Anshul Kumar, CSE IITD ECE729 : Advance Computer Architecture Lecture 26: Synchronization, Memory Consistency 25 th March, 2010

Anshul Kumar, CSE IITD slide 2 Synchronization Problem Processes run on different processors independently At some point they need to know the status of each other for –communication –mutual exclusion etc Hardware primitives required for these operations

Anshul Kumar, CSE IITD slide 3 Consider an example Bank transaction from account number A : b = read_bal (A) b = b – debit_amt if b >= bmin update_bal (A, b)

Anshul Kumar, CSE IITD slide 4 Two concurrent transactions Transaction 1 : b1 = read_bal (A) b1 = b1 – debit_amt1 if b1 >= bmin update_bal (A, b1) Transaction 2 : b2 = read_bal (A) b2 = b2 – debit_amt2 if b2 >= bmin update_bal (A, b2)

Anshul Kumar, CSE IITD slide 5 Two concurrent transactions serialize reads Transaction 1 : b1 = read_bal (A) b1 = b1 – debit_amt1 if b1 >= bmin update_bal (A, b1) and writes Transaction 2 : b2 = read_bal (A) b2 = b2 – debit_amt2 if b2 >= bmin update_bal (A, b2)

Anshul Kumar, CSE IITD slide 6 Lock for mutual exclusion Transaction 1 : aquire: x1 = read (lock) if x1 = 1 then goto aquire set (lock) do transaction1 … release: clear (lock) Transaction 2 : aquire: x2 = read (lock) if x2 = 1 then goto aquire set (lock) do transaction2 … release: clear (lock) Transaction 1 : aquire: x1 = read (lock) if x1 = 1 then goto aquire set (lock) do transaction1 … … release: clear (lock) Transaction 2 : aquire: x2 = read (lock) if x2 = 1 then goto aquire set (lock) do transaction2 … … release: clear (lock)

Anshul Kumar, CSE IITD slide 7 Lock for mutual exclusion Transaction 1 : aquire: x1 = read (lock) if x1 = 1 then goto aquire set (lock) do transaction1 … release: clear (lock) Transaction 2 : aquire: x2 = read (lock) if x2 = 1 then goto aquire set (lock) do transaction2 … release: clear (lock) Transaction 1 : aquire: x1 = read (lock) if x1 = 1 then goto aquire set (lock) do transaction1 … … release: clear (lock) Transaction 2 : aquire: x2 = read (lock) if x2 = 1 then goto aquire set (lock) do transaction2 … … release: clear (lock)

Anshul Kumar, CSE IITD slide 8 Synchronization Primitives Hardware primitive required Should have atomic read+write operation Examples: – test&set –exchange –fetch&increment –load linked, store contitional

Spin Lock with Exchange Instr. Lock: 0 indicates free and 1 indicates locked Code to lock X : r2  1 lockit: r2  X ;atomic exchange if(r2  0)  lockit ;already locked locks are cached for efficiency, coherence is used Better code to lock X : lockit: r2  X ;read lock if(r2  0)  lockit ;not available r2  1 r2  X ;atomic exchange if(r2  0)  lockit ;already locked

Anshul Kumar, CSE IITD slide 10 LD Linked & ST conditional Simpler to implement atomic exchange r2  X using LL and SC try: r3  r2 ;move exchange value LL r1, X ;load linked SC r3, X ;store conditional if(r3=0)  try ;branch, store fails r2  r1 ;put loaded value in r2 fetch&increment using LL and SC try: LL r1, X ;load locked r3  r1 + 1 ;increment SC r3, X ;store conditional if(r3=0)  try ;branch, store fails

Anshul Kumar, CSE IITD slide 11 Spin Lock with LL & SC lockit: LL r2, X ;load locked if(r2  0)  lockit ;not available r2  1 SC r2, X ;store cond if(r2 = 0)  lockit ;branch store fails performance in presence of contention? spin lock with exponential back-off reduces contention

Anshul Kumar, CSE IITD slide 12 Barrier Synchronization lock (X) if(count=0)release  0 count++ unlock(X) if(count=total){count  0;release  1} else spin(release=1) 0 1

Anshul Kumar, CSE IITD slide 13 Improved Barrier Synch. local_sense  !local_sense lock (X) count++ unlock(X) if(count = total) {count  0;release  local_sense} else {spin(release = local_sense)} tree based barrier reduces contention

Anshul Kumar, CSE IITD slide 14 Memory Consistency Problem When must a processor see the value that has been written by another processor? Atomicity of operations – system wide? Can memory operations be re-ordered? Various models : models_tutorial.ps

Anshul Kumar, CSE IITD slide 15 ExampleExample P1: A = 0 P2: B = A = 1 B = 1 L1: if(B=0)S1 L2: if(A=0)S2 Which statements among S1 and S2 are done? Both S1, S2 may be done if writes are delayed

Anshul Kumar, CSE IITD slide 16 Sequential Consistency result of any execution is same as if the operations of all processors were executed in some sequential order operations of each processor occur in the order specified by its program - it requires all memory operations to be atomic - too restrictive, high overheads

Anshul Kumar, CSE IITD slide 17 Relaxing W  R order Loads are allowed to overtake stores  Write buffering is permitted 1.Total Store Ordering : Writes are atomic 2.Processor Consistency : Writes need not be atomic - Invalidations may gradually propagate

Anshul Kumar, CSE IITD slide 18 Relaxing W  R & W  W order Partial Store Ordering Loads are allowed to overtake stores Writes can be re-ordered Memory barrier or fence are used to explicitly order any operations Further improves the performance

Anshul Kumar, CSE IITD slide 19 ExamplesExamples P1 P2 A = 1; while(flag=0); flag = 1; print A; P1 P2 A = 1; print B; B = 1; print A; SC ensures that “1” is printed TSO, PC also do so PSO does not SC ensures that if B is printed as “1” then A is also printed as “1” TSO, PC also do so PSO does not

Anshul Kumar, CSE IITD slide 20 Examples - continued P1 P2 P3 A = 1; while(A=0); while(B=0); B = 1; print A; SC ensures that “1” is printed. TSO and PSO also do that but PC does not P1 P2 A = 1; B = 1; print B; print A; SC ensures that both can’t be printed as “0”. TSO, PC and PSO do not

Anshul Kumar, CSE IITD slide 21 Relaxing all R/W order Weak Ordering or Weak Consistency Loads and Stores are not restricted to follow an order Explicit synchronization primitives are used Synchronization primitives follow a strict order Easy to achieve Low overhead

Anshul Kumar, CSE IITD slide 22 Release Consistency Further relaxation of weak ordering Synch primitives are divided into aquire and release operations R/W operations after an aquire cannot move before it but those before it can be moved after R/W operations before a release cannot move after it but those after it can be moved before

Anshul Kumar, CSE IITD slide 23 WC and RC Comparison R/W … R/W … R/W … R/W synch R/W … R/W … R/W … R/W aquire release WC RC