Showing that the reverse of a word for a regular language L is also regular
I am confused as to how I am to approach this question, i've been stuck for hours: For a word x, we use x^r to denote its reverse. For a language L, we use L^r to denote {x^r where x is in the set of L}. Show that if L is regular then so is L^r
If L is regular, then there exists some regular grammar which generates it. It can be always represented as either a left-regular grammar, or a right-regular grammar. Let's assume that it's left-regular grammar G_l(the proof for right-regular grammar is analogous).
This grammar has productions of two types; the terminating-type:
A -> a, where A is non-terminal and a is either a terminal or empty string (epsilon)
or the chaining type:
B -> Ca, where B, C are non-terminals and a is a terminal
When we apply reverse to a regular language, we basically also apply it to the tails of productions (since heads are just single non-terminals). It's going to be proved later on. So we get a new grammar G_r, with productions:
A -> a, where A is non-terminal and a is either a terminal or empty string (epsilon)
B -> aC, where B, C are non-terminals and a is a terminal
But hey, it's a right-regular grammar! So the language it accepts is also regular.
There is one thing to do - to show that reversing tails actually does the thing it's supposed to. We're going to prove it very simply:
If L contains \epsilon, then there is production 'S -> \epsilon' in G_l. Since we don't touch productions like that, it's also present in G_r.
If L contains a, a word composed of a single terminal, then it's similar to the above
If L contains aZ, where a is a terminal and Z is a word from the language constructed from chopping off the first terminals out of words in L, then L^r contains (because of changes to the chaining productions) (Z^r)a. Z is also a regular language, since it can be constructed by dropping the first "level" of left-productions from G_l, which leaves us with a regular grammar.
I hope it helped. There's also an arguably easier way of doing that by reversing edges of the relevant finite automata and changing accepting and entry states a bit.
Related
So i have this grammar :
S -> (D)
D -> EF
E -> a|b|S
F -> *D | +D | ε
First of all, books solution uses the P -> pBq , First(q) - {ε} is subset of FOLLOW(B) for the rule D -> EF but that rule has only 2 symbols do we assume ε infront of E (ε being the p in pBq)?
And secondly i can't understand how to calculate Follow(E).
FOLLOW(E) consists of every terminal symbol which can immediately follow E in some derivation step. That's the precise definition; it's not very complicated.
For a simple grammar, you should be able to figure out all the FOLLOW sets just be looking at the grammar and applying a little bit of common sense. It would probably be a good idea to do that, since it will give you a better idea of how the algorithm works.
As a side note, it's maybe worth mentioning that ε is not a thing. Or at least, it's not a grammar symbol. It's one of several conventions used to make the empty sequence visible, just like 0 is a way to make nothing visible. Sometimes that's useful, but it's important to not let it confuse you. (Abuse of notation is endemic in mathematics, which can be frustrating.)
So, what can follow E? E only appears in one place on the right-hand sdie of that grammar, in the production D → E F. So clearly any symbol which be the first symbol of F must be in FOLLOW(E). The symbols which could be at the start of F are + and *, since as mentioned, ε is not a grammar symbol. (Many definitions of FIRST allow ε to be a member of that set, along with any actual terminal symbol. That's an example of the abuse of notation I was talking about in the previous paragraph, since it makes it look like ε is a terminal symbol. But it isn't. It's nothing.)
F is what we call a "nullable" non-terminal, because it can derive the empty sequence (which was written as ε so that you can see it). In other words, it's possible for F to disappear completely in a derivation step. And if it does disappear, then E might be at the end of the production D → E F. If E is at the end of D, then it can be followed by anything which could follow D, which includes ). D can also appear at the end of a derivation of F, which means that F could be followed by anything which could follow F, a tautology which adds no information whatsoever.
So it's easy to see that FOLLOW(F) = {*, +, )}, and you can use that to check your understanding of any algorithm to compute follow sets.
