7. Relational Database Design in DBMS

Chapter 7: Relational Database Design

Chapter 7: Relational Database Design First Normal Form Pitfalls in Relational Database Design Functional Dependencies Decomposition Boyce-Codd Normal Form Third Normal Form Multivalued Dependencies and Fourth Normal Form Overall Database Design Process

First Normal Form Domain is atomic if its elements are considered to be indivisible units Examples of non-atomic domains: Set of names, composite attributes Identification numbers like CS101 that can be broken up into parts A relational schema R is in first normal form if the domains of all attributes of R are atomic Non-atomic values complicate storage and encourage redundant (repeated) storage of data E.g. Set of accounts stored with each customer, and set of owners stored with each account We assume all relations are in first normal form (revisit this in Chapter 9 on Object Relational Databases)

First Normal Form (Contd.) Atomicity is actually a property of how the elements of the domain are used. E.g. Strings would normally be considered indivisible Suppose that students are given roll numbers which are strings of the form CS0012 or EE1127 If the first two characters are extracted to find the department, the domain of roll numbers is not atomic. Doing so is a bad idea: leads to encoding of information in application program rather than in the database.

Pitfalls in Relational Database Design Relational database design requires that we find a “good” collection of relation schemas. A bad design may lead to Repetition of Information. Inability to represent certain information. Design Goals: Avoid redundant data Ensure that relationships among attributes are represented Facilitate the checking of updates for violation of database integrity constraints.

Example Consider the relation schema: Lending-schema = ( branch-name, branch-city, assets, customer-name, loan-number, amount) Redundancy: Data for branch-name, branch-city, assets are repeated for each loan that a branch makes Wastes space Complicates updating, introducing possibility of inconsistency of assets value Null values Cannot store information about a branch if no loans exist Can use null values, but they are difficult to handle.

Decomposition Decompose the relation schema Lending-schema into: Branch-schema = (branch-name, branch-city,assets) Loan-info-schema = (customer-name, loan-number, branch-name, amount) All attributes of an original schema ( R) must appear in the decomposition ( R 1 , R 2 ): R = R 1  R 2 Lossless-join decomposition. For all possible relations r on schema R r =  R1 ( r )  R2 ( r )

Example of Non Lossless-Join Decomposition Decomposition of R = (A, B) R 2 = (A) R 2 = (B) A B    1 2 1 A   B 1 2 r  A ( r )  B( r )  A (r)  B (r) A B     1 2 1 2

Goal — Devise a Theory for the Following Decide whether a particular relation R is in “good” form. In the case that a relation R is not in “good” form, decompose it into a set of relations { R 1 , R 2 , ..., R n } such that each relation is in good form the decomposition is a lossless-join decomposition Our theory is based on: functional dependencies multivalued dependencies

Functional Dependencies Constraints on the set of legal relations. Require that the value for a certain set of attributes determines uniquely the value for another set of attributes. A functional dependency is a generalization of the notion of a key.

Functional Dependencies (Cont.) Let R be a relation schema   R and   R The functional dependency    holds on R if and only if for any legal relations r (R), whenever any two tuples t 1 and t 2 of r agree on the attributes  , they also agree on the attributes  . That is, t 1 [  ] = t 2 [  ]  t 1 [  ] = t 2 [  ] Example: Consider r(A,B) with the following instance of r. On this instance, A  B does NOT hold, but B  A does hold. 4 1 5 3 7

Functional Dependencies (Cont.) K is a superkey for relation schema R if and only if K  R K is a candidate key for R if and only if K  R , and for no   K,   R Functional dependencies allow us to express constraints that cannot be expressed using superkeys. Consider the schema: Loan-info-schema = (customer-name, loan-number, branch-name, amount). We expect this set of functional dependencies to hold: loan-number  amount loan-number  branch-name but would not expect the following to hold: loan-number  customer-name

Use of Functional Dependencies We use functional dependencies to: test relations to see if they are legal under a given set of functional dependencies. If a relation r is legal under a set F of functional dependencies, we say that r satisfies F. specify constraints on the set of legal relations We say that F holds on R if all legal relations on R satisfy the set of functional dependencies F. Note: A specific instance of a relation schema may satisfy a functional dependency even if the functional dependency does not hold on all legal instances. For example, a specific instance of Loan-schema may, by chance, satisfy loan-number  customer-name.

