concurrency control, protocols, deadlocks.

Database System Concepts, 6th
Ed.
©Silberschatz, Korth and Sudarshan
See www.db-book.com for conditions on re-use
Chapter 15 : Concurrency Control
Chapter 15 : Concurrency Control

15.2
Database System Concepts - 6th
Edition
Outline
Outline
 Lock-Based Protocols
 Timestamp-Based Protocols
 Validation-Based Protocols
 Multiple Granularity
 Multiversion Schemes
 Insert and Delete Operations
 Concurrency in Index Structures

15.3
Edition
Lock-Based Protocols
Lock-Based Protocols
 A lock is a mechanism to control concurrent access to a data
item
 Data items can be locked in two modes :
1. exclusive (X) mode. Data item can be both read as well as
written. X-lock is requested using lock-X instruction.
2. shared (S) mode. Data item can only be read. S-lock is
requested using lock-S instruction.
 Lock requests are made to the concurrency-control manager
by the programmer. Transaction can proceed only after
request is granted.

15.4
Edition
Lock-Based Protocols (Cont.)
 Lock-compatibility matrix
 A transaction may be granted a lock on an item if the requested
lock is compatible with locks already held on the item by other
transactions
 Any number of transactions can hold shared locks on an item,
 But if any transaction holds an exclusive on the item no other
transaction may hold any lock on the item.
 If a lock cannot be granted, the requesting transaction is made to
wait till all incompatible locks held by other transactions have
been released. The lock is then granted.

15.5
Edition
 Example of a transaction performing locking:
T2: lock-S(A);
read (A);
unlock(A);
lock-S(B);
read (B);
unlock(B);
display(A+B)
 Locking as above is not sufficient to guarantee serializability
— if A and B get updated in-between the read of A and B, the
displayed sum would be wrong.
 A locking protocol is a set of rules followed by all
transactions while requesting and releasing locks. Locking
protocols restrict the set of possible schedules.

15.6
Edition
The Two-Phase Locking Protocol
The Two-Phase Locking Protocol
 This protocol ensures conflict-serializable schedules.
 Phase 1: Growing Phase
 Transaction may obtain locks
 Transaction may not release locks
 Phase 2: Shrinking Phase
 Transaction may release locks
 Transaction may not obtain locks
 The protocol assures serializability. It can be proved that the
transactions can be serialized in the order of their lock points
(i.e., the point where a transaction acquired its final lock).

15.7
Edition
The Two-Phase Locking Protocol (Cont.)
The Two-Phase Locking Protocol (Cont.)
 There can be conflict serializable schedules that cannot be
obtained if two-phase locking is used.
 However, in the absence of extra information (e.g., ordering of
access to data), two-phase locking is needed for conflict
serializability in the following sense:
 Given a transaction Ti that does not follow two-phase
locking, we can find a transaction Tj that uses two-phase
locking, and a schedule for Ti and Tj that is not conflict
serializable.

15.8
Edition
Lock Conversions
Lock Conversions
 Two-phase locking with lock conversions:
– First Phase:
 can acquire a lock-S on item
 can acquire a lock-X on item
 can convert a lock-S to a lock-X (upgrade)
– Second Phase:
 can release a lock-S
 can release a lock-X
 can convert a lock-X to a lock-S (downgrade)
 This protocol assures serializability. But still relies on the
programmer to insert the various locking instructions.

15.9
Edition
Automatic Acquisition of Locks
Automatic Acquisition of Locks
 A transaction Ti issues the standard read/write instruction,
without explicit locking calls.
 The operation read(D) is processed as:
if Ti has a lock on D
then
read(D)
else begin
if necessary wait until no other
transaction has a lock-X on D
grant Ti a lock-S on D;
read(D)
end

15.10
Edition
Automatic Acquisition of Locks (Cont.)
Automatic Acquisition of Locks (Cont.)
 write(D) is processed as:
if Ti has a lock-X on D
then
write(D)
else begin
if necessary wait until no other transaction has any lock on D,
if Ti has a lock-S on D
then
upgrade lock on D to lock-X
else
grant Ti a lock-X on D
write(D)
end;
 All locks are released after commit or abort

15.11
Edition
Deadlocks
Deadlocks
 Consider the partial schedule
 Neither T3 nor T4 can make progress — executing lock-S(B) causes
T4 to wait for T3 to release its lock on B, while executing lock-X(A)
causes T3 to wait for T4 to release its lock on A.
 Such a situation is called a deadlock.
 To handle a deadlock one of T3 or T4 must be rolled back
and its locks released.

