, the free buffer would be extracted from the Free List. In the case of
extracting a buffer from the Free List, this buffer must belong to a non-busy bank,
because it would be accessed as mentioned in section 3.6.4. By accomplishing this
constraint, we activate a search engine to find a buffer from a non-busy bank; it is an
easy task because the Free List is organized in a per-bank queueing scheme and we
can extract an eligible buffer at O(1).
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
The Pipelined Control Processes Micro-Architecture
63
Figure 4. 11 Free buffer extraction
Between the first and the second stage we use a buffer in order to isolate the two
stages; we call this buffer as “pipeline buffer”. The pipeline buffer consists of 10
fields: the transit_id (7 bits), the write address (22 bits), the flow_id (16 bits), the
FLU (1 bit), the next segment (7 bits), the NS flag (1 bit), the last segment (7 bits),
the pending list server flag (1 bit), the free buffer address (22 bits), and the free
buffer source (1 bit). The pipeline buffer contains all the necessary information for
the selected write transaction. This information will be used in the second stage that
executes the write transaction.
Second Stage of the Pipelined Process
The second stage of the enqueue execution process is the stage of executing the
previously selected write transaction (it is selected in the first stage). This process
handles two different cases dependent on the write transaction origination: enqueue
operation or free list update.
The case that the write transaction originates from an enqueue operation is also
subdivided into two instances. The first instance is when only an enqueue operation
is performed during a time slot, while the second instance is when an enqueue and a
dequeue operation are performed during the same time slot. If only an enqueue
operation was executed in the current time slot (there was no dequeue operation),
then a write and a read transaction would be performed to different memory
modules. The write transaction writes the segment body to the buffer memory, while
the read transaction is performed to the next pointer field of the extracted free buffer
in order to accomplish the address of the next free buffer in the free list. Otherwise,
if an enqueue and a dequeue transaction were executed in the current time slot, then,
a write and a read transaction would be executed to different memory modules. The
write transaction stores an incoming segment body to the buffer memory, while the
read transaction retrieves the departing segment body from the buffer memory. Due
to the free list bypassing, any operation related to the free list occurs. Concluding,
the second stage of the enqueue execution process performs a write and a read
transaction in the first instance, while it performs only a write transaction in the
second instance. The accommodated control information for a write and a read
Queue Management Micro-Architecture
64
transaction is written to the corresponding control buffers, during the first cycle of
this pipeline stage, as figure 4.12 shows. The information that is kept to the control
buffers of the write and read transaction is presented in the tables 4.1 and 4.2,
respectively. Additionally, the first 16-bytes part of the segment body is moved from
the Transit Buffer to the data buffer during the first cycle.
Valid
1
Valid
1
R/W
W
R/W
R
Dst.Tbl
idle
Dst.Tbl.
Free List Table
Dst.Entry
idle
Dst.Entry
free buffer addr[21:13]
Src.Tbl.
PWT
Src.Tbl.
PWT
Src.Entry
current Transit_id
Src.Entry
current Transit_id
oper. addr.
Write address (from PWT
entry indexed by transit_id)
oper. addr.
Free buffer addr[21:0]
Table 4. 1 Control buffer for a write trans
Table 4. 2 Control buffer for a read trans.
Figure 4. 12 Second stage (execute an enqueue operation)
If the write transaction, which is in execution, belonged to a pending list, it should
forward the new queue tail pointer, which is the address of the extracted free buffer,
to the proper destination. We examine two cases: the write transaction of an
intermediate entry in the pending list or the write transaction of the last entry in the
pending list. If it is an intermediate entry of the pending list, it has to forward the
queue tail pointer to the write address field of its next entry in the pending list. This
operation is performed during the first cycle. Otherwise, if the write transaction is
the last entry in the pending list, it has to write the new queue tail pointer to the
Queue table. This operation is performed during the second cycle. During the
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The Pipelined Control Processes Micro-Architecture
65
second, third, and forth cycle, the second, third, and forth 16-byte part of the
segment body are moved from the Transit Buffer to the data buffer, respectively
(write transaction). At the end of this stage, the entry of the PWT, which keeps the
executed write transaction, is released.
In the case that the write transaction originates from a free list update operation, the
accompanied control information of this transaction is loaded to the corresponding
control buffer during the first cycle. The information that the control buffer keeps
for the write transaction is performed in the table 4.3. This write transaction writes
the address of an unlinked free buffer to the next pointer field of the free list tail
buffer. The address of the unlinked free buffer is kept in the 32 most significant bits
of the corresponding transit buffer entry. The data of this transit buffer entry are
loaded to the data buffer during the four cycles of the second stage.
Valid
1
R/W
W
Dst.Tbl.
Idle
Dst.Entry
Idle
Src.Tbl.
PWT
Src.Entry
current Transit_id
Oper.
addr.
Write address (from
PWT entry indexed by
transit_id)
Table 4. 3 Control buffer for free list update (write transaction)
4.2.4
Handling Exceptional Cases during an Enqueue Operation
An exceptional case occurs when at the current time slot a newly issued enqueue
operation finds the state of the queue tail as pending, while the “server” enqueue
operation of the corresponding pending list in the PWT has started its execution
during the previous time slot and it has not completed yet (the latency of the
execution lasts up to four time slots). Even if the newly issued operation knows the
transit_id of the “server” operation of the proper pending list, the server operation is
in execution and it has not updated the corresponding queue tail, yet. This situation
causes an exceptional case where the newly issued operation can not accomplish the
required resources (the queue tail pointer) and must be handled differently.
This exceptional case is handled by using bypassing techniques. The newly issued
enqueue operation reads the pipeline buffer of the “Enqueue Operation Execution
Process”. If the transit_id, which is kept in the pending queue tail pointer field in the
Tail table, is identical with the transit_id field in the pipeline buffer, an exceptional
case has detected. In this exceptional case there are two routes of bypassing. If the
“server” operation, which is in execution, was the last entry in the pending list, then
the address of the extracted free buffer is bypassed from the pipeline buffer to the
write address field of the PWT entry that keeps the newly issued enqueue operation.
This bypassing forwards the new queue tail pointer before it is written to the Tail
Table. Otherwise, if the “server” operation was an intermediate entry in the pending
list it forwards the transit_id of the last entry in the pending list. The newly issued
operation can be linked at the tail of the pending list by using this information.
Queue Management Micro-Architecture
4.2.5
66
Dequeue Operation Issuing Process Micro-Architecture
The main task of this process is to issue a dequeue operation and to acquire the
appropriate dequeue operands. The only operand is the address of the buffer at the
head of the corresponding queue. The queue head pointer is kept in the Head Table.
The inputs of this process are the identifier of the flow that the scheduler services
the current time slot and the transit_id of a free entry in the Pending Read Table.
