ISCA－2000 海外調査報告

ＩＳＣＡ－２０００
海外調査報告
電気通信大学大学院情報システム学研究科
吉瀬謙二 [email protected]
1
会議の概要
• The 27th Annual International Symposium
on Computer Architecture, Vancouver
Canada 6月10日～1４日
– キーノート１
– パネル１
– 一般講演２９（採択率１７％）
• 参加者４４４人（日本人１３人，大学から４
人）
• http://www.cs.rochester.edu/~ISCA2k/
2
紹介する文献
• Multiple-Banked Register File Architecture
• On the Value Locality of Store Instructions
• Completion Time Multiple Branch Prediction for
Enhancing Trace Cache Performance
• Circuits for Wide-Window Superscalar Processor
• Trace Preconstruction
• A Hardware Mechanism for Dynamic Extraction and
Relayout of Program Hot Spots
• A Fully Associative Software-Managed Cache Design
• Performance Analysis of the Alpha 21264-based Compaq
ES40 System
• Piranha: A Scalable Architecture Based on Single Chip
Multiprocessing
3
Session 8 – Microarchitecture Innovations
Multiple-Banked Register File
Architecture
Jose-Lorenzo Cruz et al.
Universitat Politecnica de Catalunya, Spain
ISCA-2000 p.316-325
4
レジスタファイルの構成
Tri-state
buffer
Read
5
out1
Reg31
64
Register
RegNo
value
File
32 Registers
64
Reg2
Reg1
Reg0
• the number of registers
• the number of ports
(Read, Write)
5
decoder
5
研究の動機
• The register file access time is one of the
critical delays
• The access time of the register file depends
on the number of registers and the
number of ports
– Instruction window -> registers
– Issue width -> ports
6
研究の目的
request
RegFile
value
RegFile
Machine Cycle
• レジスタファイルのポート数を増やす
• シングル・サイクルでアクセスできる
レジスタファイルに近づける
7
Impact of Register File Architecture
4.00
SEPCint95
3.75
3.50
IPC
3.25
3.00
2.75
2.50
2.25
2.00
48
64
96
128
160
192
224
256
Physical Register File Size
8
Observation
• Processor needs many physical registers but a
very small number are actually required at
a given moment.
–
–
–
–
Registers with no value
Value used by later instructions
Last-use and overwrite
Bypass only or never read
9
Multiple-Banked Register File
Architecture
Bank 2
Bank 1
uppermost
level
Bank 2
lowest
level
Bank 1
(a) one-level
(b) multi-level (register file cache)
10
Register File Cache
Bank 1
Bank 2
• The lowest level is
always written.
• Data is moved only from
uppermost Level lower to upper level.
16 registers • Cached in upper level
based on heuristics.
lowest Level • There is a prefetch
128 registers
mechanism.
11
Caching and Fetching Policy
Locality properties of registers and memory
are very different.