Functional Dependencies (Cont.) A functional dependency is trivial if it is satisfied by all instances of a relation E.g. customer-name, loan-number  customer-name customer-name  customer-name In general,    is trivial if   

Closure of a Set of Functional Dependencies Given a set F set of functional dependencies, there are certain other functional dependencies that are logically implied by F . E.g. If A  B and B  C, then we can infer that A  C The set of all functional dependencies logically implied by F is the closure of F . We denote the closure of F by F + . We can find all of F + by applying Armstrong’s Axioms: if    , then    (reflexivity) if    , then      (augmentation) if    , and    , then    (transitivity) These rules are sound (generate only functional dependencies that actually hold) and complete (generate all functional dependencies that hold).

Example R = (A, B, C, G, H, I) F = { A  B A  C CG  H CG  I B  H } some members of F + A  H by transitivity from A  B and B  H AG  I by augmenting A  C with G, to get AG  CG and then transitivity with CG  I CG  HI from CG  H and CG  I : “union rule” can be inferred from definition of functional dependencies, or Augmentation of CG  I to infer CG  CG I, augmentation of CG  H to infer CGI  HI, and then transitivity

Procedure for Computing F + To compute the closure of a set of functional dependencies F: F + = F repeat for each functional dependency f in F + apply reflexivity and augmentation rules on f add the resulting functional dependencies to F + for each pair of functional dependencies f 1 and f 2 in F + if f 1 and f 2 can be combined using transitivity then add the resulting functional dependency to F + until F + does not change any further NOTE: We will see an alternative procedure for this task later

Closure of Functional Dependencies (Cont.) We can further simplify manual computation of F + by using the following additional rules. If    holds a nd    holds, then     holds (union) If     holds, then    holds and    holds (decomposition) If    holds a nd     holds, then     holds (pseudotransitivity) The above rules can be inferred from Armstrong’s axioms.

Closure of Attribute Sets Given a set of attributes  define the closure of  under F (denoted by  + ) as the set of attributes that are functionally determined by  under F:    is in F +     + Algorithm to compute  + , the closure of  under F result :=  ; while (changes to result ) do for each    in F do begin if   result then result := result   end

Example of Attribute Set Closure R = (A, B, C, G, H, I) F = { A  B A  C CG  H CG  I B  H } ( AG) + 1. result = AG 2. result = ABCG (A  C and A  B) 3. result = ABCGH (CG  H and CG  AGBC) 4. result = ABCGHI (CG  I and CG  AGBCH) Is AG a candidate key? Is AG a super key? Does AG  R? == Is (AG) +  R Is any subset of AG a superkey? Does A  R ? == Is (A) +  R Does G  R ? == Is (G) +  R

Uses of Attribute Closure There are several uses of the attribute closure algorithm: Testing for superkey: To test if  is a superkey, we compute  + , and check if  + contains all attributes of R . Testing functional dependencies To check if a functional dependency    holds (or, in other words, is in F + ), just check if    + . That is, we compute  + by using attribute closure, and then check if it contains  . Is a simple and cheap test, and very useful Computing closure of F For each   R, we find the closure  + , and for each S   + , we output a functional dependency   S.

Canonical Cover Sets of functional dependencies may have redundant dependencies that can be inferred from the others Eg: A  C is redundant in: {A  B, B  C, A  C} Parts of a functional dependency may be redundant E.g. on RHS: {A  B, B  C, A  CD} can be simplified to {A  B, B  C, A  D} E.g. on LHS: {A  B, B  C, AC  D} can be simplified to {A  B, B  C, A  D} Intuitively, a canonical cover of F is a “minimal” set of functional dependencies equivalent to F, having no redundant dependencies or redundant parts of dependencies

Extraneous Attributes Consider a set F of functional dependencies and the functional dependency    in F . Attribute A is extraneous in  if A   and F logically implies ( F – {    })  {(  – A )   }. Attribute A is extraneous in  if A   and the set of functional dependencies ( F – {    })  {   (  – A )} logically implies F. Note: implication in the opposite direction is trivial in each of the cases above, since a “stronger” functional dependency always implies a weaker one Example: Given F = { A  C , AB  C } B is extraneous in AB  C because { A  C, AB  C } logically implies A  C (I.e. the result of dropping B from AB  C ). Example: Given F = { A  C , AB  CD} C is extraneous in AB  CD since A B  C can be inferred even after deleting C