15.12
Edition
Deadlocks (Cont.)
Deadlocks (Cont.)
 Two-phase locking does not ensure freedom from deadlocks.
 In addition to deadlocks, there is a possibility of starvation.
 Starvation occurs if the concurrency control manager is badly
designed. For example:
 A transaction may be waiting for an X-lock on an item,
while a sequence of other transactions request and are
granted an S-lock on the same item.
 The same transaction is repeatedly rolled back due to
deadlocks.
 Concurrency control manager can be designed to prevent
starvation.

15.13
Edition
Deadlocks (Cont.)
Deadlocks (Cont.)
 The potential for deadlock exists in most locking protocols.
Deadlocks are a necessary evil.
 When a deadlock occurs there is a possibility of cascading roll-
backs.
 Cascading roll-back is possible under two-phase locking. To
avoid this, follow a modified protocol called strict two-phase
locking -- a transaction must hold all its exclusive locks till it
commits/aborts.
 Rigorous two-phase locking is even stricter. Here, all locks
are held till commit/abort. In this protocol transactions can be
serialized in the order in which they commit.

15.14
Edition
Implementation of Locking
Implementation of Locking
 A lock manager can be implemented as a separate process to
which transactions send lock and unlock requests
 The lock manager replies to a lock request by sending a lock
grant messages (or a message asking the transaction to roll
back, in case of a deadlock)
 The requesting transaction waits until its request is answered
 The lock manager maintains a data-structure called a lock
table to record granted locks and pending requests
 The lock table is usually implemented as an in-memory hash
table indexed on the name of the data item being locked

15.15
Edition
Lock Table
Lock Table
 Dark blue rectangles indicate granted
locks; light blue indicate waiting requests
 Lock table also records the type of lock
granted or requested
 New request is added to the end of the
queue of requests for the data item, and
granted if it is compatible with all earlier
locks
 Unlock requests result in the request
being deleted, and later requests are
checked to see if they can now be
granted
 If transaction aborts, all waiting or granted
requests of the transaction are deleted
 lock manager may keep a list of locks
held by each transaction, to
implement this efficiently

15.16
Edition
Deadlock Handling
Deadlock Handling
 System is deadlocked if there is a set of transactions such that
every transaction in the set is waiting for another transaction in
the set.
 Deadlock prevention protocols ensure that the system will never
enter into a deadlock state. Some prevention strategies :
 Require that each transaction locks all its data items before it
begins execution (predeclaration).
 Impose partial ordering of all data items and require that a
transaction can lock data items only in the order specified by
the partial order.

15.17
Edition
More Deadlock Prevention Strategies
More Deadlock Prevention Strategies
 Following schemes use transaction timestamps for the sake of
deadlock prevention alone.
 wait-die scheme — non-preemptive
 older transaction may wait for younger one to release data item.
(older means smaller timestamp) Younger transactions never
Younger transactions never wait for older ones; they are rolled
back instead.
 a transaction may die several times before acquiring needed data
item
 wound-wait scheme — preemptive
 older transaction wounds (forces rollback) of younger transaction
instead of waiting for it. Younger transactions may wait for older
ones.
 may be fewer rollbacks than wait-die scheme.

15.18
Edition
Deadlock prevention (Cont.)
Deadlock prevention (Cont.)
 Both in wait-die and in wound-wait schemes, a rolled back
transactions is restarted with its original timestamp. Older transactions
thus have precedence over newer ones, and starvation is hence
avoided.
 Timeout-Based Schemes:
 a transaction waits for a lock only for a specified amount of time. If
the lock has not been granted within that time, the transaction is
rolled back and restarted,
 Thus, deadlocks are not possible
 simple to implement; but starvation is possible. Also difficult to
determine good value of the timeout interval.