• Non-bypass caching
– バイパスロジックから読まれていない結果のみを上位レベルに
格納
• Ready caching
– まだ発行されていない命令で必要とされている値のみを上位レ
ベルに格納
• Fetch-on-demand
– 必要となった時点で値を上位レベルに転送
• Prefetch-first-pair -> next slide
12
Prefetch-first-pair
(1) p1 = P2 + P3
(2) P4 = P3 + P6 •
(3) P7 = P1 + P8 •
命令(1) から(3) は，プロセッサ内でリ
ネームステージを経過している．
P1 ～ P8 は，ハードウェアによって変換
された物理的なレジスタの番号
命令(1)の結果レジスタ P1 を最初に利用する命令(3)
のもう一つのレジスタ P8 をプリフェッチする．
命令(1)が発行される際に P8 をプリフェッチ
13
評価結果 (conf. C3)
• One-cycle single-banked
– Area 18855, cycle 5.22 ns (191 MHz)
– Read 4 port, Write 3 port
• Two-cycle single-banked
– Area 18855, cycle 2.61 ns (383 MHz)
– Read 4 port, Write 3 port
• Register file cache
– Area 20529, cycle 2.61 ns (382 MHz)
– Upper: Read 4 port, Write 4 port
– Lower: Write 4 port, Bus 2
14
評価結果
3.0
Relative Instruction Throughput
1-cycle
non-bypass caching + prefetch-first-pair
2-cycle, 1-bypass
2.0
1.0
0.0
SpecInt95
SpecFP95
15
研究の新規性
• Register File Cache の提案
– 高速動作が可能
– 上位レベルのミス率を削減することで，ア
クセスのサイクル数を１に近づける．
• ２つのキャッシュ方式と，２つのフェッチ
方式の提案
• エリアとサイクル時間を考慮した性能
評価
16
研究へのコメント
• 巨大なレジスタファイル，ポート数の増
加，アクセス時間の低減という要求
• 従来単純な構成だったレジスタファイル
に関しても，キャッシュのように複数階層
の構成が必要
• 今後，大規模なＩＬＰアーキテクチャにお
ける複雑化は避けられない？
17
Session 5a – Analysis of Workloads and Systems
On the Value Locality of Store
Instructions
Kevin M. Lepak et al.
University of Wisconsin
ISCA-2000 p.182-191
18
Value Locality (Value Prediction)
• Value locality
– a program attribute that describes the likelihood of
the recurrence of previously-seen program values
• ある命令が前回生成した演算結果（データ値）
と，今回生成するデータ値には関連がある．
(1) p1 = P2 + P3
(2) P4 = P3 + P6
(3) P7 = P1 + P8
P1の演算結果の履歴
... 1, 2, 3, 4, 5, 1, 2, 3, 4, 5, 1 ?
19
研究の目的
• Much publication has focused on load
instruction outcome.
• Examine the implications of store value
locality
– Introduce the notion of silent stores
– Introduce two new definitions of
multiprocessor true sharing
– Reduce multiprocessor data and address
traffic
20
Memory-centric and producercentric Locality
• Program structure store value locality
– The locality of values written by a particular
static store.
• Message-passing store value locality
– The locality of values written to a particular
address in data memory.
21
20%-62% of stores are silent stores
Silent store is a store that does not change the system state.
100
90
70
60
50
40
30
20
10
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ocea
barne
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oltp
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perl
ijpeg
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comp
gcc
m88k
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go
Silent Stores (%)
80
22
Silent Store Removal Mechanism
• Realistic Method
– All previous store addresses must be known.
– Load the data from the memory subsystem.
– If the data values are equal, the store is
update-silent.
• Remove from the LSQ and flag the RUU entry
– If store is silent, the store retires with no
memory access.
23
Evaluation Results
• Writeback Reduction
– Range in reduction from 81% to 0%
– Average 33% reduction
• Instruction Throughput
– Speedups of 6.3% and 6.9% for realistic
and perfect removal mechanisms.
24
New Definition of False Sharing
• Multiprocessor applications
– All of the previous definitions rely on the
specific addresses in the same block.
– No attempt is made to determine when the
invalidation of a block is unnecessary because
the value stored in the line does not change.
– Silent stores and stochastically silent stores
25
Address-based Definition of Sharing
• Cold Miss
[Dubois 1993]
– The first miss to a given block by a processor
• Essential Miss
– A cold miss is an essential miss
– If during the lifetime of a block, the processor
accesses a value defined by another processor since
the last essential miss to that block, it is an essential
miss.
• Pure True Sharing miss (PTS)
– An essential miss that is not cold.
• Pure False Sharing miss (PFS)
– A non-essential miss.
26
Update-based False Sharing (UFS)
• Essential Miss
– A cold miss is an essential miss
– If during the lifetime of a block, the
processor accesses an address which has
had a different data value defined by
another processor since the last essential
miss to that block, it is an essential miss.