Figure 4. 13 Dequeue Operation Issuing Process
The datapath of this process is illustrated in the figure 4.13. The first cycle of this
stage is idle (resource accesses scheduling, see the section 4.2.9). During the second
cycle, the flow_id indexes the Head table entry that keeps the queue head pointer
and its state fields (PH and AR flags). Simultaneously, a search engine looks for the
next free entry in the PRT. If the PH flag is set to 0, it implies that the acquired
queue head pointer is updated (correct). So, the read address and flow_id fields in
the PRT entry of the newly issued operation are written with the correct values. This
operation is ready for execution to a successive time slot; thus, the Ready and Valid
flags of this entry are set to 1. Otherwise, if the PH flag is set to 1, we examine two
cases: AR flag is set to 0 or AR is set to 1. The former case means that the “server”
of the corresponding pending list in the PRT is located in the PRT entry indexed by
the value of the acquired head pointer (operand renaming). The server of a pending
list in the PRT is the last dequeue operation that accesses the queue head pointer
field at the Head Table. The newly issued dequeue operation, by accomplishing this
information, is linked to the tail of this pending list. It is performed by writing its
transit_id to the next segment field of the pending list tail entry (in the PRT). The
latter case means that the “server” operation of the corresponding pending list is in
execution. It implies that the “server” read operation is send to the buffer memory,
but the buffer memory has not responded with the results and the queue head pointer
is still pending.
4.2.6
Dequeue Operation Execution Process Micro-Architecture
This process pipeline consists of two stages. The main task of the first stage is to
select an eligible read operation and to prepare it for execution. The second stage
undertakes the operation execution and the update the queue manager data
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The Pipelined Control Processes Micro-Architecture
67
structures.
First Stage of the Pipelined Process
The data path of the first stage is performed in the figure 4.14. During the first
cycle, two search engines search in parallel in order to find an eligible (notconflicting) read operation for both the memory modules. The results of this
searching are stored to the two share-accessed registers, read_tr_id1 and read_tr_id2,
in order to be accessible by the enqueue operation execution process. More
precisely, the couple of the search engines respond with at most two transit_ids,
which correspond to two eligible read operations kept in the PRT. During the second
cycle, the read address and the flow_id fields of the PRT entry indexed by the first
transit_id, are acquired and stored to local registers. During the third cycle, the read
address and flow_id fields of the PRT entry indexed by the second transit_id are
stored to local registers. Similar to the enqueue operation execution process, during
the third cycle, the dynamic scheduler circuit responds with the transit_ids of one
write transaction from the PWT and a read transaction from the PRT, which are
directed to different memory modules. During the forth cycle, the transit_id of the
selected read transaction is already available. If the selected read transaction belongs
to a pending list, then the related information, kept in the next segment field , is
accomplished and stored to local registers. Finally, the Busy flag of the entry that
keeps the selected read transaction is set to 1 to indicate that the corresponding
memory access is in progress.
Figure 4. 14 First Stage
We remind that control information related to the read operation (read address) and
information related to the corresponding dequeue operation (flow_id, transit_id of
the next operation in the pending list) is kept in local registers. The information that
is kept in the local registers is loaded to the pipeline buffer, at the pipeline clock
edge. The pipeline buffer has 6 fields: the read address, the flow_id, the next
segment of the pending list, the NS flag and the transit_id of the current read
operation.
Queue Management Micro-Architecture
68
Second Stage of the Pipelined Process
The datapath of the second stage of the dequeue operation execution process is
shown in the figure 4.15. The input of this stage is the pipeline buffer contents.
During the first cycle the read operation control buffer is updated. The information
that loaded is dependent on the value of the NS field of the entry that keeps the read
transaction. If NS field is set to 0, it means that the read operation either does not
belong to a pending list or it is the last one of the pending list. In each case the
information that loaded to the control buffer is performed in the table 4.4. As table
4.4 presents, the destination target is the Head table. Additionally, it has to update
the AR flag in the Head Table in order to inform a successive dequeue operation for
the same flow that this is in execution. Otherwise, if the NS field in the control
buffer is set to 1, it means that the current read operation belongs to a pending list
and there is a successive pending read operation that waits servicing. In this case, the
information that is loaded to the control buffer is shown in the table 4.5. As table 4.5
presents, the destination target is a PRT entry; it implies that, during the update
phase, the response of the buffer memory will be forwarded to a pending read
operation instead of updating the Head Table. At the end of the first cycle, the PRT
entry, which contains the current read operation, is released by setting its valid flag
to 0. The remaining three cycles are idle because a read operation needs only one
cycle to send its command to the memory controller.
Figure 4. 15 Second stage
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The Pipelined Control Processes Micro-Architecture
69
Valid
1
Valid
1
R/W
R
R/W
R
Dst.Tbl.
Head Table
Dst.Tbl.
PRT
Dst.Entry
Flow_id field
pipeline buffer)
Dst.Entry
next segment field (from the
pipeline buffer)
Src.Tbl.
PRT
Src.Tbl.
PRT
Src.Entry
transit_id field
pipeline buffer)
Src.Entry
transit_id field (from the pipeline
buffer)
read addr.
Read address field (from the
pipeline buffer)
read
addr.
Read address field (from the
pipeline buffer)
(from
(from
Table 4. 4 Control buffer
4.2.7
the
the
Table 4. 5 Control buffer
Handling Exceptional Cases during a Dequeue Operation
As mentioned above for the read transactions that originate from enqueue operations
and as we will explain later for the read transactions that originate from dequeue
operations, additional information, related to the update destination, is accompanied
with the read operation command. When the memory responds with the results, the
accompanied information notifies the proper destination data structure with the
updated values. If a newly dequeue operation was issued and the state of the
corresponding queue head pointer was pending, while the read operation that would
update the queue head pointer was in execution, then the newly dequeue operation
should be stalled until the queue head pointer is available. This stall may last the
latency of the buffer memory access that causes system performance decreasing.
The solution is the bypassing. However, we cannot change the update destination
target of the corresponding read operation control buffer, because we do not know
the exact state of the read execution progress. In other words, due to the pipelined
design of the buffer memory controller, we don’t know in which pipeline stage the
read operation is, so we cannot intervene and modify the destination target of this
operation.
The solution that we propose to handle this exceptional case is described in this
paragraph. The memory controller interface process undertakes the update
operations when the result provider is the buffer memory. If this exceptional case
occurs, a control flag that is visible/accessible from the memory controller interface
process is set to 1. Additionally, control information is kept in a local register, which
called “Exception Register” and contains the source table and transit_id of the
executing read operation and the new destination target (table and entry). When this
process tries to update a resource (data structure) it examines this control flag. If it is
set to 1, then it examines if the source table-transit_id fields of the read operation’s
control buffer are matched to the source table and transit_id of the Exception
Register. If there were no matching, the process would update the initial target.
Otherwise, if there was matching, the process would update the new target, which
was kept in the Exception Register. This solution overcomes this exceptional case
but it needs cautious design in order to maintain strict synchronization.
4.2.8
Queue Manager Interface Process Micro-Architecture
The main tasks of this process are three. The first task is to insert the access
Queue Management Micro-Architecture
70
commands to the memory controller. The second task is to receive the data response
of the buffer memory in the case of a read access. The third task is to forward the
received data of a dequeue operation to the output and to update the data structures
(Queue table, Free List table)/operation table (PRT). This process is pipelined and
consists of five stages. Each stage has latency that equals to the time slot delay.
