1. Module 6: Process Synchronization

2. Module 6: Process Synchronization
Background
The Critical-Section Problem
Peterson’s Solution
Synchronization Hardware
Semaphores
Classic Problems of Synchronization
Monitors
Synchronization Examples
Atomic Transactions

3. Objectives
To introduce the critical-section problem, whose solutions can be used to ensure the consistency of shared data
To present solutions of the critical-section problem

4. Background
Concurrent access to shared data may result in data inconsistency
Maintaining data consistency requires mechanisms to ensure the orderly execution of cooperating processes
Suppose that we wanted to provide a solution to the consumer-producer problem that fills all the buffers. We can do so by having an integer count that keeps track of the number of full buffers. Initially, count is set to 0. It is incremented by the producer after it produces a new buffer and is decremented by the consumer after it consumes a buffer.

5. Producer while (count == BUFFER.SIZE) ; // do nothing
// add an item to the buffer buffer[in] = item;
in = (in + 1) % BUFFER.SIZE;
++count;

6. Consumer while (count == 0) ; // do nothing
// remove an item from the buffer item = buffer[out];
out = (out + 1) % BUFFER.SIZE;
--count;

7. Race Condition
count++ could be implemented as register1 = count register1 = register count = register1
count-- could be implemented as register2 = count register2 = register count = register2
Consider this execution interleaving with “count = 5” initially:
T0: producer execute register1 = count {register1 = 5} T1: producer execute register1 = register {register1 = 6} T2: consumer execute register2 = count {register2 = 5} T3: consumer execute register2 = register {register2 = 4} T4: producer execute count = register1 {count = 6 } T5: consumer execute count = register2 {count = 4}

8. Solution to Critical-Section Problem
Mutual Exclusion - If process Pi is executing in its critical section, then no other processes can be executing in their critical sections.
Progress - If no process is executing in its critical section and there exist some processes that wish to enter their critical section, then the selection of the processes that will enter the critical section next cannot be postponed indefinitely.
Bounded Waiting - A bound must exist on the number of times that other processes are allowed to enter their critical sections after a process has made a request to enter its critical section and before that request is granted.
Assume that each process executes at a nonzero speed
No assumption concerning relative speed of the N processes

9. Structure of a Typical Process

10. Two Process Solutions Algorithm 1 do { while turn ≠ i do noop;
Critical section
turn = j;
Remainder section
} while (TRUE);
Here Pi and Pj share an integer variable turn initialized to 0. If turn = i, Pi can execute in its critical section.
Mutual exclusion is assured but progress requirement is not since it requires strict alternation of processes in the execution of the critical section.
E.g. if turn = i and Pi does not want to enter the CS and Pj wants to enter the CS, then it can’t.

11. Two Process Solutions Algorithm 2
Pi and Pj share an array: boolean flag[2];
flag[0] and flag[1] are initialized to false. If flag[i] is true, then it indicates Pi is ready to enter its critical section.
do {
flag[i] = true;
while flag[j] do noop;
Critical section
flag[i] = false;
Remainder section
} while (TRUE);
Here too the progress requirement is not met even though mutual exclusion is satisfied.
Need to remember that P0 and P1 are executing concurrently.
Say, T0: P0 sets flag[0] to true
T1: P1 sets flag[1] to true
Now both P0 and P1 are looping infinitely in their while loops.

12. Peterson’s Solution Two process solution
Assume that the LOAD and STORE instructions are atomic; that is, cannot be interrupted.
The two processes share two variables:
int turn;
boolean flag[2]
The variable turn indicates whose turn it is to enter the critical section.
The flag array is used to indicate if a process is ready to enter the critical section. flag[i] = true implies that process Pi is ready!