15.19
Edition
Deadlock Detection
Deadlock Detection
 Deadlocks can be described as a wait-for graph, which consists of a
pair G = (V,E),
 V is a set of vertices (all the transactions in the system)
 E is a set of edges; each element is an ordered pair Ti Tj.
 If Ti  Tj is in E, then there is a directed edge from Ti to Tj, implying
that Ti is waiting for Tj to release a data item.
 When Ti requests a data item currently being held by Tj, then the edge
Ti  Tj is inserted in the wait-for graph. This edge is removed only
when Tj is no longer holding a data item needed by Ti.
 The system is in a deadlock state if and only if the wait-for graph has a
cycle. Must invoke a deadlock-detection algorithm periodically to look
for cycles.

15.20
Edition
Deadlock Detection (Cont.)
Deadlock Detection (Cont.)
Wait-for graph without a cycle Wait-for graph with a cycle

15.21
Edition
Deadlock Recovery
Deadlock Recovery
 When deadlock is detected :
 Some transaction will have to rolled back (made a victim) to
break deadlock. Select that transaction as victim that will incur
minimum cost.
 Rollback -- determine how far to roll back transaction
 Total rollback: Abort the transaction and then restart it.
 More effective to roll back transaction only as far as
necessary to break deadlock.
 Starvation happens if same transaction is always chosen as
victim. Include the number of rollbacks in the cost factor to avoid
starvation

15.22
Edition
Multiple Granularity
Multiple Granularity
 Allow data items to be of various sizes and define a hierarchy of data
granularities, where the small granularities are nested within larger
ones
 Can be represented graphically as a tree.
 When a transaction locks a node in the tree explicitly, it implicitly locks
all the node's descendents in the same mode.
 Granularity of locking (level in tree where locking is done):
 fine granularity (lower in tree): high concurrency, high locking
overhead
 coarse granularity (higher in tree): low locking overhead, low
concurrency

15.23
Edition
Example of Granularity Hierarchy
Example of Granularity Hierarchy
The levels, starting from the coarsest (top) level are
 database
 area
 file
 record

15.24
Edition
Intention Lock Modes
Intention Lock Modes
 In addition to S and X lock modes, there are three additional lock
modes with multiple granularity:
 intention-shared (IS): indicates explicit locking at a lower level of
the tree but only with shared locks.
 intention-exclusive (IX): indicates explicit locking at a lower level
with exclusive or shared locks
 shared and intention-exclusive (SIX): the subtree rooted by that
node is locked explicitly in shared mode and explicit locking is
being done at a lower level with exclusive-mode locks.
 intention locks allow a higher level node to be locked in S or X mode
without having to check all descendent nodes.

15.25
Edition
Compatibility Matrix with Intention Lock Modes
Compatibility Matrix with Intention Lock Modes
 The compatibility matrix for all lock modes is:

15.26
Edition
Multiple Granularity Locking Scheme
Multiple Granularity Locking Scheme
 Transaction Ti can lock a node Q, using the following rules:
1. The lock compatibility matrix must be observed.
2. The root of the tree must be locked first, and may be locked in any mode.
3. A node Q can be locked by Ti in S or IS mode only if the parent of Q is
currently locked by Ti in either IX or IS mode.
4. A node Q can be locked by Ti in X, SIX, or IX mode only if the parent of Q
is currently locked by Ti in either IX or SIX mode.
5. Ti can lock a node only if it has not previously unlocked any node (that is,
Ti is two-phase).
6. Ti can unlock a node Q only if none of the children of Q are currently
locked by Ti.
 Observe that locks are acquired in root-to-leaf order, whereas they are
released in leaf-to-root order.
 Lock granularity escalation: in case there are too many locks at a particular
level, switch to higher granularity S or X lock

15.27
Edition
Timestamp-Based Protocols
Timestamp-Based Protocols
 Each transaction is issued a timestamp when it enters the system. If
an old transaction Ti has time-stamp TS(Ti), a new transaction Tj is
assigned time-stamp TS(Tj) such that TS(Ti) <TS(Tj).
 The protocol manages concurrent execution such that the time-stamps
determine the serializability order.
 In order to assure such behavior, the protocol maintains for each data
Q two timestamp values:
 W-timestamp(Q) is the largest time-stamp of any transaction that
executed write(Q) successfully.
 R-timestamp(Q) is the largest time-stamp of any transaction that
executed read(Q) successfully.