27
Stochastic False Sharing (SFS)
• It seems intuitive that
– if we define false sharing to
compensate for the effect of silent
stores that we could also define it in the
presence of stochastically silent stores
(values that are trivially predictable
via some mechanism)
28
研究の新規性
• Overall characterization of store value locality
• Notion of silent stores
• Uniprocessor speedup by squashing silent
stores
• Definition of UFS and SFS
• How to exploit UFS to reduce address and data
bus traffic on shared memory multiprocessors
29
研究へのコメント
• ストア命令のデータ値の局所性に関す
る様々な事柄をまとめている．
• 評価は初期的な構成のもので，今後の
研究の動機付けとなる．
• 並列計算機における局所性の利用に関
しては詳細な検討が必要
30
Session 2a – Exploiting Traces
Completion Time Multiple Branch
Prediction for Enhancing Trace
Cache Performance
Ryan Rakvic et al.
Carnegie Mellon University
ISCA-2000 p.47-58
31
Branch Prediction and
Multiple Branch Prediction
Basic Block
Not
Taken
Not
Taken
Taken
Taken
Control Flow Graph
Branch Prediction:
T or N?
Taken
Taken
Multiple Branch Prediction:
(T or N) (T or N) (T or N) ?
32
動機と目的
• Wide Instruction Fetching
– 4-way -> 8-way
• Multiple branch prediction
– Branch Execution Rate: about 15%
– One branch prediction per cycle is not enough.
• Tree-Based Multiple Branch Predictor (TMP)
33
用いられている分岐予測の例
gshare
Two-bit bimodal branch predictor
B-3
Taken
Tag
2BC
Not taken
B-2
B-1
Not Taken
index
10
N
00
N
01
T
T
N
Not Taken
B0
Not taken
10
Taken
Taken
TNN
N
T
T
11
Taken
B0
Global history
34
Tree-Based Multiple Branch
Predictor (TMP)
B-3
Taken
Tree-based Pattern History Table (Tee-PHT)
B-2
Tree
TTNT
Not Taken
B-1
Not Taken
B0
B1
Predicted path
TGlobal
N N history
B0
B2
B3
B0
Tree(i)
B1
B2
B3
B4
B4
35
Tree-based Pattern History Table
Two-bit bimodal branch predictor
Not taken
Tree-based
Pattern History Table (Tee-PHT)
tag Predicted Path
Tree
N
01
T
Not taken
00
01
10
11
Taken
Taken
11 T
NTNT
N
00
T
11
36
Updating of 2-bit bimodal tree-node
Old predicted path: NTNT
New predicted path: TNTT
Recently completed path: TNNN
Not taken
N
00
Not taken
N
01
T
Taken
N
T
T
01
10
T
N
10
N
T
11 T
10
11
01
Taken
N
11
00
10
T
11
10
01
11
37
Tree-PHT with second level PHT(NodePHT) for tree-node prediction
Node Pattern History Table(Node-PHT)
Tree-based
Pattern History Table (Tee-PHT)
tag Predicted Path
01...1 n bits of local history
Tree
N
2-bit
bimodal
T
01...1
• global(g)
• per-branch(p)
• shared(s)
38
研究の新規性と評価結果
• TMPの提案
– Three-level branch predictor
– Maintain a tree structure
– Completion time update
• TMPs-best
(shared)
– The number of entries in the Tree-PHT: 2K
– Local history bit: 6
– 72KB Memory
•
•
•
•
96%: 1 block
93%: 2 consecutive blocks
87%: 3 consecutive blocks
82%: 4 consecutive blocks
39
研究へのコメント
• サイクル当たり複数の分岐命令の分岐
先を予測するために，３レベルの予測
機構を提案
• 分岐予測はさらに複雑になるが
• 着実な性能向上
40
Session 6 – Circuit Considerations
Circuits for Wide-Window
Superscalar Processor
Dana S. Henry, Bradley C. Kuszmaul,
Gabriel H. Loh and Rahul Sami
ISCA-2000 p.236-247
41
Instruction Window
アウトオブオーダ実行のスーパースカラプロセッサと命令ウィンドウ
Src1 Src2
(1) p1 = P2 + P3
命令供給
Src1-Tag
Src1-Tag
Src1-Tag
Src1-Tag
実行結果（タグ，値）
命令
Valid
Valid
Valid
Valid
命令
ウィンドウ
Src2-Tag
Src2-Tag
Src2-Tag
Src2-Tag
命令，
データ
Valid
Valid
Valid
Valid
Op
Op
Op
Op
実行
• Wake-up
• Schedule
42
研究の動機
実行命令列
• 命令ウィンドウを大きくすることで命令レ
ベル並列性利用の可能性が増大
• Alpha 21264 のウィンドウサイズは35
• MIPS R10000 のウィンドウサイズは??