Testing if an Attribute is Extraneous Consider a set F of functional dependencies and the functional dependency    in F . To test if attribute A   is extraneous in  compute ({  } – A) + using the dependencies in F check that ({  } – A) + contains A; if it does, A is extraneous To test if attribute A   is extraneous in  compute  + using only the dependencies in F’ = ( F – {    })  {   (  – A )}, check that  + contains A; if it does , A is extraneous

Canonical Cover A canonical cover for F is a set of dependencies F c such that F logically implies all dependencies in F c, and F c logically implies all dependencies in F, and No functional dependency in F c contains an extraneous attribute, and Each left side of functional dependency in F c is unique. To compute a canonical cover for F : repeat Use the union rule to replace any dependencies in F  1   1 and  1   1 with  1   1  2 Find a functional dependency    with an extraneous attribute either in  or in  If an extraneous attribute is found, delete it from    until F does not change Note: Union rule may become applicable after some extraneous attributes have been deleted, so it has to be re-applied

Example of Computing a Canonical Cover R = ( A, B, C) F = {A  BC B  C A  B AB  C } Combine A  BC and A  B into A  BC Set is now {A  BC, B  C, AB  C } A is extraneous in AB  C Check if the result of deleting A from AB  C is implied by the other dependencies Yes: in fact, B  C is already present! Set is now {A  BC, B  C } C is extraneous in A  BC Check if A  C is logically implied by A  B and the other dependencies Yes : using transitivity on A  B and B  C. Can use attribute closure of A in more complex cases The canonical cover is: A  B B  C

Goals of Normalization Decide whether a particular relation R is in “good” form. In the case that a relation R is not in “good” form, decompose it into a set of relations { R 1 , R 2 , ..., R n } such that each relation is in good form the decomposition is a lossless-join decomposition Our theory is based on: functional dependencies multivalued dependencies

Decomposition Decompose the relation schema Lending-schema into: Branch-schema = (branch-name, branch-city,assets) Loan-info-schema = (customer-name, loan-number, branch-name, amount) All attributes of an original schema ( R) must appear in the decomposition ( R 1 , R 2 ): R = R 1  R 2 Lossless-join decomposition. For all possible relations r on schema R r =  R1 ( r )  R2 ( r ) A decomposition of R into R 1 and R 2 is lossless join if and only if at least one of the following dependencies is in F + : R 1  R 2  R 1 R 1  R 2  R 2

Example of Lossy-Join Decomposition Lossy-join decompositions result in information loss. Example: Decomposition of R = (A, B) R 2 = (A) R 2 = (B) A B    1 2 1 A   B 1 2 r  A ( r )  B( r )  A (r)  B (r) A B     1 2 1 2

Normalization Using Functional Dependencies When we decompose a relation schema R with a set of functional dependencies F into R 1 , R 2 ,.., R n we want Lossless-join decomposition : Otherwise decomposition would result in information loss. No redundancy : The relations R i preferably should be in either Boyce-Codd Normal Form or Third Normal Form. Dependency preservation: Let F i be the set of dependencies F + that include only attributes in R i . Preferably the decomposition should be dependency preserving , that is, ( F 1  F 2  …  F n ) + = F + Otherwise, checking updates for violation of functional dependencies may require computing joins, which is expensive.

Example R = (A, B, C) F = {A  B, B  C) Can be decomposed in two different ways R 1 = (A, B), R 2 = (B, C) Lossless-join decomposition: R 1  R 2 = {B} and B  BC Dependency preserving R 1 = (A, B), R 2 = (A, C) Lossless-join decomposition: R 1  R 2 = {A} and A  A B Not dependency preserving (cannot check B  C without computing R 1 R 2 )

Testing for Dependency Preservation To check if a dependency  is preserved in a decomposition of R into R 1 , R 2 , …, R n we apply the following simplified test (with attribute closure done w.r.t. F) result =  while (changes to result ) do for each R i in the decomposition t = ( result  R i ) +  R i result = result  t If result contains all attributes in  , then the functional dependency    is preserved. We apply the test on all dependencies in F to check if a decomposition is dependency preserving This procedure takes polynomial time, instead of the exponential time required to compute F + and ( F 1  F 2  …  F n ) +

Boyce-Codd Normal Form     is trivial (i.e.,    )  is a superkey for R A relation schema R is in BCNF with respect to a set F of functional dependencies if for all functional dependencies in F + of the form     , where   R and   R , at least one of the following holds:

