cs 152 computer architecture and engineering lecture 11 - virtual memory and caches krste asanovic...
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CS 152 Computer Architecture and
Engineering
Lecture 11 - Virtual Memory and Caches
Krste AsanovicElectrical Engineering and Computer Sciences
University of California at Berkeley
http://www.eecs.berkeley.edu/~krstehttp://inst.eecs.berkeley.edu/~cs152
3/6/2008 CS152-Spring’08 2
Last time in Lectures 10
• Modern page-based virtual memory systems provide:– Translation
» to avoid memory fragmentation and provide each user with own virtual address space
– Protection
» to protect users from each other, and to protect operating system from users
– Virtual memory
» to allow main memory to act as a cache of a larger disk memory, equivalent
• Translation and protection information stored in page tables, held in main memory
• Translation and protection information cached in “translation lookaside buffer” (TLB) to provide single cycle translation+protection check in common case
3/6/2008 CS152-Spring’08 3
Today is a review
• Many of you had questions about virtual memory and interaction with caches
• Going to go back over core of last two lectures with some more interactive explanation
3/6/2008 CS152-Spring’08 4
Simple Base and Bound Translation
Load X
ProgramAddressSpace
BoundRegister
BoundsViolation?
Main
Mem
ory
currentsegment
BaseRegister
+
PhysicalAddressEffective
Address
Base and bounds registers are visible/accessible only when processor is running in the supervisor mode
Base Physical Address
Segment Length
3/6/2008 CS152-Spring’08 5
Separate Areas for Program and Data
What is an advantage of this separation?(Scheme used on all Cray vector supercomputers prior to X1, 2002)
Load X
ProgramAddressSpace
Main
Mem
ory
datasegment
Data BoundRegister
Effective AddrRegister
Data BaseRegister
+
BoundsViolation?
Program BoundRegister
ProgramCounter
Program BaseRegister
+
BoundsViolation?
programsegment
3/6/2008 CS152-Spring’08 6
Memory Fragmentation
As users come and go, the storage is “fragmented”. Therefore, at some stage programs have to be moved around to compact the storage.
OSSpace
16K
24K
24K
32K
24K
user 1
user 2
user 3
OSSpace
16K
24K
16K
32K
24K
user 1
user 2
user 3
user 5
user 4
8K
Users 4 & 5 arrive
Users 2 & 5leave
OSSpace
16K
24K
16K
32K
24K
user 1
user 48K
user 3
free
3/6/2008 CS152-Spring’08 7
• Processor generated address can be interpreted as a pair <page number, offset>
• A page table contains the physical address of the base of each page
Paged Memory Systems
Page tables make it possible to store the pages of a program non-contiguously.
0123
01
23
Address Spaceof User-1
Page Table of User-1
10
2
3
page number offset
3/6/2008 CS152-Spring’08 8
Private Address Space per User
• Each user has a page table • Page table contains an entry for each user page
VA1User 1
Page Table
VA1User 2
Page Table
VA1User 3
Page Table
Physi
cal
Mem
ory
free
OSpages
3/6/2008 CS152-Spring’08 9
Page Tables in Physical Memory
VA1
User 1
PT User 1
PT User 2
VA1
User 2
3/6/2008 CS152-Spring’08 10
Linear Page Table
VPN Offset
Virtual addressPT Base Register
VPN
Data word
Data Pages
Offset
PPNPPN
DPNPPN
PPNPPN
Page Table
DPN
PPN
DPNDPN
DPNPPN
• Page Table Entry (PTE) contains:– A bit to indicate if a page exists
– PPN (physical page number) for a memory-resident page
– DPN (disk page number) for a page on the disk
– Status bits for protection and usage
• OS sets the Page Table Base Register whenever active user process changes
3/6/2008 CS152-Spring’08 11
Size of Linear Page Table
With 32-bit addresses, 4-KB pages & 4-byte PTEs: 220 PTEs, i.e, 4 MB page table per user 4 GB of swap needed to back up full virtual address
space
Larger pages?• Internal fragmentation (Not all memory in a page is used)
• Larger page fault penalty (more time to read from disk)
What about 64-bit virtual address space???• Even 1MB pages would require 244 8-byte PTEs (35 TB!)
What is the “saving grace” ?
sparsity of virtual address usage
3/6/2008 CS152-Spring’08 12
Hierarchical Page Table
Level 1 Page Table
Level 2Page Tables
Data Pages
page in primary memory page in secondary memory
Root of the CurrentPage Table
p1
offset
p2
Virtual Address
(ProcessorRegister)
PTE of a nonexistent page
p1 p2 offset01112212231
10-bitL1 index
10-bit L2 index
3/6/2008 CS152-Spring’08 13
Address Translation & Protection
• Every instruction and data access needs address translation and protection checks
A good VM design needs to be fast (~ one cycle) and space efficient
Physical Address
Virtual Address
AddressTranslation
Virtual Page No. (VPN) offset
Physical Page No. (PPN) offset
ProtectionCheck
Exception?
Kernel/User Mode
Read/Write
3/6/2008 CS152-Spring’08 14
Translation Lookaside BuffersAddress translation is very expensive!