The first stage undertakes the insertion of the memory access operations to the
memory controller. An access operation consists of three elements: the command
field, the address field and the data field. The read operation has not data field. The
command can be either write or read. The address field consists of five parts: the
module address (1 bit), the device address (4 bits), the bank address (4 bits), the row
address (9 bits) and the column address (6 bits). To be noted here, that the Rambus
memory organization manipulates fixed size 16-byte units. Each memory row
contains 64 units of 16-bytes size; thus, the column address has 6 bits (26=64).
However, the queue manager manipulates 64-byte units (segments). It implies that
each memory access addresses quadruples of memory units and the 2 least
significant bits of the column address is set to 00. In the case of a write operation,
the data field consists of four data buffers; each buffer size is 64 bytes.
The remaining stages of the interface operation are dedicated to serve only the read
operations. The only requirement of a write transaction is to be inserted in the
controller. The memory core writing is performed by means of three modules: the
memory controller, the Rambus memory core interface (RAC), and the memory
core. Otherwise, a read transaction requires from the queue management interface
process to insert a read command and to receive and manipulate the memory
responses. The difficulty of the data-receiving task rises because the memory
responses will be available in a window time slot after the read transactions’
insertion. The minimum delay of memory response is 110ns (2.75 time slots), while
the maximum is 140ns (3.5 time slots) and the window size is 40ns. The latency of
rambus memory access is constant but this variability caused by the transactions’
shifting in the memory controller. Even if the memory transactions are inserted into
the memory controller at the beginning of a time slot, it may initiate them in a
successive cycle due to the memory turn around overhead. This overhead occurs
because read after write and write after read operations cannot be initiated back to
back. Instead, a time interval of 5ns must intervene among read and write
alternating. The main characteristic of the transactions shifting is that it is
accumulative, which means that the memory controller must remember the memory
transaction history. More details on this subject will be referred in section 4.3.5.
The main point of the second stage existence is the provision of a time slot delay.
Noted that the insertion of a memory access operation is performed during the first
cycle of the first stage. The remaining three cycles provide delay. The first two
cycles of the third stage is idle in order to provide additional delay. This delay holds
each read operation until the time where the memory responds with the data. The
tasks that are performed during the third and fourth cycle of the third stage and
during the first and second cycle of the fourth stage have many similarities because
they belong to the critical window time slot to which the memory will respond with
the data. Each of these tasks are split into three functions: the detection of a new
data block from the memory, the received data forwarding toward the output link,
and the update of the queue manager data structures.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
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The Pipelined Control Processes Micro-Architecture
Figure 4. 16 Detection circuit
The detection function can be performed by the circuit, which is illustrated in the
figure 4.16. The memory controller interface has an output for each memory module
that indicates the receiving of a new data block from the corresponding memory
module. Each of two outputs has 1 bit size and alternates its value at the beginning
of a new data block receiving. We remind that each data block size is 64 bytes. Each
time a read operation response originated by a dequeue operation, the receiving data
block must be forwarded to the output or alternatively it must be loaded to the
output buffer. However, it is possible both memory modules to respond with the
data almost simultaneously. The example of the figure 4.17 illustrates this case. It
implies that the data of both memory modules want to access the output buffer
simultaneously and cause a conflict. As the figure shows, the occurrence of this
conflict is not a usual case because we don’t send two read operations as well as two
write operations toward the two memory modules at the same time slot. This conflict
looks like with the classical case of the critical section accessing by multiple
processes in the computer operating systems. Similar to the operating systems we
implement an arbitration process that allows the access of the output buffer (critical
section) to only one process. This arbiter architecture is described in section 4.2.9.
Finally, the update of the queue management data structures can be performed only
during the fourth cycle of the third or the forth pipeline stage in order to schedule
this task and use the resources more efficiently, see the section 4.2.9. The overall
process pipeline is illustrated in the figure 4.18.
Figure 4. 17 Ingress Module Output Access Conflict
Queue Management Micro-Architecture
72
Figure 4. 18 Queue Management Interface Process in Pipeline Fashion
4.2.9
Resource Conflicts among Queue Management Processes
Multiple processes access the queue manager data structures. For example the
second and the third enqueue processes access the Tail table. The Head table is
accessed by the two dequeue processes and by the QM interface process. The PWT
is accessed by all the enqueue processes, while the PRT is accessed by all the
dequeue processes and the QM interface process. Finally, the first enqueue and the
QM interface process access the Free List Table, while the Free List Cache is
accessed by the first and third enqueue processes and by the second dequeue
process. The case of a resource accessing from multiple parallel processes causes the
danger of concurrent accessing. When many processes try to access the same
resource concurrently, a conflict occurs. In the computer science, the operating
system overcomes this conflict by using techniques of mutually exclusive access of
the critical section. In our system we overcome this issue by twofold ways. The first
manner is to schedule the resource accesses of the multiple processes to different
clock cycles. It implies a strict synchronization among processes. On the other hand,
two independent processes can access different entries of the same table (memory
block) concurrently, by using dual ported memories. Finally, in the cases that two or
more processes want to access the same entry of a resource, we use mutually
exclusive techniques originated in the computer operating systems, but they are
implemented in hardware.
Critical Section Access Arbiter Architecture
In general, an arbiter schedules the resource accesses of multiple processes. We
show the implementation of an arbiter that handles the resource accesses of two
different processes. Our arbiter has two request input lines from the two processes. It
grants only one process to access the critical section at each time. The grant signal
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The Pipelined Control Processes Micro-Architecture
73
consists of two bits: the first bit indicates whether the critical section is busy or not
(key state) and the second bit indicates the process that accesses the critical section
(key owner). We also use some state registers, such as the owner state, the pending
request, and the pending request_id. A process may be in the critical section for at
most four clock cycles (clock cycle: 10ns); thus the owner state register indicates the
time that the process has been in the critical section (2-bits size). The pending
request indicates if a process waits to achieve the access to the critical section while
an other process has already been in the critical section. The pending request_id
indicates the identifier of the pending process. The arbitration cycle lasts a clock
cycle (10ns). The arbiter also uses a final state machine in order to schedule the
incoming requests. The FSM of our arbiter is illustrated in the table 4.6.
in
Current state
Next state
Req
est
Key
state
key
owne
r
owne
r
state
Pendin
g
request
Pendin
g
request
id
Next
key
state
Next
key
owne
r
Next
owne
r
state
Next
Pendin
g
request
Next
Pending
request
id
0 0
1
S
3
1
A
1
A
0
0
X
0 0
1
S
?3
1
C
1
S
+1
1
C
0 0
1
X
3
0
X
0
X
X
0
X
0 0
1
S
?3
0
X
1
S
+1
0
X
0 0
0
X
X
1
A
1
A
0
0
X
0 0
0
X
X
0
X
0
X
X
0
X
0 1
1
S
3
1
A
1
A
0
1
0
0 1
1
S
3
0
X
1
0
0
0
X
0 1
1
S
?3
0
X
1
S
+1
1
0
0 1
0
X
X
1
A
1
A
0
1
0
0 1
0
X
X
0
X
1
0
0
0
X
1 0
1
S
3
1
A
1
A
0
1
111
1 0
1
S
3
0
X
1
1
0
0
X
1 0
1
S
?3
0
X
1
S
+1
1
1
1 0
0
X
X
1
A
1
A
0
1
1
1 0
0
X
X
0
X
1
1
0
0
X
1 1
1
S
3
0
X
1
0
0
1
1
1 1
0
X
X
0
X
1
0
0
1
1
Table 4. 6 The Arbiter FSM
4.2.10 Search Engines Architecture
Throughout the description of the queue management architecture in the chapter 3,
we referred to the requirement for implementing a fast search engine, which traces
the queue management operation table (Pending Write Table, Pending Read Table,
Free List Table) entries in parallel. Remember that these search engines mainly
Queue Management Micro-Architecture
74
search to find a write or a read transaction that will not cause a memory bank
conflict. So, a search operation may be split into two simple search functions, which
are performed simultaneously. The first function searches for matching on a fraction
of examined bits, while the second function searches on the remaining bits for not
matching. Not matching search is referred to the conflicting cases.