Example R = (A, B, C) F = {A  B B  C } Key = { A } R is not in BCNF Decomposition R 1 = ( A, B), R 2 = (B, C) R 1 and R 2 in BCNF Lossless-join decomposition Dependency preserving

Testing for BCNF To check if a non-trivial dependency     causes a violation of BCNF 1. compute  + (the attribute closure of  ), and 2. verify that it includes all attributes of R , that is, it is a superkey of R . Simplified test : To check if a relation schema R is in BCNF, it suffices to check only the dependencies in the given set F for violation of BCNF, rather than checking all dependencies in F + . If none of the dependencies in F causes a violation of BCNF, then none of the dependencies in F + will cause a violation of BCNF either. However, using only F is incorrect when testing a relation in a decomposition of R E.g. Consider R ( A, B, C, D ), with F = { A  B, B  C } Decompose R into R 1 ( A,B ) and R 2 ( A,C,D ) Neither of the dependencies in F contain only attributes from ( A,C,D ) so we might be mislead into thinking R 2 satisfies BCNF. In fact, dependency A  C in F + shows R 2 is not in BCNF.

BCNF Decomposition Algorithm result := { R }; done := false; compute F + ; while (not done) do if (there is a schema R i in result that is not in BCNF) then begin let     be a nontrivial functional dependency that holds on R i such that    R i is not in F + , and    =  ; result := ( result – R i )  ( R i –  )  (  ,  ); end else done := true; Note: each R i is in BCNF, and decomposition is lossless-join.

Example of BCNF Decomposition R = (branch-name, branch-city, assets, customer-name, loan-number, amount) F = {branch-name  assets branch-city loan-number  amount branch-name} Key = {loan-number, customer-name} Decomposition R 1 = ( branch-name, branch-city, assets) R 2 = (branch-name, customer-name, loan-number, amount) R 3 = (branch-name, loan-number, amount) R 4 = (customer-name, loan-number) Final decomposition R 1 , R 3 , R 4

Testing Decomposition for BCNF To check if a relation R i in a decomposition of R is in BCNF, Either test R i for BCNF with respect to the restriction of F to R i (that is, all FDs in F + that contain only attributes from R i ) or use the original set of dependencies F that hold on R , but with the following test: for every set of attributes   R i , check that  + (the attribute closure of  ) either includes no attribute of R i -  , or includes all attributes of R i . If the condition is violated by some     in F , the dependency    (  + -   )  R i can be shown to hold on R i , and R i violates BCNF. We use above dependency to decompose R i

BCNF and Dependency Preservation R = (J, K, L) F = {JK  L L  K } Two candidate keys = JK and JL R is not in BCNF Any decomposition of R will fail to preserve JK  L It is not always possible to get a BCNF decomposition that is dependency preserving

Third Normal Form: Motivation There are some situations where BCNF is not dependency preserving, and efficient checking for FD violation on updates is important Solution: define a weaker normal form, called Third Normal Form. Allows some redundancy (with resultant problems; we will see examples later) But FDs can be checked on individual relations without computing a join. There is always a lossless-join, dependency-preserving decomposition into 3NF.

Third Normal Form A relation schema R is in third normal form (3NF) if for all:    in F + at least one of the following holds:    is trivial (i.e.,    )  is a superkey for R Each attribute A in  –  is contained in a candidate key for R. (NOTE: each attribute may be in a different candidate key) If a relation is in BCNF it is in 3NF (since in BCNF one of the first two conditions above must hold). Third condition is a minimal relaxation of BCNF to ensure dependency preservation (will see why later).

3NF (Cont.) Example R = (J, K, L) F = { JK  L, L  K } Two candidate keys: JK and JL R is in 3NF JK  L JK is a superkey L  K K is contained in a candidate key BCNF decomposition has (JL) and (LK) Testing for JK  L requires a join There is some redundancy in this schema Equivalent to example in book: Banker-schema = (branch-name, customer-name, banker-name) banker-name  branch name branch name customer-name  banker-name

Testing for 3NF Optimization: Need to check only FDs in F , need not check all FDs in F + . Use attribute closure to check for each dependency    , if  is a superkey. If  is not a superkey, we have to verify if each attribute in  is contained in a candidate key of R this test is rather more expensive, since it involve finding candidate keys testing for 3NF has been shown to be NP-hard Interestingly, decomposition into third normal form (described shortly) can be done in polynomial time