15.28
Edition
Timestamp-Based Protocols (Cont.)
 The timestamp ordering protocol ensures that any conflicting read
and write operations are executed in timestamp order.
 Suppose a transaction Ti issues a read(Q)
1. If TS(Ti)  W-timestamp(Q), then Ti needs to read a value of Q
that was already overwritten.
 Hence, the read operation is rejected, and Ti is rolled back.
2. If TS(Ti)  W-timestamp(Q), then the read operation is
executed, and R-timestamp(Q) is set to max(R-timestamp(Q),
TS(Ti)).

15.29
Edition
 Suppose that transaction Ti issues write(Q).
1. If TS(Ti) < R-timestamp(Q), then the value of Q that Ti is
producing was needed previously, and the system assumed that
that value would never be produced.
 Hence, the write operation is rejected, and Ti is rolled back.
2. If TS(Ti) < W-timestamp(Q), then Ti is attempting to write an
obsolete value of Q.
 Hence, this write operation is rejected, and Ti is rolled back.
3. Otherwise, the write operation is executed, and W-timestamp(Q)
is set to TS(Ti).

15.30
Edition
Example Use of the Protocol
Example Use of the Protocol
A partial schedule for several data items for transactions with
timestamps 1, 2, 3, 4, 5

15.31
Edition
Correctness of Timestamp-Ordering Protocol
Correctness of Timestamp-Ordering Protocol
 The timestamp-ordering protocol guarantees serializability since all
the arcs in the precedence graph are of the form:
Thus, there will be no cycles in the precedence graph
 Timestamp protocol ensures freedom from deadlock as no
transaction ever waits.
 But the schedule may not be cascade-free, and may not even be
recoverable.

15.32
Edition
Recoverability and Cascade Freedom
Recoverability and Cascade Freedom
 Problem with timestamp-ordering protocol:
 Suppose Ti aborts, but Tj has read a data item written by Ti
 Then Tj must abort; if Tj had been allowed to commit earlier, the
schedule is not recoverable.
 Further, any transaction that has read a data item written by Tj
must abort
 This can lead to cascading rollback --- that is, a chain of rollbacks
 Solution 1:
 A transaction is structured such that its writes are all performed at
the end of its processing
 All writes of a transaction form an atomic action; no transaction
may execute while a transaction is being written
 A transaction that aborts is restarted with a new timestamp
 Solution 2: Limited form of locking: wait for data to be committed
before reading it
 Solution 3: Use commit dependencies to ensure recoverability

15.33
Edition
Thomas
Thomas’
’ Write Rule
Write Rule
 Modified version of the timestamp-ordering protocol in which obsolete
write operations may be ignored under certain circumstances.
 When Ti attempts to write data item Q, if TS(Ti) < W-timestamp(Q),
then Ti is attempting to write an obsolete value of {Q}.
 Rather than rolling back Ti as the timestamp ordering protocol
would have done, this {write} operation can be ignored.
 Otherwise this protocol is the same as the timestamp ordering
protocol.
 Thomas' Write Rule allows greater potential concurrency.
 Allows some view-serializable schedules that are not conflict-
serializable.