Search Engines in the Queue Management Architecture
Let study the cases of searching. During an enqueue operation the following four
searches must be performed. The first search engine must achieve a free entry in the
pending write table in order to insert a new enqueue operation and to allocate the
corresponding buffer in the Transit Buffer for keeping an incoming segment. The
first search engine must examine only the valid field (1-bit) of the PWT entries.
The second search engine is required during the phase of selecting a non-conflicting
write transaction, during the enqueue execution process. This search must examine
multiple fields of the pending write table entries. Firstly, it must find a valid, not
busy, and ready entry. Secondly, it must select an operation that it will not cause
module/bank conflict in the buffer memory. It can be achieved by selecting an
operation that is going to access a non-busy bank.
Because we select the eligible pending write operations for execution, kept in the
PWT entries, in a round robin order, we use an additional (third) search engine to
move the round robin pointer to the next pending operation. This search engine
examines only the Valid flag of the PWT entries.
Finally, the extraction of a free buffer from the free list, during an enqueue
operation, requires a parallel searching in the free list table. Note that, the free buffer
must not belong to a busy bank. The free list is organized as a set of 512 queues (per
bank queueing). Selecting an eligible free buffer requires searching to 512 entries of
the free list table.
During a dequeue operation, three searches must be performed in the pending read
table. The first search engine finds a free entry in the pending read table for keeping
a newly issued dequeue operation. The second search engine selects an eligible
dequeue operation to execute. Because we service the eligible dequeue operations in
a round robin order, an additional (third) search engine is required in order to shift
the round robin pointer to the next pending operation in the PRT.
Parallel Search Engines
Search engines that examines multiple entries of a memory block in parallel uses
Content Addressable Memories (CAMs) along with a priority encoder. In order to
construct a common search engine structure for all the required seven search engines
(4 for the enqueue operation and 3 for dequeue operation), we will study which
fields of the PWT, PRT, and Free List table must be examined from each search
engine.
The first and the third search engine of the enqueue operation examine only a 1-bit
field in the PWT entries, the Valid flag. The former for matching (V=0), and the
latter for matching (V=1). The second search engine of the enqueue operation
lookups 12 bits in the PWT entries. It is looking for an entry that is valid (V=1) and
contains a ready enqueue operation (R=1). This operation must be not busy (B=0),
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which means that it has not already been in execution. Additionally, it examines the
9 most significant bits of the writing address in order to avoid a conflict. These bits
identify the module, the device, and the bank to which the writing buffer belongs (1
bit for the RIMM module, 4 bits for the device and 4 bits for the bank).
In the case of a free buffer extraction, 10 bits must be examined in the Free List
Table. One empty bit is examined for checking whether a bank in the free list is
empty of buffers or not. Next, the 9 most significant bits of the free buffer address
are examined in order to avoid a bank conflict. Similar to the above reference, these
9 bits identify the module (1), the device (4), and the bank (4) to which the free
buffer belongs.
Figure 4. 19 The Search Engine Block Diagram
The first and the third search engine of the dequeue operation examine only a 1-bit
field in the PRT entries, the Valid flag. The former for matching (V=0) and the latter
for matching (V=1). The second search engine of the dequeue operation lookups 12
bits in the PRT entries. It is looking for an entry that is valid (V=1) and contains a
ready dequeue operation (R=1). This operation must be not busy (B=0), which
means that it has not already been in execution. Additionally, it examines the 9 most
significant bits of the read address in order to avoid a conflict. These bits identify the
module, the device, the bank to which the reading buffer belongs (1 bit for the
RIMM module, 4 bits for the device and 4 bits for the bank).
The examined fields of all operation tables must be extracted from these tables and
must be located to CAMs in order to perform parallel lookups to these tables entries.
The total memory requirements are 128x12 bits for PWT, 128x12 bits for PRT,
512x10 bits for Free List Table. The results of each search operation are stored to an
one-dimension array. This array contains so many entries as the number of the
searched table entries. The array elements that correspond to the searched tables
matching entries are set to 1, while the remaining are set to 0. Next, the search
results, which are kept in the array, are driven to a priority encoder in order to be
identified the matching entry with the highest priority, as the figure 4.19 shows. The
structure of a priority encoder is described in the section 4.4.11.
The Priority Encoder Data Structure
Queue Management Micro-Architecture
76
The priority encoder is a function that has a N-bit binary number input and a
log2(N)-bit binary number output. The output number points the location of the
most significant 1 in the input number. For example if the input number is the
00101101 and the output is the number 110. In the case of a search engine the
priority encoder is required to index the first matching with the highest priority. For
a binary number the highest priority is identical to the most significant bit. Instead,
the search engine needs the flexibility to determine itself the highest priority point.
The figure 4.20 illustrates this flexibility. At the leftmost scheme the priority
encoder determines the upper element with the highest priority and the bottom
element with the lowest; at the rightmost scheme the middle element has the highest
priority, while the priority of the subsequent elements decreased in a circular order.
The priority encoder implementation in hardware is modular. The priority encoder
primitive (cell) has two versions as figure 4.21 shows. The priority encoder cell T1
is the complement of the T2 counterpart and vice versa. T1 and T2 are combined to
build a two-bit priority encoder, figure 4.22. In the same way, a multiple bit priority
encoder can be built. Figure 4.23 shows an 8-bit priority encoder which determines
the upper element with the highest priority.
Figure 4. 20 Priority Alternatives
Figure 4. 21 T1 and T2 Priority Encoder Cells
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Figure 4. 22 Two-bit Priority Encoder Figure 4. 23 Eight-bit Priority Encoder
Priority Encoder Modification
The priority encoder of figure 4.23 sets the highest priority to the upper element. We
modify this scheme in order to make it more flexible. The figure 4.24 shows this
modification for a 4-bit priority encoder. We locate T1 and T2 cells alternately and
insert between them a 2x1 multiplexer. The select inputs of the four multiplexers are
driver by a 2x4 decoder. The decoder input defines the point that the search will start
(the point with the highest priority). The upper cell is wired with the lower cell in
order to perform a circular searching.
Figure 4. 24 The Modified Priority Encoder
4.2.11 Free List Organization Alternatives
The free list organization has two main alternative implementations: the bitmap
organization and the queueing organization.