13. Algorithm for Process Pi
while (true) {
flag[i] = true;
turn = j;
while (flag[j] && turn == j);
critical section
flag[i] = false;
remainder section }

14. Peterson’s Solution contd..
This algorithm meets all 3 conditions that a solution to CS problem must meet.
Say Pi wants to enter its CS. It can be prevented only if flag[j] is true and turn = j. If Pj is not ready to enter its CS, then flag[j] is false and Pi will enter its CS.
If Pj has set flag[j] to true and is also executing in its while loop, then either turn = i or turn = j (depends on whether Pi or Pj accessed turn last) and so mutual exclusion is assured.
If say Pj enters it CS then upon exiting it will set flag[j] to false so Pi can now enter its CS.
If Pj resets flag[j] to true if it wants to enter CS again, then it will also turn = i. So Pi will enter its CS (progress is assured) after at most 1 entry by Pj (bounded waiting is assured).

15. Semaphore
A semaphore S is a global variable that is accessed by only process Pi at a time.
It is an integer variable that, apart from initialization, is accessed only through two atomic operations: wait and signal.
Two standard operations modify S: wait () and signal ()
Originally called P() and V()
wait (s): while S ≤ 0 do noop;
S = S – 1;
signal (s): S = S + 1;
wait and signal are executed indivisibly (while one process is modifying the semaphore, no other process can simultaneously modify it.)

16. Semaphore Usage – Mutual Exclusion
Semaphores are used to deal with the n-process critical section problem.
n processes can share a semaphore, mutex, initialized to 1.
Each Pi executes as follows:
do {
wait (mutex);
Critical section
signal (mutex);
Remainder section
} while (TRUE);

17. Semaphore Usage - Synchronization
Semaphores can also be used to solve synchronization problems.
Say there are 2 concurrently executing processes: P1 with a statement S1 and P2 with S2. S2 needs to be executed only after S1.
Solution: Let P1 and P2 share a semaphore, synch, initialized to 0.
Process P1: S1
signal (synch);
Process P2: wait (synch);
S2;
Because synch is initialized to 0, P2 will execute S2 only after P1 has invoked signal (synch) after executing S1.

18. Disadvantage of these mutual exclusion solutions
These solutions all require busy waiting i.e. while a process is in its CS, all other processes trying to enter CS must loop continuously in the entry code (wait on mutex). Thus CPU cycles are wasted. This is also called spinlock (Pi “spins” waiting on mutex lock).
Advantage of spin lock is that no context switch is required. So this can be used where applications need to hold the lock for a short duration.
Note that applications may spend lots of time in critical sections and therefore this may not be a good solution.

19. Semaphore Implementation with no Busy waiting
When a process executes wait (s) and finds the semaphore to be non-positive, it blocks itself.
The block operation places a process into a waiting queue associated with semaphore and the state of the process is changed to the waiting state.
A blocked process is restarted when another process executes a signal operation. This is done by a wakeup operation.
Two operations:
block – place the process invoking the operation on the appropriate semaphore waiting queue.
wakeup – remove one of processes in the waiting queue and place it in the ready queue.

20. Semaphore Implementation with no Busy waiting (Cont.)
Semaphore definition:
An integer value (initialized to 1) and a list of processes
Implementation of wait():
Implementation of signal():
Note: If the semaphore value is negative, its magnitude is the # of processes waiting on that semaphore (because of switching of order of decrement and test in wait(s)).
In the classical definition of wait(s), semaphore is never negative.

21. Deadlocks
The implementation of a semaphore with a waiting queue can result in a deadlock.
Let S and Q be two semaphores initialized to 1
P0 P1
wait (S); wait (Q);
wait (Q); wait (S);
: :
signal (S); signal (Q);
signal (Q); signal (S);

22. Classical Problems of Synchronization
Bounded-Buffer Problem
Readers and Writers Problem
Dining-Philosophers Problem

23. Bounded-Buffer Problem using Semaphores
N buffers, each can hold one item
Semaphore mutex initialized to the value 1 to provide mutual exclusion for access to the buffer pool.
Semaphore full initialized to the value 0 to count the # of full buffers.
Semaphore empty initialized to the value N to count the # of empty buffers.