Now, I don't know what book you are referring to (and it would have been courteous to mention that in your original question; sources should always be correctly cited). But the book I happen to have in front of me --the Dragon Book-- has a pretty similar algorithm. The Dragon book uses a simple convention for writing statements like that. Probably your book does, too, but it might not be the same convention. You should check what it says and make sure that you typed the copied statement correctly, respecting whatever formatting used to indicate what the symbols stand for.
In the Dragon book, some of the conventions include:
Lower case characters at the start of the alphabet. –a, b, c,…– are terminals (as well as actual symbols like * and +).
Upper case characters at the start of the alphabet. –A, B, C,…– are non-terminals.
S is the start symbol.
Upper case characters at the end of the alphabet. –X, Y, Z– stand for arbitrary grammar symbols (either terminals or non-terminals).
$ is the marker used to indicate the end of the input.
Lower-case Greek letters –α, β, γ,…– are possibly-empty strings of grammar symbols.
The phrase "possibly empty" is very important, so I'm repeating it.
With that convention, they write the rules for computing the FOLLOW set:
Place $ in FOLLOW(S).
For every production A → αBβ, copy everything from FIRST(&beta) except ε into FOLLOW(B).
If there is a production A → αB or a production A → αBβ where FIRST(β) contains ε, place everything in FOLLOW(A) into FOLLOW(B).
As mentioned above, α is a possibly-empty string of grammar symbols. So it might not be visible.
Keep doing steps 2 and 3 until no new symbols are added to any follow set.
I'm pretty sure that the algorithm in your book differs only in notation conventions.
Goal: find a way to formally define a grammar that recognizes elements from a set 0 or 1 times in any order. Subsequently, I want to parse it and generate an AST as well.
For example: Say the set of valid strings in my language is {A, B, C}. I want to define a grammar that recognizes all valid permutations of any number of those elements.
Syntactically valid strings would include:
(the empty string)
A,
B A, and
C A B
Syntactically invalid strings would include:
A A, and
B A C B
To be clear, defining all possible permutations explicitly in a CFG is unacceptable for my purposes, since larger sets would be impossible to maintain.
From what I understand, such a language fails the pumping lemma for context free languages, so the solution will not be context free or regular.
Update
What I'm after is called a "permutation language", which Benedek Nagy has done some theoretical work on as an extension to context free languages.
Regarding a parser generator, I've only found talk of implementing parsers with a permutation phase (link). Parsers evidently have an exponential lower bound on the size of resulting CFG, and I haven't found any parser generators that support it anyhow.
A sort-of solution to this problem was written in ANTLR. It uses semantic predicates to 'code around' the issue.
Assuming that the set of alternative strings is fixed and known in advance, say of size n, one can come up with a (non context-free) grammar of size O(n!). This is not asymptotically smaller than enumerating all permutations, so I suppose it cannot be considered a good solution. I believe that this grammar can be reformulated as a context-sensitive grammar (although in the form I'm suggesting below it is not).
For the example {a, b, c} mentioned in the question, one such grammar is the following. I'm using lower case letters for terminal symbols and upper case letters for non-terminals, as is customary. S is the initial non-terminal symbol.
S ::= XabcY
XabcY ::= aXbcY | bXacY | cXabY
XabY ::= ab | ba
XacY ::= ac | ca
XbcY ::= bc | cb
Non-terminals X and Y enclose the substring in the production which has not been finalized yet; this substring will eventually be replaced by a permutation of the terminals that are given between X and Y (in some arbitrary order).
I can read on Wikipedia the formal definition of a Production, however when you start reading that article, it makes an assumption about prior knowledge.
Wikipedia defines it as follows:
A production or production rule in computer science is a rewrite rule specifying a symbol substitution that can be recursively performed to generate new symbol sequences.
This assumes that I know and understand what a rewrite rule is. I don't, and if I click the link, I get into another fairly technical explanation.
Can someone explain to me in plain English what a Production actually is?
Note: I have made many attempts to understand this, but I don't think I've succeeded. From what I can tell it rewrites the given string in terms of grammar rules. Not sure if I'm correct.