Queue Management Micro-Architecture
78
Bitmap Organization
We explain the bitmap organization by giving an example. Next we apply the
bitmap organization to our memory system. For the example we consider that the
free list consists of 64 buffers. The state of these buffers is kept in an 8x8 array as
the figure 4.25 shows. The buffer emptiness is indicated by the “1” and the opposite
by “0”. By applying the “or” operation to the elements of a row and storing the
result to the corresponding entry of a 2x4 array, showed in the middle part of the
figure, we compress the information of the initial array. The middle array has
information only for the rows of the initial array that have at least an empty buffer.
Continuing this process, by applying the “or” operation to the first and second row
of the middle array, we update the rightmost array. Selecting an element with 1 of an
one-dimension array is a trivial issue; we can use a priority encoder. If we select an
element with 1 from the rightmost array, we go back to the corresponding row of the
middle array and select an element with 1 of this row. Continuing similarly, we
choose a row from the leftmost array, which has at least one free buffer; all the
previous steps ensure this situation. Finally, we select an element with 1 from this
row. The location of this element in the array points to a free buffer. Applying the
bitmap organization to our buffer memory, which contains 4 million buffers, will be
extremely expensive.
Figure 4. 25 Bitmap Free List organization
Per-Bank Queueing
Another alternative implementation of the free list is to organize the free buffers in
single linked lists (queues). A queueing organization requires two pointers: a pointer
to the queue head and a pointer to the queue tail. It also requires to assign a pointer
to each free buffer, which we call next pointer. Each free buffer next pointer field
indicates to the next free buffer in the list, as the figure 4.26 shows. In order to build
a more flexible scheme, we organize free buffers of the memory in a per- bank
queueing scheme, as the figure 4.27 shows.
Figure 4. 26 Free list organization as a linked list
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Rambus Memory Technology
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Figure 4. 27 Per-bank queueing Free List Organization
4.3 Rambus Memory Technology
Rambus is a new DRAM technology that provides high throughput and great
capacity. The Rambus memory architecture consists of four main elements: the
memory controller, the Rambus Interface (RAC), the Rambus Channel, and the
RDRAM devices. The Rambus Interface is implemented on both the memory
controller and the Rambus Channel. The memory controller and the RDRAM
devices are connected via the channel. The controller is located at one end and the
RDRAM devices are distributed along the channel, which is parallel terminated at
its characteristic impedance at the other end; the channel terminator eliminates any
reflection. The Rambus memory architecture is shown in the figure 4.x. The
common channel consists of a 16-bit data bus and a 8-bit control and address bus (3
row and 5 column pins). The channel clock cycle is 400MHz but data and control
transfers are performed at both clock edges. Therefore, the channel data transfer rate
is 12.8Gbps (16 bits * 2 * 400MHz). As figure 4.28 shows there are two clocks: the
Clock to Master (CTM) and the Clock from Master (CFM). The former travels
toward the controller and the latter travels away from the controller. In a read
operation the data travel through the channel in parallel to the CTM, while in a write
operation the data travel in parallel to the CFM in order to minimize clock to data
skew. The data transfers are performed at the granularity of 16-bytes data units per
10ns, while the control/address transfers at the granularity of 8-bytes units per 10ns.
Figure 4. 28 Rambus Technology
Queue Management Micro-Architecture
4.3.1
80
Read and Write Operations in a Pipelined Fashion
A read operation is shown in figure 4.29. Generally, a completely random access is
performed by the assertion of the ACT command across the ROW pins followed by
a read command sent across the COL pins. After the data is read, a precharge
command is executed to prepare that bank for another completely random read. Data
is always returned in a fixed number of cycles from the end of the read command.
Figure 4. 29 Read Transaction
Write transaction timing is very much like read transactions. The control packets are
sent in the same way as the read command. However, one significant difference
between the RDRAM and a conventional DRAM is that write data is delayed to
match the timing of a read transaction in order to maximize the usable bandwidth
on the data pins. On a conventional SDRAM, the write-read transaction alternating
causes a gap on the data bus from the write data to the read data. The RDRAM
avoids this gap by sending the data later in time. A write command on the COL pins
tells the RDRAM that data will be written to the device on an exact number of clock
cycles later. This data would be written to the core as soon as the data is received.
The write transaction is shown in the figure 4.30
Figure 4. 30 Write Transaction
Each of the commands on the control bus may be pipelined, allowing much higher
throughput. The ACT commands can completely absorb the ROW pins, allowing
16-byte random transfers to occur. In order to completely fill the data bus, column
command would be continuously sent on the COL pins. Except for small gaps of 5ns
required for bus turn-around going from a write to a read, these busses can be fully
utilized. In the figure 4.31 is presented an example of interleaved write and read
transactions. The transaction transfer granularity is 64-byte data blocks.
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Rambus Memory Technology
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Figure 4. 31 Interleaved read and write transactions
4.3.2
Rambus Memory Device Architecture
The Rambus device (RDRAM) consists of two major blocks: a core block build
from banks and sense amplifiers similar to those found in other types of DRAM, and
a Direct Rambus interface block which permits an external controller to access this
core at up to 12.8Gbps.
The control registers supply the RDRAM configuration information to the controller
and they select the operating modes of the device. The RDRAM contains 17 sense
amplifiers. Each sense amplifier consists of 512 bytes of fast storage and can hold
one-half of one row of one bank of the RDRAM. Each sense amplifier is shared
between two adjacent banks of the RDRAM. This introduces the restriction that
adjacent banks may not be simultaneously accessed.
The RQ Pins carry control and address information. They are split into two groups:
the row group and the column group. The row group pins carry the row-command,
while the column group pins carry the column-command. The row command may
initiate an ACT (active) command or a PRER (precharge) command. An ACT
command causes one of the 512 rows of the selected bank to be loaded to its
associated sense amplifiers, while a PRER command causes the selected bank to
release its two associated sense amplifiers, permitting a different row in that bank to
be activated, or permitting adjacent banks to be activated.
The column command may initiate a read command or a write command. The read
command causes one of the 64 blocks of one of the sense amplifiers to be
transmitted on the data pins of the Rambus channel. The write command causes a
block received from the data pins of the Rambus channel to be loaded in to the write
buffer. The data in the write buffer is automatically retired to one of the 64 blocks
of one of the sense amplifiers during a subsequent column commands. A retire can
take place during a read, write, or no-operation to another device, or during a write
or no-operation to the same device. The write buffer will not retire during a read to
the same device.
Queue Management Micro-Architecture
4.3.3
82
Rambus Memory Module Architecture
The Rambus memory upgrade packaging consists of the Rambus RIMM memory
module and connector, which are similar in size to existing DIMM memory modules
and connectors. The RIMM module supports up to eight RDRAM on each side. A
single RIMM module can accommodate up to 128 Mbytes of memory using 64-Mbit
RDRAM devices.
4.3.4
Rambus Memory Interface
The Direct Rambus ASIC Cell (Direct RAC) is a library macro cell used in ASIC
designs to interface the core logic of a CMOS ASIC memory controller to a highspeed Rambus Channel. The RAC typically resides in a portion of the ASIC I/O pad
ring and converts the high-speed (800 MHz) Rambus Signal Level on the Rambus
channel into lower-speed CMOS-level signals usable by the memory controller
designer. The RAC functions as a high performance parallel-to serial and serial-toparallel converter performing the packing and unpacking functions of high
frequency data packets into wider and synchronous 144-bit data words.