24. Bounded-Buffer Problem using Semaphores
Producer process:
do {
produce an item in nextp;
wait (empty);
wait (mutex);
Add nextp to buffer;
signal (mutex);
signal (full);
} while (TRUE);

25. Bounded-Buffer Problem using Semaphores
Consumer process:
do {
wait (full);
wait (mutex);
Remove an item from buffer to nextc;
signal (mutex);
signal (empty);
consume the item nextc;
} while (TRUE);

26. Readers-Writers Problem
A data set (file or record or variable) is shared among a number of concurrent processes
Readers – only read the data set; they do not perform any updates
Writers – can both read and write
Problem – allow multiple readers to read at the same time. Mutual exclusion is needed to prevent a writer and some other process from simultaneously accessing the shared object (e.g. airline reservation record). This synchronization is called the readers and writers problem.
No reader should be kept waiting unless a writer has already obtained permission to use the shared object (could lead to writer starvation.)

27. Readers-Writers Problem
Shared Data
Semaphore mutex initialized to 1 (ensures mutual exclusion when the variable readCount is updated)
Semaphore wrt initialized to 1 (common to both readers and writers. It is used to provide mutual exclusion to writers. It is also used by the first or last reader that enters or exits the CS)
Integer readCount initialized to 0 (the # of readers currently reading the object)

28. Readers-Writers Problem
Writer process:
do {
wait (wrt); ….
writing is performed;
signal (wrt);
} while (TRUE);

29. Readers-Writers Problem
Reader process:
do {
wait (mutex);
readCount++; // # of readers
if (readCount == 1)
wait (wrt);
signal (mutex); …
reading is performed;
readCount--;
if (readCount == 0)
signal (wrt);
} while (TRUE);
Once one reader gets in by wait(wrt), then wrt is 0 and other readers can get in but no writer. When last reader gets out the wrt is incremented to 1. So now a writer can get in. But once 1 reader gets in , readers can keep getting in but not a writer. The only way a writer can get in is if no reader is reading. So writers can starve. It can keep on waiting.

30. Dining-Philosophers Problem
Problem represents the need to allocate several resources among several processes in a deadlock and starvation free manner.
Solution: Represent each chopstick by a semaphore. A philosopher tries to grab a chopstick by executing a wait() on that semaphore. A chopstick is released by a signal operation.
Semaphore chopStick [5] initialized to 1

31. Dining-Philosophers Problem (Cont.)
The structure of Philosopher i:
do {
wait (chopstick[i]);
wait (chopstick[i+1 mod 5]); …
eat
signal(chopstick[i]);
signal (chopstick[i+1 mod 5]);
think;
} while (TRUE);
Disadvantage: while this solution guarantees that no two neighbors are eating simultaneously, it might create a deadlock if say all 5 philosophers grab their left chopsticks. Now no philosopher can get the chopstick to the right.

32. Java Synchronization
Java provides synchronization at the language-level.
Each Java object has an associated lock.
This lock is acquired by invoking a synchronized method.
This lock is released when exiting the synchronized method.
Threads waiting to acquire the object lock are placed in the entry set for the object lock.

33. Java Synchronization
Each object has an associated entry set.

34. Java Synchronization
Synchronized insert() and remove() methods – Incorrect!

35. Java Synchronization wait/notify()
When a thread invokes wait(): 1. The thread releases the object lock; 2. The state of the thread is set to Blocked; 3. The thread is placed in the wait set for the object.
When a thread invokes notify(): 1. An arbitrary thread T from the wait set is selected; 2. T is moved from the wait to the entry set; 3. The state of T is set to Runnable.

36. Java Synchronization
Entry and wait sets

37. Java Synchronization – wait/notify
Synchronized insert() method – Correct!

38. Java Synchronization – wait/notify
Synchronized remove() method – Correct!

39. Java Synchronization - Bounded Buffer

40. Java Synchronization
The call to notify() selects an aribitrary thread from the wait set. It is possible the selected thread is in fact not waiting upon the condition for which it was notified.
The call notifyAll() selects all threads in the wait set and moves them to the entry set.
In general, notifyAll() is a more conservative strategy than notify().