• サイズの大きい命令ウィンドウを構成す
ることは困難
命令
ウィンドウ • Power4, two 4-issue processors
• Intel Itanium, VLIW techniques
43
研究の目的
• 高速動作する大きなサイズ（１２８）の命令
ウィンドウを実現する
• Log-depth cyclic segment prefix (CSP) circuit
の提案
• Log-depth cyclic segment prefix circuitとサイ
クル時間の関係を議論
• 大きなサイズの命令ウィンドウによる性能向
上を議論
44
Gate-delay cyclic segmented
prefix (CSP)
Tail
out 0
in 0
s0
Head
out 1
in 1
s1
out 2
in 2
s2
out 3
in 3
s3
An 4-entry wrap-around
Reordering buffer with
Adjacent, linear gated-delay
cyclic segmented prefix.
45
Commit login using CSP
Tail
Head
done
out 0
in 0
s0
done
out 1
in 1
s1
done
done
out 2
in 2
s2
done
Not
done
out 3
in 3
s3
Tail
Head
done
Not
done
46
Wake-up logic for logical register R5
Tail
out 0
D: R4=R5+R7 in 0
s0
Not done
Head
out 1
A: R5=R8+R1 in 1
s1
done
out 2
B: R1=R5+R1 in 2
s2
Not done
out 3
C: R5=R3+R3 in 3
s3
Not done
47
Scheduler logic scheduling two FUs
Tail
Head
D: R4=R5+R7
request
A: R5=R8+R1
request
B: R1=R5+R1
2
+
2
+
1
+
C: R5=R3+R3
request
1
Logarithmic
gate-delay
implementations
 p.240 - 241 参照
+
48
評価結果
•
•
•
•
•
１２８エントリの命令ウィンドウを設計
Commit logic: 1.41 ns (709 MHz)
Wakeup logic: 1.54 ns (649 MHz)
Schedule logic: 1.69 ns (591 MHz)
現在のプロセス技術を用いて500MHz
以上の動作速度を達成
49
研究へのコメント
• １２８エントリの命令ウィンドウの実現可
能性を示した．
• 従来，命令ウィンドウのエントリ数を増や
すことは困難と考えられてきた．
• この点を覆すという意味で面白い．
50
Session 2a – Exploiting Traces
Trace Preconstruction
Quinn Jacobson et al.
Sun Microsystems
ISCA-2000 p.37-46
51
Trace Cache

Traces are snapshots of short segments
of the dynamic instruction stream.
I-Cache
Branch
Predict
Trace
Constructor
Instruction
Fetch
Next
Trace
Predictor
Trace Cache
Execution
Engine
52
Trace Cache Slow Path


If the trace cache does not have the needed
trace, the slow path is used.
When the trace cache is busy, the slow path
hardware is idle.
I-Cache
Branch
Predict
Trace
Constructor
Slow Path
Instruction
Fetch
Next
Trace
Predictor
Trace Cache
Execution
Engine
53
目的
• Trace cache enables high bandwidth, low
latency instruction supply.