To explain what a production is I'd like to introduce a bit of context first.
The dragon book states that a context free grammar has 4 components:
a set of terminal symbols (tokens)
a set of non-terminal symbols (syntactic variables)
a set of productions of the form: non-term --> sequence of terminals and non-terminals
a non-terminal symbol designated as the start symbol
It is also said that parsing is the problem of taking a string of terminals (the source code) and figuring out what are the steps required to derive this string of terminals from the start symbol of the grammar.
Now that this has been said, a production is essentially a possible (intermediate) step. I say possible because some symbols can derive into different sequences.
For example, let's make a simple grammar to represent an arbitrarily long sequences of a's ending with a b. The 4 components of this grammar would be:
Terminals: a, b
Non-terminals: S, X
Rules: S --> X, X --> aX, X --> ab
Start symbol: S
From the description I gave above "aaaab" should be derivable from this grammar. Let's see if that holds up. We start from, the start symbol, and then apply productions until a) we get the final sequence, b) we exhaust all possibilities without succeeding (meaning the sequence is not "grammatically correct").
S
X (after applying S --> X)
aX (after applying X --> aX)
aaX (after applying X --> aX)
aaaX (after applying X --> aX)
aaaab (after applying X --> ab)
And we're done, we got the original sequence. So as you can see we re-wrote the non-terminal symbols by applying rules (one of them we applied recursively) which transformed the sequence into a new sequence of symbols at every step and we did so until we had the final sequence.
A rewrite rule is a method of replacing subterms of a formula with other terms. In their most basic form, they consist of a set of objects, plus relations on how to transform those objects.
An example of a rewrite rule could look like:
A → B
Now as for what this actually does! You are right on your note, take for example a list of things (and 2 rewrite rules):
X, Y, Z
X → Y
Y → Z
Which would result in:
Z, Z, Z
A production rule is a rewrite rule because it is a method of replacing subterms of a formula (probably a string in your case). A production rule could look like:
X, Y, Z
X → aX
By using the rule in such a way it becomes possible to apply recursion (create new sequences) as it will keep replacing itself:
aX, Y, Z
aaX, Y, Z
aaaX, Y, Z
As for the question you are asking, you could say: "A production rule is a replacement rule for formulas that uses recursion to create new sequences".
What are FIRST and FOLLOW sets? What are they used for in parsing?
Are they used for top-down or bottom-up parsers?
Can anyone explain me FIRST and FOLLOW SETS for the following set of grammar rules:
E := E+T | T
T := T*V | T
V := <id>
They are typically used in LL (top-down) parsers to check if the running parser would encounter any situation where there is more than one way to continue parsing.
If you have the alternative A | B and also have FIRST(A) = {"a"} and FIRST(B) = {"b", "a"} then you would have a FIRST/FIRST conflict because when "a" comes next in the input you wouldn't know whether to expand A or B. (Assuming lookahead is 1).
On the other side if you have a Nonterminal that is nullable like AOpt: ("a")? then you have to make sure that FOLLOW(AOpt) doesn't contain "a" because otherwise you wouldn't know if to expand AOpt or not like here: S: AOpt "a" Either S or AOpt could consume "a" which gives us a FIRST/FOLLOW conflict.
FIRST sets can also be used during the parsing process for performance reasons. If you have a nullable nonterminal NullableNt you can expand it in order to see if it can consume anything, or it may be faster to check if FIRST(NullableNt) contains the next token and if not simply ignore it (backtracking vs predictive parsing). Another performance improvement would be to additionally provide the lexical scanner with the current FIRST set, so the scanner does not try all possible terminals but only those that are currently allowed by the context. This conflicts with reserved terminals but those are not always needed.
Bottom up parsers have different kinds of conflicts namely Reduce/Reduce and Shift/Reduce. They also use item sets to detect conflicts and not FIRST,FOLLOW.