41. Java Synchronization
notify() may not notify the correct thread!

42. Java Synchronization - Readers-Writers

43. Java Synchronization - Readers-Writers
Methods called by readers

44. Java Synchronization - Readers-Writers
Methods called by writers

45. Java Synchronization Rather than synchronizing an entire method, Block
synchronization allows blocks of code to be declared as
synchronized

46. Java Synchronization
Block synchronization using wait()/notify()

47. Concurrency Features in Java 5
Prior to Java 5, the only concurrency features in Java were Using synchronized/wait/notify.
Beginning with Java 5, new features were added to the API:
Reentrant Locks
Semaphores
Condition Variables

48. Concurrency Features in Java 5
Reentrant Locks

49. Concurrency Features in Java 5
Semaphores

50. Concurrency Features in Java 5
A condition variable is created by first creating a
ReentrantLock and invoking its newCondition() method:
Once this is done, it is possible to invoke the
await() and signal() methods

51. Concurrency Features in Java 5
doWork() method with condition variables

52. Synchronization Examples
Solaris
Windows XP
Linux
Pthreads

53. Solaris Synchronization
Implements a variety of locks to support multitasking, multithreading (including real-time threads), and multiprocessing
Uses adaptive mutexes for efficiency when protecting data from short code segments
Uses condition variables and readers-writers locks when longer sections of code need access to data
Uses turnstiles to order the list of threads waiting to acquire either an adaptive mutex or reader-writer lock

54. Windows XP Synchronization
Uses interrupt masks to protect access to global resources on uniprocessor systems
Uses spinlocks on multiprocessor systems
Also provides dispatcher objects which may act as either mutexes and semaphores
Dispatcher objects may also provide events
An event acts much like a condition variable

55. Linux Synchronization
Prior to kernel Version 2.6, disables interrupts to implement short critical sections
Version 2.6 and later, fully preemptive
Linux provides:
semaphores
spin locks

56. Pthreads Synchronization
Pthreads API is OS-independent
It provides:
mutex locks
condition variables
Non-portable extensions include:
read-write locks
spin locks

57. Atomic Transactions System Model Log-based Recovery Checkpoints
Concurrent Atomic Transactions

58. Transactional Memory
Memory transaction is a series of read-write operations that are atomic.
We replace
With
The atomic{S} statement ensures the statements in S execute as a transaction.

59. System Model
Assures that operations happen as a single logical unit of work, in its entirety, or not at all
Related to field of database systems
Challenge is assuring atomicity despite computer system failures
Transaction - collection of instructions or operations that performs single logical function
Here we are concerned with changes to stable storage – disk
Transaction is series of read and write operations
Terminated by commit (transaction successful) or abort (transaction failed) operation
Aborted transaction must be rolled back to undo any changes it performed

60. Types of Storage Media
Volatile storage – information stored here does not survive system crashes
Example: main memory, cache
Nonvolatile storage – Information usually survives crashes
Example: disk and tape
Stable storage – Information never lost
Not actually possible, so approximated via replication or RAID to devices with independent failure modes
Goal is to assure transaction atomicity where failures cause loss of information on volatile storage

61. Log-Based Recovery
Record to stable storage information about all modifications by a transaction
Most common is write-ahead logging
Log on stable storage, each log record describes single transaction write operation, including
Transaction name
Data item name
Old value
New value
<Ti starts> written to log when transaction Ti starts
<Ti commits> written when Ti commits
Log entry must reach stable storage before operation on data occurs