• Trace cache performance may suffer due
to capacity and compulsory misses.
• By performing a function analogous to
prefetching, trace preconstruction
augments a trace cache.
54
Preconstruction Method

For preconstruction to be
successful


Region start points must identify
instructions that the actual execution
path will reach in the future.
Heuristic


loop back edge
procedure call
55
Preconstruction Example
Region Start Point:
a
JAL
h
i
Br2
j
Br4
b
d e g
c
Br1
d
e
f
g
CFG
JAL Br1
Br2
a b c c c c c d e g h i i i
Br3
JMP
f
g
Region 2
j
Region 3
h i i i i i
j
j
j
j
Region 1
56
Preconstruction Efficiency
• Slow-path dynamic branch predictor
– To reduce the number of paths
– Assume a bimodal branch predictor
– Only the strongly biased path is followed during
preconstruction.
• Trace Alignment
– ..cccdeg...
– <c c c> <d e g> or <. c c> <c d e> or ...
– In order for pre-constructed trace to be useful, it must
align with the actual execution path
57
Conclusions
• One implementation of trace
preconstruction
– SPECint95 benchmarks with large working
set sizes
– Reduce the trace cache miss rates from 30%
to 80%
– 3% to 10% overall performance improvement
58
Session 2a – Exploiting Traces
A Hardware Mechanism for
Dynamic Extraction and
Relayout of Program Hot Spots
Matthew C. Merten et al.
Coordinated Science Lab
ISCA-2000 p.59-70
59
Hot Spot
• 頻繁に実行される命令コードHot Spot
– 10対90, 1対99 の法則
60
目的
• Dynamic extraction and relayout of
program hot spots
– A hardware-driven profiling scheme for
identifying program hot spots to support
runtime optimization, ISCA 1999
• Improve instruction fetch rate
61
Branch Address
=
&
+I
0
-D
1
Touched Bit
Call ID
Remapped FT Addr
FV Valid Bit
Remapped Taken Addr
Taken Valid Bit
Cand Flag
Taken Counter
Reset Timer
Exec Counter
Refresh Timer
Branch Tag
Branch Behavior Buffer with
new fields shaded in gray
Hot Spot
Detection
Counter
Saturating
Adder
At Zero:
62
Hot Spot Detected
Memory-Based Code Cache
• Use a region of memory called a code cache to
contain the translated code.
• Standard paging mechanisms are used to
manage a set of virtual pages reserved for the
code cache.
• The code cache pages can be allocated by the
operating system and marked as read-only
executable code.
• The remapping hardware must be allowed to
write this region.
63
Hardware Interaction
Fetch
Icache
Branch
Predictor
Decode ... Execute
In-order Retire
Profile
Information
BTB
TGU
BBB
Update
Memory
Trace
Generation
Unit
Branch
Behavior
Buffer
Code Cache
64
Code Deployment
• Original code cannot be altered
• Transfer to the optimized code is handled
by the Branch Target Buffer.
– BTB target for the entry point branch is
updated with the address of the entry point
target in the code cache.
• Optimized code consists of only the
original code.
65
Trace Generation Algorithm
• Scan Mode
– Search for a trace entry point. This is the initial
mode following hot spot detection
• Fill Mode
– Construct a trace by writing each retired
instructions into the memory-based code cache.
• Pending Mode
– Pause trace construction until a new path is
executed.