Your grammar would't work with LL-parsers because it contains left recursion. But the FIRST sets for E, T and V would be {id} (assuming your T := T*V | T is meant to be T := T*V | V).
Answer :
E->E+T|T
left recursion
E->TE'
E'->+TE'|eipsilon
T->T*V|T
left recursion
T->VT'
T'->*VT'|epsilon
no left recursion in
V->(id)
Therefore the grammar is:
E->TE'
E'->+TE'|epsilon
T->VT'
T'->*VT'|epsilon
V-> (id)
FIRST(E)={(}
FIRST(E')={+,epsilon}
FIRST(T)={(}
FIRST(T')={*,epsilon}
FIRST(V)={(}
Starting Symbol=FOLLOW(E)={$}
E->TE',E'->TE'|epsilon:FOLLOW(E')=FOLLOW(E)={$}
E->TE',E'->+TE'|epsilon:FOLLOW(T)=FIRST(E')={+,$}
T->VT',T'->*VT'|epsilon:FOLLOW(T')=FOLLOW(T)={+,$}
T->VT',T->*VT'|epsilon:FOLLOW(V)=FIRST(T)={ *,epsilon}
Rules for First Sets
If X is a terminal then First(X) is just X!
If there is a Production X → ε then add ε to first(X)
If there is a Production X → Y1Y2..Yk then add first(Y1Y2..Yk) to first(X)
First(Y1Y2..Yk) is either
First(Y1) (if First(Y1) doesn't contain ε)
OR (if First(Y1) does contain ε) then First (Y1Y2..Yk) is everything in First(Y1) except for ε as well as everything in First(Y2..Yk)
If First(Y1) First(Y2)..First(Yk) all contain ε then add ε to First(Y1Y2..Yk) as well.
Rules for Follow Sets
First put $ (the end of input marker) in Follow(S) (S is the start symbol)
If there is a production A → aBb, (where a can be a whole string) then everything in FIRST(b) except for ε is placed in FOLLOW(B).
If there is a production A → aB, then everything in FOLLOW(A) is in FOLLOW(B)
If there is a production A → aBb, where FIRST(b) contains ε, then everything in FOLLOW(A) is in FOLLOW(B)
Wikipedia is your friend. See discussion of LL parsers and first/follow sets.
Fundamentally they are used as the basic for parser construction, e.g., as part of parser generators. You can also use them to reason about properties of grammars, but most people don't have much of a need to do this.
I want to test whether two languages have a string in common. Both of these languages are from a subset of regular languages described below and I only need to know whether there exists a string in both languages, not produce an example string.
The language is specified by a glob-like string like
/foo/**/bar/*.baz
where ** matches 0 or more characters, and * matches zero or more characters that are not /, and all other characters are literal.
Any ideas?
thanks,
mike
EDIT:
I implemented something that seems to perform well, but have yet to try a correctness proof. You can see the source and unit tests
Build FAs A and B for both languages, and construct the "intersection FA" AnB. If AnB has at least one accepting state accessible from the start state, then there is a word that is in both languages.
Constructing AnB could be tricky, but I'm sure there are FA textbooks that cover it. The approach I would take is:
The states of AnB is the cartesian product of the states of A and B respectively. A state in AnB is written (a, b) where a is a state in A and b is a state in B.
A transition (a, b) ->r (c, d) (meaning, there is a transition from (a, b) to (c, d) on symbol r) exists iff a ->r c is a transition in A, and b ->r d is a transition in B.
(a, b) is a start state in AnB iff a and b are start states in A and B respectively.
(a, b) is an accepting state in AnB iff each is an accepting state in its respective FA.
This is all off the top of my head, and hence completely unproven!
I just did a quick search and this problem is decidable (aka can be done), but I don't know of any good algorithms to do it. One is solution is:
Convert both regular expressions to NFAs A and B
Create a NFA, C, that represents the intersection of A and B.
Now try every string from 0 to the number of states in C and see if C accepts it (since if the string is longer it must repeat states at one point).
I know this might be a little hard to follow but this is only way I know how.