The RAC has six main unidirectional CMOS input/output buses to the memory
controller side: a receive control data bus (RDQ – 64 bits), a transmit control data
bus (TDQ – 64 bits), two receive data buses (RdataA and RdataB, 72 bits each one),
and two transmit data buses (TdataA and RdataB, 72 bits each one). The RAC has
three main bi-directional RSL buses to the Rambus Channel side: a control bus (RQ
– 8 bits), and two data buses (DQA and DQB – 9bits each one). The RQ bus is
further subdivided into a three-bit ROW bus and five bit COL bus, while RdataQ
and TdataQ buses are subdivided into 24-bit and 40-bit buses.
The performed converting functions of serializing the memory controller commands
– data toward the Rambus channel and paralleling the Rambus channel data
responses toward the memory controller by the RAC interface are assisted by some
other control signals. The RAC function is mainly subdivided into two operations:
the Transmit and the Receive operations. A write transaction from the controller
maps to a Transmit operation in the RAC, while a read transaction maps to a
transmit and a receive operation. In the case of the write transaction, the transmit
operation transmits the write command and the data toward the rambus channel. In
the case of the read transaction, the transmit operation transmits the read command
toward the rambus channel, while the receive operation transmits the responding
data from the Rambus channel to the memory controller.
The memory controller must send a control command to the RAC in order to inform
the RAC that it will send a command (read/write) packet to the transmit control bus
or a data packet to the transmit data buses at a fixed delay after this control
command. Each control command is synchronized to the controller clock cycle
(100MHz) and it registers the time delay of loading the command-data packet. Note
that the signals TSEL and RSEL indicate the time that the parallel data packets will
be transmitted or received to/from the RAC interface. This figure indicates the
conversion of parallel command-data packets to serial command-data through the
RAC interface.
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Rambus Memory Technology
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4.3.5
Rambus Memory Controller
The Rambus memory controller resides at the one end of the Rambus Channel. It
directly connects to the RAC interface as an input-output cell. The controller
provides the protocol for performing Read and Write transactions to the Rambus
DRAMs. It translates address, data and command into the Rambus protocol. It
presents a simplified, high level 128/144-bit data path to the controller designer. It
also supports all control functions including memory initialization and memory’s
control buffers configuration, refreshing, and transactions interleaving.
We design a Rambus memory controller, which is oriented to support the
requirements of the queue manager system. For example, due to the data segment
size (64-bytes), the data block transfers have 64-byte size (the transfer granularity is
64-bytes). So each memory access addresses four contiguous 16-bytes units in the
memory. Additionally, since the buffer memory consists of two RIMM modules, the
memory controller is split into two Rambus controllers, which work independently.
Finally the memory controller system provides an appropriate high level interface to
the remaining queue management system in order to simplify its communication
with the buffer memory. We remind that the queue manager communicates with the
memory controller by means of the queue manager interface process.
Rambus Memory Controller Micro-Architecture
In order to perform the controller architecture, we summarize its main tasks. It has
to manage the appropriate control signals in order to perform a read or a write
transaction. It has to send synchronized control, command, address and data packets
at the appropriate inputs of the RAC interface. It has to refresh the memory banks
after accessing them. Finally, it must efficiently interleave the memory transactions
in a pipelined fashion in order to achieve full memory throughput utilization, and
concurrently, it must satisfy the strict Rambus timing constraints. An issue that
complicates the memory controller implementation is that it has to delay the
insertion of a new write operation which follows a read operation for a half clock
cycle due to the turn around overhead. Since this delaying function behaves
accumulatively, the controller must remember the history of the previously inserted
transactions. The minimum gap between two successive operations is a half clock
cycle (50ns); it implies that memory transactions may be inserted at both rising and
falling edge of the clock cycle. In order to exploit the buffer memory high
throughput, we insert a memory access (read or write) per time slot and per memory
module. The exact cycle where the controller inserts a new memory transaction in a
time slot is dependent on the history of the previous memory accesses.
In order to represent the history of memory accesses in hardware, we define two
parameters: the state, and the transaction type. The state consists of three bits (23 = 8
possible states). The two most significant state bits indicate the cycle in a time slot
that the new operation was inserted; since the time slot consists of four clock cycles
there are four states. The least significant state bit indicates whether the new
transaction is inserted at the rising or at the falling edge of the clock cycle. The
operation type determines if the new operation is a read or a write access. The
definition of the cycle in a time slot that the new operation will be inserted is
dependent on the state and type of the previous operation and the type of the new
operation. Upon a transaction arrival at the beginning of a time slot, the memory
controller examines the state and type of the previously inserted transaction
Queue Management Micro-Architecture
84
accompanied with the type of the current transaction in order to determine the state
of the new transaction; determining the state of a new operation is identical of
determining the cycle that it will be inserted. The FSM that describes the above
function is illustrated in the figure 4.32. As figure 4.32 shows, the state parameter is
split into two fields: cycle number (2-bit size) and clock edge (1 bit size). If clock
edge field is 1, it corresponds to the rising clock edge, while if it is set to 0, it
corresponds to the falling clock edge.
Figure 4. 32 Transaction Insertion FSM
We remind that a new read or write transaction may be inserted at the rising or the
falling edge of the four clock cycles in a time slot. The figure 4.33 shows a timediagram of a read transaction. In this diagram we show that the signaling is identical
independent the exact time that the read transaction is inserted. Note that the
Trowsel, Tcolsel, Tdatasel, Rdatasel are control signals that indicate the timing a
row command, a column command, a transmitted data packet, or a received data
packet will be loaded to the RAC interface. The figure 4.34 shows the time-diagram
of a write transaction.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
85
Rambus Memory Technology
Figure 4. 33 Read Transaction time-diagram
Figure 4. 34 Write Transaction time-diagram
5 Verilog Description & Simulation
In order to verify our architecture we describe the multi-queue management
architecture at OC-192 line rate using the hardware description language “Verilog”.
We wrote a behavioral model that simulates the queue manager architecture at a
clock cycle accurate level. A basic choice of designing the architecture was the
choice of the core clock cycle period. We assumed that the queue manager
architecture operates with a clock frequency of 100 MHz for the following reasons:
?? The choice of the Rambus DRAM technology for the buffer memory
determines our architecture clock frequency. The external clock of the RAC
interface cell is 100 MHz in order to achieve 12.8 Gbps throughput (write or
read 16-bytes per 10 ns)
?? We assumed that the access latency of the on-chip memories plus the
combinational logic latency fit within a clock cycle of 10 ns with relative easy.
This assumption is realistic for 0.18 micron technology.
?? We assumed that the search engines of our architecture can operate at or above
100 MHz. This assumption is realistic because modern CAM memories
promise 100 million searches per second [26].
Another crucial issue for the queue management architecture description is the
choice for the pipeline stage length. This choice is determined by the size of packet
segments. More precisely, in order to achieve queue management throughput of 25.6
Gbps we have to enqueue and dequeue a packet segment per time slot. The time slot
corresponds to the time interval for writing or reading an 64-bytes segment from the
Rambus memory, which equals to 40 ns; thus the pipeline stage length equals to 40
ns.