66
State Transition
From Profile Mode
Hot spot detection
New Entry Point
(jcc || jmp) &&
candidate && taken
Fill
Mode
Scan
Mode
Cold Branch
(jcc || jmp) && !candidate
End Trace
2-> jcc && candidate && both_dir_remapped
3-> jcc && !candidate && off_path_branch_cnt
> max_off_path_allowed
4-> red && red_addr_mismatch
5-> jcc && candidate && recursive_call
Merge jcc && candidate
&& other_dir_not_remapped
&& exec_dir_reampped
Pending
Mode
Continue jcc && addr_matches_pending_target
&& exec_dir_not_remapped
67
Trace Example with Optimization
(1/3)
Branch Taken
A
B
C1
A1
B1
A2
C2
A3
B2
C3
D1
Execution order during remapping
C
D
(b) Original code layout
68
Trace Example with Optimization
(2/3)
Entrance: C1
RM-A1
RM-B1
Branch Taken
C1
A1
B1
A2
C2
A3
B2
C3
D1
Execution order during remapping
RM-A2
(c) Trace generated by basic remapping
RM-C2
RM-D1
69
Trace Example with Optimization
(3/3)
Entrance: C1
RM-A1
RM-B1
RM-A2
RM-C3
C1
A1
B1
A2
C2
A3
B2
C3
D1
Execution order during remapping
RM-A3
RM-B2
RM-C3
(d) The application of two remapping
optimizations, patching and branch replication
RM-D1
70
Fetched IPC for various fetch
mechanisms
8
Fetch IPC
6
4
2
0
IC:64KB
IC:64KB
remap
IC:64KB,
TC:8KB
IC:64KB,
TC:8KB
remap
IC:64KB,
TC:128KB
IC:64KB,
TC:128KB
remap
71
Conclusion
• Detect the hot spots to perform
– code straightening
– partial function inlining
– loop unrolling
• Achieve significant fetch performance
improvement at little extra hardware cost
• Create opportunities for more aggressive
optimizations in the future
72
研究へのコメント
• 昨年度の Hot Spot の検出から， Hot
Spotの最適化に踏み込んだ研究
• 最適化のアルゴリズムが複雑？
• 最適化のバリエーションは無限？
– ソフトウェアによる最適化
– 命令セット間の変換
–等
73
Session 3 – Memory Hierarchy Improvement
A Fully Associative SoftwareManaged Cache Design
Erik G. Hallnor et al.
The University of Michigan
ISCA-2000 p.107-116
74
Data Cache Hierarchy
MPU
100 million
Transistor
L1 Data Cache
L2 Data Cache
1 MB
Off-chip
Data Cache
2-way set-associative data cache
Data Address
Tag
Offset
Tag
Tag
Tag
Tag
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Tag
Tag
Tag
Tag
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Data
Main Memory (DRAM)
75
研究の目的
• Processor-memory gap
– miss latencies approaching a thousand
instruction execution time
• On-chip SRAM caches in excess of one
megabyte [Alpha 21364]
• Re-examination of how secondary caches
are organized and managed
– Practical, fully associative, software managed
secondary cache system
76
Indirect Index Cache (IIC)
TE
Data Address
Tag
TAG
STATUS
INDEX
REPL
Offset
Secondary Storage
for Chaining
TE
Chain
Primary Hash Table
hash
TE
TE
TE
TE
Chain
=?
Data Array
Data
Hit?
=?
=?
=?
=?
Hit?
Hit?
Hit?
Hit?
Data Value
77
Generational Replacement
• Blocks are grouped into a small number of
prioritized pools.
– Blocks that are referenced regularly are
promoted into higher-priority pools.
– Demote unreferenced blocks into lower-priority
pools.
Ref = 0
Ref = 1
Pool 0
(lowest priority)
Fresh pool
Ref = 1
Ref = 1
Pool 1
Ref = 1
Pool 3
(Highest priority)
Pool 2
78
Ref = 0
Ref = 0
Ref = 0
Generational Replacement
Algorithm
• Each pool is variable-length FIFO queue of blocks.
• On a hit, only the block's reference bit is updated.
• On a each miss, the algorithm checks the head of
each pool FIFO.
– Head block's reference bit is set -> Next higher-priority
pool, reference bit=0
– Head block's reference bit is not set -> Next lowerpriority pool, reference bit=0
79
Evaluation Result and
Conclusion
• Substantial reduction in miss rates(8-85%)
relative to a conventional 4-way associative
LRU cache.