In a two-level page table, each reference becomes several memory accesses
Solution: Cache translations in TLBTLB hit Single Cycle Translation
TLB miss Page Table Walk to refill
VPN offset
V R W D tag PPN
physical address PPN offset
virtual address
hit?
(VPN = virtual page number)
(PPN = physical page number)
3/6/2008 CS152-Spring’08 15
TLB Designs
• Typically 32-128 entries, usually fully associative– Each entry maps a large page, hence less spatial locality across pages
more likely that two entries conflict
– Sometimes larger TLBs (256-512 entries) are 4-8 way set-associative
• Random or FIFO replacement policy
• TLB Reach: Size of largest virtual address space that can be simultaneously mapped by TLB
Example: 64 TLB entries, 4KB pages, one page per entry
TLB Reach = _____________________________________________?64 entries * 4 KB = 256 KB (if contiguous)
3/6/2008 CS152-Spring’08 16
Address Translation in CPU Pipeline
• Software handlers need restartable exception on TLB fault
• Handling a TLB miss needs a hardware or software mechanism to refill TLB
• Need mechanisms to cope with the additional latency of a TLB:
– slow down the clock
– pipeline the TLB and cache access
– virtual address caches
– parallel TLB/cache access
PCInst TLB
Inst. Cache D Decode E M
Data TLB
Data Cache W+
TLB miss? Page Fault?Protection violation?
TLB miss? Page Fault?Protection violation?
3/6/2008 CS152-Spring’08 17
Handling a TLB Miss
Software (MIPS, Alpha)TLB miss causes an exception and the operating system walks the page tables and reloads TLB. A privileged “untranslated” addressing mode used for walk
Hardware (SPARC v8, x86, PowerPC)A memory management unit (MMU) walks the page tables and reloads the TLB
If a missing (data or PT) page is encountered during the TLB reloading, MMU gives up and signals an exception for the original instruction
3/6/2008 CS152-Spring’08 18
Virtual Memory
• More than just translation and protection
• Use disk to extend apparent size of main memory
• Treat DRAM as cache of disk contents
• Only need to hold active working set of processes in DRAM, rest of memory image can be swapped to disk
• Inactive processes can be completely swapped to disk (except usually the root of the page table)
• Combination of hardware and software used to implement this feature
• (ATLAS was first implementation of this idea)
3/6/2008 CS152-Spring’08 19
Page Fault Handler
• When the referenced page is not in DRAM:– The missing page is located (or created)
– It is brought in from disk, and page table is updated Another job may be run on the CPU while the first job waits
for the requested page to be read from disk
– If no free pages are left, a page is swapped out Pseudo-LRU replacement policy
• Since it takes a long time to transfer a page (msecs), page faults are handled completely in software by the OS
– Untranslated addressing mode is essential to allow kernel to access page tables
3/6/2008 CS152-Spring’08 20
Caching vs. Demand Paging
CPU cacheprimarymemory
secondarymemory
Caching Demand pagingcache entry page framecache block (~32 bytes) page (~4K bytes)cache miss rate (1% to 20%) page miss rate (<0.001%)cache hit (~1 cycle) page hit (~100 cycles)cache miss (~100 cycles) page miss (~5M cycles)a miss is handled a miss is handled in hardware mostly in software
primarymemory
CPU
3/6/2008 CS152-Spring’08 21
Address Translation:putting it all together
Virtual Address
TLBLookup
Page TableWalk
Update TLBPage Fault(OS loads page)
ProtectionCheck
PhysicalAddress(to cache)
miss hit
the page is memory memory denied permitted
ProtectionFault
hardwarehardware or softwaresoftware
SEGFAULT
Restart instruction
3/6/2008 CS152-Spring’08 22
Protection and Translation
• These have been combined in a modern virtual memory system, but are really separate functions
• Question, does translation itself provide enough protection?
3/6/2008 CS152-Spring’08 23
CS152 Administrivia
• Tuesday Mar 18, Quiz 3– Virtual memory hierarchy lectures Lab 8-10
3/6/2008 CS152-Spring’08 24
Address Translation in CPU Pipeline
• Software handlers need restartable exception on TLB fault
• Handling a TLB miss needs a hardware or software mechanism to refill TLB
• Need mechanisms to cope with the additional latency of a TLB:
– slow down the clock
– pipeline the TLB and cache access
– virtual address caches
– parallel TLB/cache access
PCInst TLB
Inst. Cache D Decode E M
Data TLB
Data Cache W+
TLB miss? Page Fault?Protection violation?
TLB miss? Page Fault?Protection violation?
3/6/2008 CS152-Spring’08 25
Virtual Address Caches
• one-step process in case of a hit (+)• cache needs to be flushed on a context switch unless address
space identifiers (ASIDs) included in tags (-)• aliasing problems due to the sharing of pages (-)• maintaining cache coherence (-) (see later in course)
CPU PhysicalCache
TLB PrimaryMemory
VAPA
Alternative: place the cache before the TLB
CPU
VA
(StrongARM)VirtualCache
PATLB
PrimaryMemory
3/6/2008 CS152-Spring’08 26
Aliasing in Virtual-Address Caches
VA1
VA2
Page Table
Data Pages
PA
VA1
VA2
1st Copy of Data at PA
2nd Copy of Data at PA
Tag Data
Two virtual pages share one physical page
Virtual cache can have two copies of same physical data. Writes to one copy not visible
to reads of other!
