Managing Hierarchies of Database Elements
(18.6)
Presented by
Sarat Dasika
(114)
February 16, 2012
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Managing Hierarchies of Database Elements
Locks with Multiple Granularity
Warning (Intention) Locks
Database Elements Organized in Hierarchy
Rules of Warning Protocol
Group Modes of Intention Locks
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Two problems that arise with locks when there is a tree structure to the data are:
When the tree structure is a hierarchy of lockable elements
◦ Determine how locks are granted for both large elements (relations) and smaller elements (blocks containing tuples or individual tuples)
When the data itself is organized as a tree
(B-tree indexes)
◦ This will be discussed in the next section

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A database element can be a relation, block or a tuple
Different systems use different database elements to determine the size of the lock
Thus some may require small database elements such as tuples or blocks and others may require large elements such as relations

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Consider a database for a bank
◦ Choosing relations as database elements means we would have one lock for an entire relation
◦ If we were dealing with a relation having account balances, this kind of lock would be very inflexible and thus provide very little concurrency
◦ Why? Because balance transactions require exclusive locks and this would mean only one transaction occurs for one account at any time
◦ But as each account is independent of others we could perform transactions on different accounts simultaneously

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◦ Thus it makes sense to have block element for the lock so that two accounts on different blocks can be updated simultaneously
Another example is that of a document
◦ With similar arguments as above, we see that it is better to have large element (a complete document) as the lock in this case

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These are required to manage locks at different granularities
◦ In the bank example, if the a shared lock is obtained for the relation while there are exclusive locks on individual tuples, unserializable behavior occurs
The rules for managing locks on hierarchy of database elements constitute the

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These involve both
(S and X) and
(IS and IX) locks
The rules are:
◦ Begin at the root of hierarchy
◦ Request the S/X lock if we are at the desired element
◦ If the desired element id further down the hierarchy, place a warning lock (IS if S and IX if X)
◦ When the warning lock is granted, we proceed to the child node and repeat the above steps until desired node is reached

IS
IX
S
X
IS IX
Yes Yes Yes No
Yes Yes No
Yes No
No No
S
Yes
No
X
No
No
No
• The above matrix applies only to locks held by other transactions

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An element can request S and IX locks at the same time if they are in the same transaction
(to read entire element and then modify sub elements)
This can be considered as another lock mode,
SIX, having restrictions of both the locks i.e.
No for all except IS
SIX serves as the group mode

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Consider a transaction T
◦ Select * from
( as follows relation; then moving to individual tuples them table
◦ Here, IS lock is first acquired on the entire attribute = ‘abc’ where
1 attribute1 = ‘abc’
), S lock in acquired on each of
Consider another transaction T
◦ Update
2 table set attribute2 = ‘def’ where attribute1 = ‘ghi’
◦ Here, it requires an IX lock on relation and since
T
1
’s IS lock is compatible, IX is granted

◦ On reaching the desired tuple (ghi), as there is no lock, it gets X too
◦ If T2 was updating the same tuple as T1, it would have to wait until T1 released its S lock

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This arises when transactions create new sub elements of lockable elements
Since we can lock only existing elements the new elements fail to be locked
The problem created in this situation is explained in the following example

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Consider a transaction T
◦ Select sum(length) from
‘abc’
3 table where attribute1 =
◦ This calculates the total length of all tuples having attribute1
◦ Thus, T
3 tuples acquires IS for relation and S for targeted

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This is not a concurrency problem since the serial order (T3, T4) is maintained
But if both T3 and T4 were to write an element X, it could lead to unserializable behavior
◦ r3(t1);r3(t2);w4(t3);w4(X);w3(L);w3(X)
◦ r3 and w3 are read and write operations by T3 and w4 are the write operations by T4 and L is the total length calculated by T3 (t1 + t2)
◦ At the end, we have result of T3 as sum of lengths of t1 and t2 and X has value written by T3
◦ This is not right; if value of X is considered to be that written by T3 then for the schedule to be serializable, the sum of lengths of t1, t2 and t3 should be considered

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◦ Else if the sum is total length of t1 and t2 then for the schedule to be serializable, X should have value written by T4
This problem arises since the relation has a phantom tuple (the new inserted tuple), which should have been locked but wasn’t since it didn’t exist at the time locks were taken
The occurrence of phantoms can be avoided if all insertion and deletion transactions are treated as write operations on the whole relation