• IIC/generational replacement cache could
be competitive with a conventional cache at
today's DRAM latencies, and will
outperform as CPU-relative latencies grow.
80
研究へのコメント
• 大規模なオンチップの２次キャッシュを
想定
• ソフトウェアが支援しているが，プログラ
マや実行コードからはキャッシュに見え
る．
• 命令セットやプログラマの支援を考慮し
た検討も必要ではないか．
81
Session 5a – Analysis of Workloads and Systems
Performance Analysis of the
Alpha 21264-based Compaq
ES40 System
Zarka Cvetanovic, R.E.Kessler
Compaq Computer Corporation
ISCA-2000 p.192-202
82
研究の動機
Shared Memory Multiprocessor Comparison
3000
Compaq ES40/21264 667MHz
2500
HP PA-8500 440MHz
SPECfp_rate95
SUN USparc-II 400MHz
Compaq ES40
2000
1500
1000
500
0
1-CPU
2-CPU
4-CPU
83
研究の目的
• Evaluation of Compaq ES40 shared
memory multiprocessor
– Up to four Alpha 21264 CPU
• Quantitatively show the performance
– Instruction Per Cycle
– Branch mispredicts
– Cache misses
84
Alpha 21264 Microprocessor
• 15-million transistors, 4-way out-oforder superscalar
• 80 in-flight instructions
• 35 instruction window
• hybrid (Local and global) branch
prediction
• two-clusters integer execution core
• load hit/miss, store/load prediction
85
Instruction Cache Miss Rate
Comparison
8KB direct-mapped -> 64KB two-way associative
TPM
Transactions Per Minute
SPECint95
Alpha 21164
SPECfp95
Alpha 21264
0
5
10
15 20 25 30 35 40 45
Icache Misses per 1000 Retires
50
55
60
86
Branch Mispredicts Comparison
2bit predictor -> Local and global hybrid predictor
TPM
SPECint95
Alpha 21164
SPECfp95
Alpha 21264
0
2
4
6
8
10
12
14
16
18
Branch Mispredicts per 1000 Retires
20
22
87
IPC Comparison
Alpha 21164
TPM
Alpha 21264
SPECint95
SPECfp95
0.0
0.5
1.0
1.5
Retired Instructions Per Cycle
2.0
88
Compaq ES40 Block Diagram
L2
8MB
CPU
Alpha 21264
L2
CPU
Alpha 21264
L2
L2
64b
8 Data
Switches
Memory Bank
Memory Bank
(Crossbar-based)
Memory Bank
Memory Bank
CPU
Alpha 21264
CPU
Alpha 21264
256b
Control Chip
64b
PCI-Chip
PCI
PCI-Chip
PCI
89
Inter-processor Communication
• Write-invalidate cache coherence
– 21264 passes the L2 miss requests to the
control chip.
– The control chip simultaneously forwards the
request to DRAM and other 21264s.
– Other 21264s check for necessary coherence
violations and respond.
– The control chip responds with data from
DRAM or another 21264.
90
STREAM Copy Bandwidth
• The STREAM benchmark
• Measure the best memory bandwidth in
megabytes per second
–
–
–
–
COPY: a(i) = b(i)
SCALE: a(i) = q*b(i)
SUM: a(i) = b(i) + c(i)
TRAID: a(i) = b(i) + q*c(i)
– The general rule for STREAM is that each array must be
at least 4x the size of the sum of all the last-level caches
used in the run, or 1 Million elements -- whichever is larger.