General Solution: Disallow aliases to coexist in cache
Software (i.e., OS) solution for direct-mapped cache
VAs of shared pages must agree in cache index bits; this ensures all VAs accessing same PA will conflict in direct-mapped cache (early SPARCs)
3/6/2008 CS152-Spring’08 27
Concurrent Access to TLB & Cache
Index L is available without consulting the TLBcache and TLB accesses can begin simultaneously
Tag comparison is made after both accesses are completed
Cases: L + b = k L + b < k L + b > k
VPN L b
TLB Direct-map Cache 2L
blocks2b-byte block
PPN Page Offset
=hit?
DataPhysical Tag
Tag
VA
PA
VirtualIndex
k
3/6/2008 CS152-Spring’08 28
Virtual-Index Physical-Tag Caches: Associative Organization
Is this scheme realistic?
VPN a L = k-b b
TLBDirect-map
2L blocks
PPN Page Offset
=hit?
Data
Phy.Tag
Tag
VA
PA
VirtualIndex
kDirect-map
2L blocks
2a
=2a
After the PPN is known, 2a physical tags are compared
3/6/2008 CS152-Spring’08 29
Concurrent Access to TLB & Large L1The problem with L1 > Page size
Can VA1 and VA2 both map to PA ?
VPN a Page Offset b
TLB
PPN Page Offset b
Tag
VA
PA
Virtual Index
L1 PA cacheDirect-map
= hit?
PPNa Data
PPNa Data
VA1
VA2
3/6/2008 CS152-Spring’08 30
A solution via Second Level Cache
Usually a common L2 cache backs up both Instruction and Data L1 caches
L2 is “inclusive” of both Instruction and Data caches
CPU
L1 Data Cache
L1 Instruction
CacheUnified L2
Cache
RF Memory
Memory
Memory
Memory
3/6/2008 CS152-Spring’08 31
Anti-Aliasing Using L2: MIPS R10000
VPN a Page Offset b
TLB
PPN Page Offset b
Tag
VA
PA
Virtual Index L1 PA cacheDirect-map
= hit?
PPNa Data
PPNa Data
VA1
VA2
Direct-Mapped L2
PA a1 Data
PPN
into L2 tag
• Suppose VA1 and VA2 both map to PA and VA1 is already in L1, L2 (VA1 VA2)
• After VA2 is resolved to PA, a collision will be detected in L2.
• VA1 will be purged from L1 and L2, and VA2 will be loaded no aliasing !
3/6/2008 CS152-Spring’08 32
Virtually-Addressed L1:Anti-Aliasing using L2
VPN Page Offset b
TLB
PPN Page Offset b
Tag
VA
PA
VirtualIndex & Tag
PhysicalIndex & Tag
L1 VA Cache
L2 PA Cache L2 “contains” L1
PA VA1 Data
VA1 Data
VA2 Data
“VirtualTag”
Physically-addressed L2 can also be used to avoid aliases in virtually-addressed L1
3/6/2008 CS152-Spring’08 33
Variable-Sized Page Support
Level 1 Page Table
Level 2Page Tables
Data Pages
page in primary memorylarge page in primary memory page in secondary memoryPTE of a nonexistent page
Root of the CurrentPage Table
p1
offset
p2
Virtual Address
(ProcessorRegister)
p1 p2 offset01112212231
10-bitL1 index
10-bit L2 index
3/6/2008 CS152-Spring’08 34
Variable-Size Page TLBSome systems support multiple page sizes.
VPN offset
physical address PPN offset
virtual address
hit?
V RWD Tag PPN L
3/6/2008 CS152-Spring’08 35
Atlas Revisited
• One PAR for each physical page
• PAR’s contain the VPN’s of the pages resident in primary memory
• Advantage: The size is proportional to the size of the primary memory
• What is the disadvantage ?
VPN
PAR’s
PPN
3/6/2008 CS152-Spring’08 36
Hashed Page Table:Approximating Associative Addressing
hashOffset
Base of Table
+PA of PTE
PrimaryMemory
VPN PID PPN
Page Table
VPN d Virtual Address
VPN PID DPN
VPN PID
PID
• Hashed Page Table is typically 2 to 3 times larger than the number of PPN’s to reduce collision probability
• It can also contain DPN’s for some non-resident pages (not common)
• If a translation cannot be resolved in this table then the software consults a data structure that has an entry for every existing page
3/6/2008 CS152-Spring’08 37
Acknowledgements
• These slides contain material developed and copyright by:
– Arvind (MIT)
– Krste Asanovic (MIT/UCB)
– Joel Emer (Intel/MIT)
– James Hoe (CMU)
– John Kubiatowicz (UCB)
– David Patterson (UCB)
• MIT material derived from course 6.823
• UCB material derived from course CS252