5.1 Hardware Implementation Cost
The Rambus memory model in Verilog along with the model for the RAC interface
cell was kindly provided to us from Rambus Inc. The memory model has the
following parameters:
?? The number of RDRAM chips in the RIMM module
?? The capacity of each RDRAM chip (64 Mb, 72 Mb, 128 Mb etc.)
?? The operation speed
The memory initialization is achieved by loading a file of appropriate format1 in the
memory.
The Verilog hardware description of our memory controller uses 3800 code lines,
while our queue management architecture uses 4200 code lines, which are
distributed as follow:
?? 1500 code lines for the enqueue control processes
?? 1500 code lines for the dequeue control processes
?? 1200 code lines for the queue management interface process
1
The file format is :[@address data]. Both address and data are hexadecimal numbers
Verilog Description & Simulation
88
We estimated the on-chip memory requirements for the datapath chip. Table 5.1
shows this estimation per memory block. The PWT, the PRT as well as the
Head/Tail table are split into multiple separate tables in order to allow parallel
accessing.
Memory
Block
Internal
Organization
Memory
ports
Head/Tail
Table
2 x 64K x 32
1 port
50 mm2
Pending
Write
Table
3 x 128 x 32
2 ports
0.36 mm2
Pending
Read Table
2 x 128 x 32
2 ports
0.24 mm2
Transit
Buffer
15 x 128 x 32
2 ports
1.84 mm2
total
ASIC area
(0.18µm)
4 Mbits
Figure 5. 1 Datapath chip memory requirements
We also estimate the hardware complexity of our architecture in terms of gates and
flip-flops for 64 Kflows, as shown in table 5.2
Processes
Gates
Flip-Flops
Packet entry
3K
4K
Enqueue issue
7K
10 K
Enqueue execution
12 K
15 K
Dequeue issue
10 K
14 K
Dequeue execution
13 K
15 K
Queue management
interface
15 K
22 K
60 K
80 K
Total
Figure 5. 2 Hardware complexity of our architecture
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Verification
89
5.2 Verification
In order to verify the design that simulated, the queue management architecture at
cycle accurate level, using test patterns that simulate incoming traffic at 10 Gbps
maximum load. We assumed that the packet segmentation is performed externally of
the architecture model, i.e the test patterns contain segment arrivals rather than
packet arrivals. The test patterns parameters are:
?? The input load
?? The segment arrival distribution
?? The maximum packet size
?? The flow identifiers of the incoming packets
?? The header processing delay variability for incoming packets
The test patterns were generated by using the C programming language and stored in
files. The files’ format is the following:
time slot
packet_id
segment_id
segment type
flow_id
Header processing delay
1
1
1
0
1500
5 (time slots)
2
1
2
1
1500
3
1
3
3
1500
4
2
1
3
16383
1
5
3
1
3
0
2
The test pattern files have the following information: at times slot 1 the first segment
of the first packet arrived. The segment type identifies the type the incoming
segment, which mean that it identifies if the incoming segment is the first, an
intermediate or the last segment of a packet. This information is required in order to
organize the incoming segments into packet queues at the time of segment arrivals.
Instead we have to wait the arrival of the next segment in order to identify the tail of
the last packet and the head of a new packet. The flow_id and processing delay
fields identify the flow, which the packet (packet segment) belongs, and the delay,
which the packet suffers during its header processing period.
Except for the test pattern generation, the header processor and scheduler
simulation is required. Both header processor and scheduler can be simulated as
devices that schedule the enqueue and dequeue operations, correspondingly. The
input of the header processor device is the triplet of the packet identifier, flow
identifier and processing delay. The header processor schedules the incoming
enqueue operations according to their processing delays. The input of the scheduler
is the state of system queues. The scheduler defines the order of the packet
departures from the active flows (non-empty queues). Using calendar queue data
structures may simulate both header processor and scheduler.
The system architecture verification is split into four stages. In the first stage we
implemented and verified the Rambus memory controller. The verification is
performed by writing memory segments and then reading them in order to compare
writing and reading data. In the second stage we implemented the six control
processes of the queue manager and verified each process separately by using short
Verilog Description & Simulation
90
simulation runs. In the third stage we verified all the enqueue and all the dequeue
control processes separately. In the fourth stage, we verified all the system processes
by using short simulation runs. Verification consists of the following steps:
1. we generate test patterns that simulates the incoming traffic in C language
2. we organize the incoming packet segments into queues according to the
flows that they belong and save the result to the fileA
3. we apply these test patterns to the queue manager architecture model in
Verilog
4. we organize the outgoing packet segments of the Verilog model into queues
according to the flows that they belong and save the result to the fileB
5. we compare the fileA and fileB. The results of this comparison verifies:
?? the packet/segment loss
?? the packet segments output order
?? if the incoming segments were linked to the proper queues
The test patterns that were successfully run through the queue manager behavioral
model were short1, due to time constraints of this master thesis. As a consequence,
our design has been debugged only partially, up to now.
1
test patterns consisting of 128 segments pass through our queue manager behavioural model
successfully. We detected some bags in the way of updating memory blocks of our architecture
(queue management interface process).
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
6 Conclusions and Open Topics
This thesis studied the architecture of high-speed switches and routers that support
Quality of Service (QoS) guarantees. It concentrates on the architecture of ingress
and egress interface cards at OC-192 (10 Gbps) line rate and describes the supported
functions. It presents an effective chip partitioning for the ingress module that
economizes on chip-to-chip communication, so that pin count and power
consumption are reduced. We then focused on the queue management subsystem in
the ingress and egress line cards. We believe that the provision of QoS guarantees
usually requires flow isolation that can be effectively achieved by using per-flow
queueing. Per-flow queueing for thousands of flows was considered an excessively
expensive architecture up to a few years ago. Modern technology, however, provides
the means to implement such architectures within a fraction of an integrated chip
(IC). This thesis studied the implementation of such architectures at OC-192 (10
Gbps) line rates. We showed that, although challenging, this implementation is
feasible, using advanced hardware techniques that were developed for
supercomputers in the 60’s and are used in high-end microprocessors now-a-days.
We assumed Rambus dynamic RAM (RDRAM) technology for buffer memory, in
order to provide large capacity & high throughput at low pin count. In order to
effectively use DRAM buffer memory, we schedule the memory accesses in the
presence of bank conflicts. We use multiple, pipelined control processes to achieve
out-of-order scheduling of DRAM accesses. The pipeline data dependencies among
successive operations are handled by using Tomasulo’s dynamic scheduling
techniques (operand renaming). We also report methods of economizing off-chip
memories and chip pins by locating the queue management pointers in the buffer
memory and using free list bypassing and buffer Preallocation. Finally, we described
our architecture using behavioral Verilog (Hardware Description Language), at a
clock accurate level, and we estimate the hardware complexity at 60 K gates and 80
Kflip-flops for 64 Kflows. Unfortunately, we run short simulations due to time
constraints of this master thesis.
Debugging our architecture for large test patterns and modifying the verilog
tescription code to a synthesizable code are two challenging issues for future work.