3NF Decomposition Algorithm Let F c be a canonical cover for F; i := 0; for each functional dependency    in F c do if none of the schemas R j , 1  j  i contains   then begin i := i + 1; R i :=   end if none of the schemas R j , 1  j  i contains a candidate key for R then begin i := i + 1; R i := any candidate key for R; end return (R 1 , R 2 , ..., R i )

3NF Decomposition Algorithm (Cont.) Above algorithm ensures: each relation schema R i is in 3NF decomposition is dependency preserving and lossless-join Proof of correctness is at end of this file ( click here )

Example Relation schema: Banker-info-schema = (branch-name, customer-name, banker-name, office-number) The functional dependencies for this relation schema are: banker-name  branch-name office-number customer-name branch-name  banker-name The key is: { customer-name, branch-name }

Applying 3NF to Banker-info-schema The for loop in the algorithm causes us to include the following schemas in our decomposition: Banker-office-schema = (banker-name, branch-name, office-number) Banker-schema = (customer-name, branch-name, banker-name) Since Banker-schema contains a candidate key for Banker-info-schema, we are done with the decomposition process.

Comparison of BCNF and 3NF It is always possible to decompose a relation into relations in 3NF and the decomposition is lossless the dependencies are preserved It is always possible to decompose a relation into relations in BCNF and the decomposition is lossless it may not be possible to preserve dependencies.

Comparison of BCNF and 3NF (Cont.) Example of problems due to redundancy in 3NF R = (J, K, L) F = { JK  L, L  K } J j 1 j 2 j 3 null L l 1 l 1 l 1 l 2 K k 1 k 1 k 1 k 2 A schema that is in 3NF but not in BCNF has the problems of repetition of information (e.g., the relationship l 1 , k 1 ) need to use null values (e.g., to represent the relationship l 2 , k 2 where there is no corresponding value for J ).

Design Goals Goal for a relational database design is: BCNF. Lossless join. Dependency preservation. If we cannot achieve this, we accept one of Lack of dependency preservation Redundancy due to use of 3NF Interestingly, SQL does not provide a direct way of specifying functional dependencies other than superkeys. Can specify FDs using assertions, but they are expensive to test Even if we had a dependency preserving decomposition, using SQL we would not be able to efficiently test a functional dependency whose left hand side is not a key.

Testing for FDs Across Relations If decomposition is not dependency preserving, we can have an extra materialized view for each dependency   in F c that is not preserved in the decomposition The materialized view is defined as a projection on   of the join of the relations in the decomposition Many newer database systems support materialized views and database system maintains the view when the relations are updated. No extra coding effort for programmer. The functional dependency    is expressed by declaring  as a candidate key on the materialized view. Checking for candidate key cheaper than checking    BUT: Space overhead: for storing the materialized view Time overhead: Need to keep materialized view up to date when relations are updated Database system may not support key declarations on materialized views

Multivalued Dependencies There are database schemas in BCNF that do not seem to be sufficiently normalized Consider a database classes(course, teacher, book) such that ( c,t,b )  classes means that t is qualified to teach c, and b is a required textbook for c The database is supposed to list for each course the set of teachers any one of which can be the course’s instructor, and the set of books, all of which are required for the course (no matter who teaches it).

There are no non-trivial functional dependencies and therefore the relation is in BCNF Insertion anomalies – i.e., if Sara is a new teacher that can teach database, two tuples need to be inserted (database, Sara, DB Concepts) (database, Sara, Ullman) Multivalued Dependencies (Cont.) course teacher book database database database database database database operating systems operating systems operating systems operating systems Avi Avi Hank Hank Sudarshan Sudarshan Avi Avi Jim Jim DB Concepts Ullman DB Concepts Ullman DB Concepts Ullman OS Concepts Shaw OS Concepts Shaw classes

Therefore, it is better to decompose classes into: Multivalued Dependencies (Cont.) course teacher database database database operating systems operating systems Avi Hank Sudarshan Avi Jim teaches course book database database operating systems operating systems DB Concepts Ullman OS Concepts Shaw text We shall see that these two relations are in Fourth Normal Form (4NF)

Multivalued Dependencies (MVDs) Let R be a relation schema and let   R and   R. The multivalued dependency    holds on R if in any legal relation r(R), for all pairs for tuples t 1 and t 2 in r such that t 1 [  ] = t 2 [  ], there exist tuples t 3 and t 4 in r such that: t 1 [  ] = t 2 [  ] = t 3 [  ] t 4 [  ] t 3 [  ] = t 1 [  ] t 3 [ R –  ] = t 2 [ R –  ] t 4 [  ] = t 2 [  ] t 4 [ R –  ] = t 1 [ R –  ]