15.34
Edition
Validation-Based Protocol
Validation-Based Protocol
 Execution of transaction Ti is done in three phases.
1. Read and execution phase: Transaction Ti writes only to
temporary local variables
2. Validation phase: Transaction Ti performs a ''validation test''
to determine if local variables can be written without violating
serializability.
3. Write phase: If Ti is validated, the updates are applied to the
database; otherwise, Ti is rolled back.
 The three phases of concurrently executing transactions can be
interleaved, but each transaction must go through the three phases in that
order.
 Assume for simplicity that the validation and write phase occur
together, atomically and serially
 I.e., only one transaction executes validation/write at a time.
 Also called as optimistic concurrency control since transaction executes
fully in the hope that all will go well during validation

15.35
Edition
Validation-Based Protocol (Cont.)
Validation-Based Protocol (Cont.)
 Each transaction Ti has 3 timestamps
 Start(Ti) : the time when Ti started its execution
 Validation(Ti): the time when Ti entered its validation phase
 Finish(Ti) : the time when Ti finished its write phase
 Serializability order is determined by timestamp given at validation time;
this is done to increase concurrency.
 Thus, TS(Ti) is given the value of Validation(Ti).
 This protocol is useful and gives greater degree of concurrency if
probability of conflicts is low.
 because the serializability order is not pre-decided, and
 relatively few transactions will have to be rolled back.

15.36
Edition
Validation Test for Transaction
Validation Test for Transaction T
Tj
j
 If for all Ti with TS (Ti) < TS (Tj) either one of the following condition
holds:
 finish(Ti) < start(Tj)
 start(Tj) < finish(Ti) < validation(Tj) and the set of data items
written by Ti does not intersect with the set of data items read
by Tj.
then validation succeeds and Tj can be committed. Otherwise,
validation fails and Tj is aborted.
 Justification: Either the first condition is satisfied, and there is no
overlapped execution, or the second condition is satisfied and
 the writes of Tj do not affect reads of Ti since they occur after Ti
has finished its reads.
 the writes of Ti do not affect reads of Tj since Tj does not read
any item written by Ti.

15.37
Edition
Schedule Produced by Validation
Schedule Produced by Validation
 Example of schedule produced using validation

15.38
Edition
Multiversion Schemes
Multiversion Schemes
 Multiversion schemes keep old versions of data item to increase
concurrency.
 Multiversion Timestamp Ordering
 Multiversion Two-Phase Locking
 Each successful write results in the creation of a new version of the
data item written.
 Use timestamps to label versions.
 When a read(Q) operation is issued, select an appropriate version of
Q based on the timestamp of the transaction, and return the value of
the selected version.
 reads never have to wait as an appropriate version is returned
immediately.

15.39
Edition
Multiversion Timestamp Ordering
Multiversion Timestamp Ordering
 Each data item Q has a sequence of versions <Q1, Q2,...., Qm>. Each
version Qk contains three data fields:
 Content -- the value of version Qk.
 W-timestamp(Qk) -- timestamp of the transaction that created
(wrote) version Qk
 R-timestamp(Qk) -- largest timestamp of a transaction that
successfully read version Qk
 When a transaction Ti creates a new version Qk of Q, Qk's W-
timestamp and R-timestamp are initialized to TS(Ti).
 R-timestamp of Qk is updated whenever a transaction Tj reads Qk, and
TS(Tj) > R-timestamp(Qk).

15.40
Edition
Multiversion Timestamp Ordering (Cont)
Multiversion Timestamp Ordering (Cont)
 Suppose that transaction Ti issues a read(Q) or write(Q) operation. Let Qk
denote the version of Q whose write timestamp is the largest write timestamp
less than or equal to TS(Ti).
1. If transaction Ti issues a read(Q), then the value returned is the content
of version Qk.
2. If transaction Ti issues a write(Q)
1. if TS(Ti) < R-timestamp(Qk), then transaction Ti is rolled back.
2. if TS(Ti) = W-timestamp(Qk), the contents of Qk are overwritten
3. else a new version of Q is created.
 Observe that
 Reads always succeed
 A write by Ti is rejected if some other transaction Tj that (in the serialization
order defined by the timestamp values) should read
Ti's write, has already read a version created by a transaction older than Ti.
 Protocol guarantees serializability

15.41
Edition
Multiversion Two-Phase Locking
Multiversion Two-Phase Locking
 Differentiates between read-only transactions and update transactions
 Update transactions acquire read and write locks, and hold all locks up
to the end of the transaction. That is, update transactions follow rigorous
two-phase locking.
 Each successful write results in the creation of a new version of the
data item written.
 Each version of a data item has a single timestamp whose value is
obtained from a counter ts-counter that is incremented during
commit processing.
 Read-only transactions are assigned a timestamp by reading the current
value of ts-counter before they start execution; they follow the
multiversion timestamp-ordering protocol for performing reads.