91
STREAM Copy Bandwidth
< 3 GBytes / sec
Memory Copy Bandwidth (MB/sec)
3000
2750
2500
2250
Compaq ES40/21264 667MHz
Compaq ES40/21264 500MHz
SUN Ultra Enterprise 6000
AlphaServer 4100 5/600
2547
2000
1750
1500
1197
1250
1000
750
500
250
0
470
263
1-CPU
2-CPU
3-CPU
4-CPU
92
研究結果
•
•
•
•
Five times the memory bandwidth
Microprocessor enhancement
Compiler enhancement
Compaq ES40 provides 2 to 3 times the
performance of the AlphaServer 4100
93
研究へのコメント
• ネットワークにクロスバを採用すること
で性能向上を達成
• Alpha 21264 の挙動を紹介する文献と
して興味深い
• 実マシンの詳細な性能評価として興味
深い
• アイデアの新しさはない
94
Session 7 – Extracting Parallelism
Piranha: A Scalable Architecture
Based on Single Chip
Multiprocessing
Luiz Andre Barroso
Compaq Computer Corporation
ISCA-2000 p.282-293
95
研究の動機
• Complex processors
– Higher development cost
– Longer design times
• On-line transaction Processing (OLTP)
– Little instruction-level parallelism
– Thread-level or process-level parallelism
• Semiconductor integration density
• Chip multiprocessing (CMP)
96
研究の目的
• Piranha: a research prototype at Compaq
– Targeted at parallel commercial workloads
– Chip multiprocessing architectures
– Small team, modest investment, short design
time
• General-purpose microprocessors or
Piranha?
97
Output
Queue
Router
Input
Queue
Packet Switch
Interconnect Links
Single-chip Piranha processing node
Home
Engine
System
Control
Direct
Rambus Array
CPU7
iL1 dL1
iL1 dL1
Intra-Chip Switch
Remote
Engine
Chip
CPU0
L20
L27
MC0
MC7
0
1
RDRAM
RDRAM
31
RDRAM
0
1
RDRAM
RDRAM
31 RDRAM
98
Alpha CPU Core and L1 Caches
• CPU: Single-issue, in-order, 500MHz
datapath
• iL1, dL1: 64KB two-way set-associative
– Single-cycle latency
• TBL: 256 entries, 4-way set-associative
• 2-bit state field per cache line for MESI
protocol
99
Intra-Chip Switch (ICS)
•
•
•
•
ICS manages 27 clients.
Conceptually, ICS is a crossbar.
Eight internal datapaths
Capacity of 32 GB/sec
– 500MHz, 64bit bus (8Byte), 8 datapaths
– 500MHz x 8Byte x 8 = 32 GB/sec
– 3 times the available memory bandwidth
100
Second-Level Cache
• L2: 1MB unified instruction/data cache
– Physically partitioned into eight banks
• Each bank is 8-way set-associative
• Non-inclusive on-chip cache hierarchy
– Keep a duplicate copy of the L1 tags and state
• L2 controllers are responsible for coherence
within a chip.
101
Memory Controller
• Eight memory controllers
• Each RDRAM channel has a maximum
1.6GB/sec
• Maximum local memory bandwidth of
12.8GB/sec
102
Single-chip Piranha versus out-oforder processor
400
300
On-Line Transaction
Processing (OLTP)
250
233
Normalized Execution Time
350
350
DSS
(Query 6 of
the TPC-D)
191
200
145
150
100
100
100
44
34
50
0
P1
500MHz
1-issue
INO
OOO
1GHz
1-issue
1GHz
500MHz
4-issue 1-issue
P8
P1
INO
500MHz 1GHz
1-issue
1-issue
OOO
1GHz
4-issue
P8
500MHz
103
1-issue
Example System Configuration
P chip
P chip
P chip
I/O chip
P chip
P chip
P chip
I/O chip
A Piranha system with six processing(8 CPUs each) and two I/O chips.
104
研究の新規性
• Detailed evaluation of database workloads in
the context of CMP
• Simple processor, standard ASIC design
– short time, small team size, small investment
105
研究へのコメント
• トランザクション処理では命令レベル
の並列性を利用できない。
• CMPが強い？
• 普及するのは時間の問題？
106

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