MVD (Cont.) Tabular representation of   

Example Let R be a relation schema with a set of attributes that are partitioned into 3 nonempty subsets. Y, Z, W We say that Y  Z (Y multidetermines Z) if and only if for all possible relations r(R) < y 1 , z 1 , w 1 >  r and < y 2 , z 2 , w 2 >  r then < y 1 , z 1 , w 2 >  r and < y 2 , z 2 , w 1 >  r Note that since the behavior of Z and W are identical it follows that Y  Z if Y  W

Example (Cont.) In our example: course  teacher course  book The above formal definition is supposed to formalize the notion that given a particular value of Y ( course ) it has associated with it a set of values of Z (teacher) and a set of values of W (book) , and these two sets are in some sense independent of each other. Note: If Y  Z then Y  Z Indeed we have (in above notation) Z 1 = Z 2 The claim follows.

Use of Multivalued Dependencies We use multivalued dependencies in two ways: 1. To test relations to determine whether they are legal under a given set of functional and multivalued dependencies 2. To specify constraints on the set of legal relations. We shall thus concern ourselves only with relations that satisfy a given set of functional and multivalued dependencies. If a relation r fails to satisfy a given multivalued dependency, we can construct a relations r  that does satisfy the multivalued dependency by adding tuples to r.

Theory of MVDs From the definition of multivalued dependency, we can derive the following rule: If    , then    That is, every functional dependency is also a multivalued dependency The closure D + of D is the set of all functional and multivalued dependencies logically implied by D . We can compute D + from D , using the formal definitions of functional dependencies and multivalued dependencies. We can manage with such reasoning for very simple multivalued dependencies, which seem to be most common in practice For complex dependencies, it is better to reason about sets of dependencies using a system of inference rules (see Appendix C).

Fourth Normal Form A relation schema R is in 4NF with respect to a set D of functional and multivalued dependencies if for all multivalued dependencies in D + of the form    , where   R and   R, at least one of the following hold:    is trivial (i.e.,    or    = R)  is a superkey for schema R If a relation is in 4NF it is in BCNF

Restriction of Multivalued Dependencies The restriction of D to R i is the set D i consisting of All functional dependencies in D + that include only attributes of R i All multivalued dependencies of the form   (   R i ) where   R i and    is in D +

4NF Decomposition Algorithm result: = { R }; done := false; compute D + ; Let D i denote the restriction of D + to R i while ( not done ) if (there is a schema R i in result that is not in 4NF) then begin let    be a nontrivial multivalued dependency that holds on R i such that   R i is not in D i , and  ; result := ( result - R i )  ( R i -  )  (  ,  ); end else done := true; Note: each R i is in 4NF, and decomposition is lossless-join

Example R =( A, B, C, G, H, I ) F ={ A  B B  HI CG  H } R is not in 4NF since A  B and A is not a superkey for R Decomposition a) R 1 = ( A, B ) ( R 1 is in 4NF) b) R 2 = ( A, C, G, H, I ) ( R 2 is not in 4NF) c) R 3 = ( C, G, H ) ( R 3 is in 4NF) d) R 4 = ( A, C, G, I ) ( R 4 is not in 4NF) Since A  B and B  HI , A  HI , A  I e) R 5 = ( A, I ) ( R 5 is in 4NF) f) R 6 = (A, C, G) (R 6 is in 4NF)

Further Normal Forms Join dependencies generalize multivalued dependencies lead to project-join normal form ( PJNF) (also called fifth normal form ) A class of even more general constraints, leads to a normal form called domain-key normal form . Problem with these generalized constraints: are hard to reason with, and no set of sound and complete set of inference rules exists. Hence rarely used

Overall Database Design Process We have assumed schema R is given R could have been generated when converting E-R diagram to a set of tables. R could have been a single relation containing all attributes that are of interest (called universal relation ). Normalization breaks R into smaller relations. R could have been the result of some ad hoc design of relations, which we then test/convert to normal form.