15.42
Edition
Multiversion Two-Phase Locking (Cont.)
Multiversion Two-Phase Locking (Cont.)
 When an update transaction wants to read a data item:
 it obtains a shared lock on it, and reads the latest version.
 When it wants to write an item
 it obtains X lock on; it then creates a new version of the item and
sets this version's timestamp to .
 When update transaction Ti completes, commit processing occurs:
 Ti sets timestamp on the versions it has created to ts-counter + 1
 Ti increments ts-counter by 1
 Read-only transactions that start after Ti increments ts-counter will see
the values updated by Ti.
 Read-only transactions that start before Ti increments the
ts-counter will see the value before the updates by Ti.
 Only serializable schedules are produced.

15.43
Edition
MVCC: Implementation Issues
MVCC: Implementation Issues
 Creation of multiple versions increases storage overhead
 Extra tuples
 Extra space in each tuple for storing version information
 Versions can, however, be garbage collected
 E.g. if Q has two versions Q5 and Q9, and the oldest active
transaction has timestamp > 9, than Q5 will never be required
again

15.44
Edition
Snapshot Isolation
Snapshot Isolation
 Motivation: Decision support queries that read large amounts of data
have concurrency conflicts with OLTP transactions that update a few
rows
 Poor performance results
 Solution 1: Give logical “snapshot” of database state to read only
transactions, read-write transactions use normal locking
 Multiversion 2-phase locking
 Works well, but how does system know a transaction is read only?
 Solution 2: Give snapshot of database state to every transaction,
updates alone use 2-phase locking to guard against concurrent
updates
 Problem: variety of anomalies such as lost update can result
 Partial solution: snapshot isolation level (next slide)
 Proposed by Berenson et al, SIGMOD 1995
 Variants implemented in many database systems
– E.g. Oracle, PostgreSQL, SQL Server 2005

15.45
Edition
Snapshot Isolation
Snapshot Isolation
 A transaction T1 executing with Snapshot
Isolation
 takes snapshot of committed data at
start
 always reads/modifies data in its own
snapshot
 updates of concurrent transactions are
not visible to T1
 writes of T1 complete when it commits
 First-committer-wins rule:
 Commits only if no other concurrent
transaction has already written data
that T1 intends to write.
T1 T2 T3
W(Y := 1)
Commit
Start
R(X)  0
R(Y) 1
W(X:=2)
W(Z:=3)
Commit
R(Z)  0
R(Y)  1
W(X:=3)
Commit-Req
Abort
Concurrent updates not visible
Own updates are visible
Not first-committer of X
Serialization error, T2 is rolled back

15.46
Edition
Snapshot Read
Snapshot Read
 Concurrent updates invisible to snapshot read

15.47
Edition
Snapshot Write:
Snapshot Write: First Committer Wins
 Variant: “First-updater-wins”
 Check for concurrent updates when write occurs by locking item
– But lock should be held till all concurrent transactions have finished
 (Oracle uses this plus some extra features)
 Differs only in when abort occurs, otherwise equivalent

15.48
Edition
Benefits of SI
Benefits of SI
 Reading is never blocked,
 and also doesn’t block other txns activities
 Performance similar to Read Committed
 Avoids the usual anomalies
 No dirty read
 No lost update
 No non-repeatable read
 Predicate based selects are repeatable (no phantoms)
 Problems with SI
 SI does not always give serializable executions
 Serializable: among two concurrent txns, one sees the effects
of the other
 In SI: neither sees the effects of the other
 Result: Integrity constraints can be violated

15.49
Edition
Snapshot Isolation
Snapshot Isolation
 E.g. of problem with SI
 T1: x:=y
 T2: y:= x
 Initially x = 3 and y = 17
 Serial execution: x = ??, y = ??
 if both transactions start at the same time, with snapshot
isolation: x = ?? , y = ??
 Called skew write
 Skew also occurs with inserts
 E.g:
 Find max order number among all orders
 Create a new order with order number = previous max + 1