Through this thesis, we introduced ideas of implementing header processing and
scheduling functions (in chapter 2). There are interesting ideas on header processing
and flow classification at high speed that can be implemented in hardware. Also, the
implementation of a scheduling discipline, which handles thousands of flows
(guaranteed service & best effort) and support traffic shaping/ policing functions at
OC-192 line rates, is another interesting area for future research.
Conclusions and Open Topics
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
92
Flow Classification
93
7 Appendix A
7.1 Flow Classification
7.1.1
Recursive Flow Classification (RFC)
The flow classification function can be viewed as a mapping of S bits in the packet
header identifier to T bits of the corresponding class_id. The main aim of the RFC
[17] is to perform this mapping over several stages, as shown in the figure 2b.1. This
mapping is performed recursively; at each stage the algorithm performs a reduction
by mapping one set of values to a smaller set. The RFC algorithm has P phases; each
phase is consisting of a set of parallel memory lookups. Each lookup is a reduction
in the sense that the value returned by the memory lookup is shorter than the index
of the memory access.
Figure 2b. 1 Recursive Flow Classification
The RFC performance can be tuned with two parameters: the number of phases and
the way the memory access results of one phase are combined to index the memories
of the next phase in order to yield the best reduction. The latter can be achieved by
combining the memory access results with the most correlation without causing
unreasonable memory consumption. As the number of phases increases the total
amount of memory decreases but the number of memory accesses per classification
increases. An important disadvantage of this algorithm is the big time for
preprocessing and updating the memory contents. The classification rates that the
RFC performs are 30 million packets per second (for 40-bytes minimum size
packets).
Appendix A
7.1.2
94
Flow Classification by using Hashing functions
The flow classification is a process that examines several fields of the packet’s
header in order to classify the packet to the proper flow. Since the set of examined
header fields has no hierarchical structure and due to the large size of this set, the
longest prefix lookup algorithms can not be applied to the flow classification case.
Instead, hashing algorithms are more efficient to be applied. The basic block in the
hashing architecture is a simple hash table, where each table entry contains a pointer
to a routing table with the forwarding information. The index into the table is
computed as a hash function over the packet identifier. A hashing scheme has the
property that a hash function can map several identifiers into the same table location.
The approach of [18] uses several parallel hash paths, where each path consists of a
hash table and a hash function. In this scheme, a given identifier can only appear in
one of the paths at a time. Therefore a lookup of a given identifier will succeed in at
most one of the paths, and consequently all paths can be searched in parallel. Figure
2b.2 shows a lookup engine with three parallel paths. The hash functions are
denoted as H1, H2, and H3.
Figure 2b. 2 Flow Classification by Hashing
A hash table entry contains an index into a second level table, the address table,
where the full packet identifier is stored together with the forwarding information
for the destination. When the lookup engine has detected a hit in a hash table, the
packet identifier is compared to the original identifier in the table, making sure that
they are the same. The hash calculation, the memory lookup, the table lookup and
the comparison are all independent operations and can work in parallel, thus the
lookup can easily be pipelined to increase the throughput. The modification of the
hash architecture in a pipelined fashion is also performed in the figure 2b.2. The
performance of this pipeline is determined by the slowest pipeline stage, which is
the hash memory stage; thus, this pipeline is capable to perform one lookup per
memory cycle. The most common used hash memories are the Content Addressable
Memories (CAM), which can perform parallel lookups. The modern CAMs provide
up to 100 million searches per second and a single module can handle up to ½
million entries.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
7.2 IP Routing Lookup
7.2.1
Multi-stage IP routing by using Small SRAM Blocks
In this architecture [18], the longest prefix matching lookup is based on a
tree representation of the routing table, where the tree is searched from
shorter prefixes to longer. In contrast to other tree-based schemes, this tree
representation uses a small, fixed number of prefix lengths; it implies that the
prefix tree consists of a fixed number of levels. The implementation of such
a prefix tree is similar to the data structure commonly used for page tables in
virtual memory systems. The prefix tree structure is partitioned into several
tables, figure 2b.3, where each table represents a prefix length. The figure
2b.3 shows a prefix tree with four levels. An entry in a table either represents
a valid route and then contains a pointer into a table defining the next hop, or
represents a part of a route and then contains a pointer to a new table. The
time it takes to lookup a route depends on the number of levels in a prefix
tree. Every level requires one memory access, so with fewer levels, the
lookup is faster. However, fewer levels require more memory due to prefix
expansion; each prefix with different length from the fixed defined prefix
lengths is expanded into several longer prefixes.
Figure 2b. 3 Multi-stage IP routing
The performance of the design is mainly limited by the speed of the memory
accesses. Since the design can be pipelined with one memory access per pipeline
stage, it can perform lookups at the rate of one lookup per memory cycle.
Furthermore, the maximum delay of a lookup is the memory cycle time times the
number of stages in the pipeline. By using SRAM with memory cycle time of 100ns,
it is possible to process 10 million packets per second. Assuming that an average IP
packet is 1000 – 2000 bits long, this means that each lookup can deal with 10-
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Appendix A
96
20Gbps worth of traffic. The memory consumption is parameterized on the number
of levels in the prefix tree.
7.2.2
Two-stage IP routing by using Large DRAM Blocks
This approach [20] shows a longest prefix match address lookup architecture, which
needs one or at most two memory accesses. By examining the statistics of backbone
switch routing tables, it is verified that there are very few routes with prefixes longer
than 24-bits. The majority (99.93%) of the prefixes is 24-bits or less. Based on the
above results, the proposed routing scheme makes use of the two tables shown in
figure 2b.4, both stored in DRAM. The first table (called TBL24) stores all the
possible route prefixes that are up to- and including- 24-bits long. This table has 224
entries, addressed from 0.0.0 to 255.255.255. The second table (TBLlong) stores all
the route prefixes in the routing table that are longer than 24-bits. Upon a packet
arrival the 32-bit destination address is extracted. The 24 most significant bits
indexes the first table. If the destination route has prefix length up to 24-bits, the
access to the TBL24 is adequate to define the next hop. If the route prefix length is
greater than 24-bits, then a second access to the TBLlong is required in order to
accomplish the next hop.
Figure 2b. 4 Two-stage IP routing
Although (in general) two memory accesses are required per routing lookup, these
accesses are in separate memories, allowing the scheme to be pipelined. By
pipelining this scheme we can achieve a routing rate of one lookup per memory
access. This longest prefix matching architecture provides high performance with
simple hardware by using the memory inefficiently. The total requirement of
memory is 33Mbytes of DRAM. Additionally, an important advantage of this
scheme is that by using simple hardware logic, the routing table update operation is
quite simple.
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
8 Appendix B
8.1 Block Diagrams of Queue Management Processes
Figure 8. 1 Block Diagram of Packet Entry Process
Appendix B
98
Figure 8. 2 Block Diagram of Enqueue Issue Process
Figure 8. 3 Block Diagram of Enqueue Execution Process (first stage)
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Figure 8. 4 Block Diagram of Enqueue Execution Process (second stage)
Figure 8. 5 Block Diagram of Dequeue Issue Process
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
Appendix B
Figure 8. 6 Block Diagram of Dequeue Execution Process (first stage)
Figure 8. 7 Block Diagram of Dequeue Execution Process (second stage)
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
100
2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
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2000 FORTH-ICS, Crete, Greece – TR-279 – November 2000
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