ER Model and Normalization When an E-R diagram is carefully designed, identifying all entities correctly, the tables generated from the E-R diagram should not need further normalization. However, in a real (imperfect) design there can be FDs from non-key attributes of an entity to other attributes of the entity E.g. employee entity with attributes department-number and department-address , and an FD department-number  department-address Good design would have made department an entity FDs from non-key attributes of a relationship set possible, but rare --- most relationships are binary

Universal Relation Approach Dangling tuples – Tuples that “disappear” in computing a join. Let r 1 ( R 1 ), r 2 ( R 2 ), …., r n ( R n ) be a set of relations A tuple r of the relation r i is a dangling tuple if r is not in the relation:  Ri ( r 1 r 2 … r n ) The relation r 1 r 2 … r n is called a universal relation since it involves all the attributes in the “universe” defined by R 1  R 2  …  R n If dangling tuples are allowed in the database, instead of decomposing a universal relation, we may prefer to synthesize a collection of normal form schemas from a given set of attributes.

Universal Relation Approach Dangling tuples may occur in practical database applications. They represent incomplete information E.g. may want to break up information about loans into: (branch-name, loan-number) (loan-number, amount) (loan-number, customer-name) Universal relation would require null values, and have dangling tuples

Universal Relation Approach (Contd.) A particular decomposition defines a restricted form of incomplete information that is acceptable in our database. Above decomposition requires at least one of customer-name, branch-name or amount in order to enter a loan number without using null values Rules out storing of customer-name, amount without an appropriate loan-number (since it is a key, it can't be null either!) Universal relation requires unique attribute names unique role assumption e.g. customer-name , branch-name Reuse of attribute names is natural in SQL since relation names can be prefixed to disambiguate names

Denormalization for Performance May want to use non-normalized schema for performance E.g. displaying customer-name along with account-number and balance requires join of account with depositor Alternative 1: Use denormalized relation containing attributes of account as well as depositor with all above attributes faster lookup Extra space and extra execution time for updates extra coding work for programmer and possibility of error in extra code Alternative 2: use a materialized view defined as account depositor Benefits and drawbacks same as above, except no extra coding work for programmer and avoids possible errors

Other Design Issues Some aspects of database design are not caught by normalization Examples of bad database design, to be avoided: Instead of earnings ( company-id, year, amount ), use earnings-2000, earnings-2001, earnings-2002 , etc., all on the schema ( company-id, earnings ). Above are in BCNF, but make querying across years difficult and needs new table each year company-year ( company-id, earnings-2000, earnings-2001, earnings-2002 ) Also in BCNF, but also makes querying across years difficult and requires new attribute each year. Is an example of a crosstab , where values for one attribute become column names Used in spreadsheets, and in data analysis tools

Proof of Correctness of 3NF Decomposition Algorithm

Correctness of 3NF Decomposition Algorithm 3NF decomposition algorithm is dependency preserving (since there is a relation for every FD in F c ) Decomposition is lossless join A candidate key ( C ) is in one of the relations R i in decomposition Closure of candidate key under F c must contain all attributes in R . Follow the steps of attribute closure algorithm to show there is only one tuple in the join result for each tuple in R i

Correctness of 3NF Decomposition Algorithm (Contd.) Claim: if a relation R i is in the decomposition generated by the above algorithm, then R i satisfies 3NF. Let R i be generated from the dependency   Let   B be any non-trivial functional dependency on R i . (We need only consider FDs whose right-hand side is a single attribute.) Now, B can be in either  or  but not in both. Consider each case separately.

Correctness of 3NF Decomposition (Contd.) Case 1: If B in  : If  is a superkey, the 2nd condition of 3NF is satisfied Otherwise  must contain some attribute not in  Since   B is in F + it must be derivable from F c , by using attribute closure on  . Attribute closure not have used   - if it had been used,  must be contained in the attribute closure of  , which is not possible, since we assumed  is not a superkey. Now, using  (  - {B}) and   B, we can derive   B (since     , and B   since   B is non-trivial) Then, B is extraneous in the right-hand side of   ; which is not possible since   is in F c . Thus, if B is in  then  must be a superkey, and the second condition of 3NF must be satisfied.

Correctness of 3NF Decomposition (Contd.) Case 2: B is in  . Since  is a candidate key, the third alternative in the definition of 3NF is trivially satisfied. In fact, we cannot show that  is a superkey. This shows exactly why the third alternative is present in the definition of 3NF. Q.E.D.

The Relation branch-customer customer-loan

Tabular Representation of 

Relation bc : An Example of Reduncy in a BCNF Relation

7. Relational Database Design in DBMS

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7. Relational Database Design in DBMS