15.50
Edition
Snapshot Isolation Anomalies
Snapshot Isolation Anomalies
 SI breaks serializability when txns modify different items, each based on a
previous state of the item the other modified
 Not very common in practice
 E.g., the TPC-C benchmark runs correctly under SI
 when txns conflict due to modifying different data, there is usually also
a shared item they both modify too (like a total quantity) so SI will abort
one of them
 But does occur
 Application developers should be careful about write skew
 SI can also cause a read-only transaction anomaly, where read-only
transaction may see an inconsistent state even if updaters are serializable
 We omit details
 Using snapshots to verify primary/foreign key integrity can lead to
inconsistency
 Integrity constraint checking usually done outside of snapshot

15.51
Edition
SI In Oracle and PostgreSQL
 Warning: SI used when isolation level is set to serializable, by Oracle, and
PostgreSQL versions prior to 9.1
 PostgreSQL’s implementation of SI (versions prior to 9.1) described in
Section 26.4.1.3
 Oracle implements “first updater wins” rule (variant of “first committer
wins”)
 concurrent writer check is done at time of write, not at commit time
 Allows transactions to be rolled back earlier
 Oracle and PostgreSQL < 9.1 do not support true serializable
execution
 PostgreSQL 9.1 introduced new protocol called “Serializable Snapshot
Isolation” (SSI)
 Which guarantees true serializabilty including handling predicate
reads (coming up)

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 Can sidestep SI for specific queries by using select .. for update in Oracle
and PostgreSQL
 E.g.,
1. select max(orderno) from orders for update
2. read value into local variable maxorder
3. insert into orders (maxorder+1, …)
 Select for update (SFU) treats all data read by the query as if it were
also updated, preventing concurrent updates
 Does not always ensure serializability since phantom phenomena can
occur (coming up)
 In PostgreSQL versions < 9.1, SFU locks the data item, but releases locks
when the transaction completes, even if other concurrent transactions are
active
 Not quite same as SFU in Oracle, which keeps locks until all
 concurrent transactions have completed

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Insert and Delete Operations
Insert and Delete Operations
 If two-phase locking is used :
 A delete operation may be performed only if the transaction
deleting the tuple has an exclusive lock on the tuple to be deleted.
 A transaction that inserts a new tuple into the database is given an
X-mode lock on the tuple
 Insertions and deletions can lead to the phantom phenomenon.
 A transaction that scans a relation
 (e.g., find sum of balances of all accounts in Perryridge)
and a transaction that inserts a tuple in the relation
 (e.g., insert a new account at Perryridge)
(conceptually) conflict in spite of not accessing any tuple in
common.
 If only tuple locks are used, non-serializable schedules can result
 E.g. the scan transaction does not see the new account, but
reads some other tuple written by the update transaction

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Insert and Delete Operations (Cont.)
Insert and Delete Operations (Cont.)
 The transaction scanning the relation is reading information that indicates
what tuples the relation contains, while a transaction inserting a tuple
updates the same information.
 The conflict should be detected, e.g. by locking the information.
 One solution:
 Associate a data item with the relation, to represent the information
about what tuples the relation contains.
 Transactions scanning the relation acquire a shared lock in the data
item,
 Transactions inserting or deleting a tuple acquire an exclusive lock on
the data item. (Note: locks on the data item do not conflict with locks on
individual tuples.)
 Above protocol provides very low concurrency for insertions/deletions.
 Index locking protocols provide higher concurrency while
preventing the phantom phenomenon, by requiring locks
on certain index buckets.