62. Log-Based Recovery Algorithm
Using the log, system can handle any volatile memory errors
Undo(Ti) restores value of all data updated by Ti
Redo(Ti) sets values of all data in transaction Ti to new values
Undo(Ti) and redo(Ti) must be idempotent
Multiple executions must have the same result as one execution
If system fails, restore state of all updated data via log
If log contains <Ti starts> without <Ti commits>, undo(Ti)
If log contains <Ti starts> and <Ti commits>, redo(Ti)

63. Checkpoints Log could become long, and recovery could take long
Checkpoints shorten log and recovery time.
Checkpoint scheme:
Output all log records currently in volatile storage to stable storage
Output all modified data from volatile to stable storage
Output a log record <checkpoint> to the log on stable storage
Now recovery only includes Ti, such that Ti started executing before the most recent checkpoint, and all transactions after Ti All other transactions already on stable storage

64. Concurrent Transactions
Must be equivalent to serial execution – serializability
Could perform all transactions in critical section
Inefficient, too restrictive
Concurrency-control algorithms provide serializability

65. Serializability Consider two data items A and B
Consider Transactions T0 and T1
Execute T0, T1 atomically
Execution sequence called schedule
Atomically executed transaction order called serial schedule
For N transactions, there are N! valid serial schedules

66. Schedule 1: T0 then T1

67. Nonserial Schedule Nonserial schedule allows overlapped execute
Resulting execution not necessarily incorrect
Consider schedule S, operations Oi, Oj
Conflict if access same data item, with at least one write
If Oi, Oj consecutive and operations of different transactions & Oi and Oj don’t conflict
Then S’ with swapped order Oj Oi equivalent to S
If S can become S’ via swapping nonconflicting operations
S is conflict serializable

68. Schedule 2: Concurrent Serializable Schedule

69. Locking Protocol
Ensure serializability by associating lock with each data item
Follow locking protocol for access control
Locks
Shared – Ti has shared-mode lock (S) on item Q, Ti can read Q but not write Q
Exclusive – Ti has exclusive-mode lock (X) on Q, Ti can read and write Q
Require every transaction on item Q acquire appropriate lock
If lock already held, new request may have to wait
Similar to readers-writers algorithm

70. Two-phase Locking Protocol
Generally ensures conflict serializability
Each transaction issues lock and unlock requests in two phases
Growing – obtaining locks
Shrinking – releasing locks
Does not prevent deadlock

71. Timestamp-based Protocols
Select order among transactions in advance – timestamp-ordering
Transaction Ti associated with timestamp TS(Ti) before Ti starts
TS(Ti) < TS(Tj) if Ti entered system before Tj
TS can be generated from system clock or as logical counter incremented at each entry of transaction
Timestamps determine serializability order
If TS(Ti) < TS(Tj), system must ensure produced schedule equivalent to serial schedule where Ti appears before Tj

72. Timestamp-based Protocol Implementation
Data item Q gets two timestamps
W-timestamp(Q) – largest timestamp of any transaction that executed write(Q) successfully
R-timestamp(Q) – largest timestamp of successful read(Q)
Updated whenever read(Q) or write(Q) executed
Timestamp-ordering protocol assures any conflicting read and write executed in timestamp order
Suppose Ti executes read(Q)
If TS(Ti) < W-timestamp(Q), Ti needs to read value of Q that was already overwritten
read operation rejected and Ti rolled back
If TS(Ti) ≥ W-timestamp(Q)
read executed, R-timestamp(Q) set to max(R-timestamp(Q), TS(Ti))

73. Timestamp-ordering Protocol
Suppose Ti executes write(Q)
If TS(Ti) < R-timestamp(Q), value Q produced by Ti was needed previously and Ti assumed it would never be produced
Write operation rejected, Ti rolled back
If TS(Ti) < W-tiimestamp(Q), Ti attempting to write obsolete value of Q
Write operation rejected and Ti rolled back
Otherwise, write executed
Any rolled back transaction Ti is assigned new timestamp and restarted
Algorithm ensures conflict serializability and freedom from deadlock

74. Schedule Possible Under Timestamp Protocol

75. End of Chapter 6