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Index Locking Protocol
Index Locking Protocol
 Index locking protocol:
 Every relation must have at least one index.
 A transaction can access tuples only after finding them through one or
more indices on the relation
 A transaction Ti that performs a lookup must lock all the index leaf
nodes that it accesses, in S-mode
 Even if the leaf node does not contain any tuple satisfying the index
lookup (e.g. for a range query, no tuple in a leaf is in the range)
 A transaction Ti that inserts, updates or deletes a tuple ti in a relation r
 must update all indices to r
 must obtain exclusive locks on all index leaf nodes affected by the
insert/update/delete
 The rules of the two-phase locking protocol must be observed
 Guarantees that phantom phenomenon won’t occur

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Next-Key Locking
Next-Key Locking
 Index-locking protocol to prevent phantoms required locking entire leaf
 Can result in poor concurrency if there are many inserts
 Alternative: for an index lookup
 Lock all values that satisfy index lookup (match lookup value, or
fall in lookup range)
 Also lock next key value in index
 Lock mode: S for lookups, X for insert/delete/update
 Ensures that range queries will conflict with inserts/deletes/updates
 Regardless of which happens first, as long as both are concurrent

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Concurrency in Index Structures
Concurrency in Index Structures
 Indices are unlike other database items in that their only job is to help in
accessing data.
 Index-structures are typically accessed very often, much more than
other database items.
 Treating index-structures like other database items, e.g. by 2-phase
locking of index nodes can lead to low concurrency.
 There are several index concurrency protocols where locks on internal
nodes are released early, and not in a two-phase fashion.
 It is acceptable to have nonserializable concurrent access to an
index as long as the accuracy of the index is maintained.
 In particular, the exact values read in an internal node of a
B+
-tree are irrelevant so long as we land up in the correct leaf
node.

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Concurrency in Index Structures (Cont.)
Concurrency in Index Structures (Cont.)
 Example of index concurrency protocol:
 Use crabbing instead of two-phase locking on the nodes of the B+
-tree, as
follows. During search/insertion/deletion:
 First lock the root node in shared mode.
 After locking all required children of a node in shared mode, release the lock
on the node.
 During insertion/deletion, upgrade leaf node locks to exclusive mode.
 When splitting or coalescing requires changes to a parent, lock the parent in
exclusive mode.
 Above protocol can cause excessive deadlocks
 Searches coming down the tree deadlock with updates going up the tree
 Can abort and restart search, without affecting transaction
 Better protocols are available; see Section 16.9 for one such protocol, the B-link
tree protocol
 Intuition: release lock on parent before acquiring lock on child
 And deal with changes that may have happened between lock release
and acquire

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Weak Levels of Consistency
Weak Levels of Consistency
 Degree-two consistency: differs from two-phase locking in that S-locks
may be released at any time, and locks may be acquired at any time
 X-locks must be held till end of transaction
 Serializability is not guaranteed, programmer must ensure that no
erroneous database state will occur]
 Cursor stability:
 For reads, each tuple is locked, read, and lock is immediately
released
 X-locks are held till end of transaction
 Special case of degree-two consistency

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Weak Levels of Consistency in SQL
Weak Levels of Consistency in SQL
 SQL allows non-serializable executions
 Serializable: is the default
 Repeatable read: allows only committed records to be read, and
repeating a read should return the same value (so read locks should
be retained)
 However, the phantom phenomenon need not be prevented
– T1 may see some records inserted by T2, but may not see
others inserted by T2
 Read committed: same as degree two consistency, but most
systems implement it as cursor-stability
 Read uncommitted: allows even uncommitted data to be read
 In many database systems, read committed is the default consistency
level
 has to be explicitly changed to serializable when required
 set isolation level serializable

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Transactions across User Interaction
 Many applications need transaction support across user interactions
 Can’t use locking
 Don’t want to reserve database connection per user
 Application level concurrency control
 Each tuple has a version number
 Transaction notes version number when reading tuple
 select r.balance, r.version into :A, :version
from r where acctId =23
 When writing tuple, check that current version number is same as the
version when tuple was read
 update r set r.balance = r.balance + :deposit
where acctId = 23 and r.version = :version
 Equivalent to optimistic concurrency control without validating read set
 Used internally in Hibernate ORM system, and manually in many applications
 Version numbering can also be used to support first committer wins check of
snapshot isolation
 Unlike SI, reads are not guaranteed to be from a single snapshot

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End of Module 16

15.63
Edition
Deadlocks
Deadlocks
 Consider the following two transactions:
T1: write (X) T2: write(Y)
write(Y) write(X)
 Schedule with deadlock

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