Link Search Menu Expand Document

Limited Expressivity for finite automata

DFA (Deterministic Finite Automata) are very efficient at parsing regular languages. Unfortunately the expressivity of regular languages is limited.

For example, the language \(L_{ab} = {a^nb^n}\) with \(n \geq 1\) is not regular.

Indeed, a DFA has a finite number of states. Recognizing words in \(L_{ab}\) requires keeping track of the number of \(a\) symbols to ensure that it matches the number of \(b\) symbols. Since \(n\) is arbitrarily large, this is not possible with a finite number of states.

More formally, this can be proved by contradiction with the pumping lemma (Lemme de l’étoile).


Suppose \(L_{ab}\) is regular.

The pumping lemma tells us that there is a \(n \geq 1\) such that every string \(w \in L_{ab}\) with \(|w| \geq n\) can be written as \(w = xyz\) satisfying

  • \[|y| \geq 1\]
  • \[|xy| \leq n\]
  • \[\forall k \geq 0, xy^kz \in L_{ab}\]

Let us consider \(w = a^nb^n\) which belongs to \(L_{ab}\) and is longer than \(n\). Therefore \(a^nb^n\) can be decomposed as three substrings \(xyz\) satisfying above properties. Because \(|xy| \leq n\), \(x\) and \(y\) are only composed of \(a\) symbols.

Let us consider \(w'=xy^2z\). Since \(y\) is only composed of \(a\) symbols, \(w'\) has exactly \(n\) symbols \(b\) and strictly more than \(n\) symbols \(a\). Therefore \(w'\) does not belong to \(L_{ab}\), contradicting the pumping lemma.

Most programming languages include recursively nested constructions. For example nested parenthesis, such as ((3+(7/9)) - 2) or nested control blocks such as if a then (if b then c else d) else f. These constructs with arbitrarily deep recursive nesting cannot be parsed with DFA, as we just saw.

Since we cannot use DFA, parsers will use the more expressive formalism of Context Free Grammars.

Context Free Grammars


A context-free grammar \(G\) is defined by four elements, \(G=(V, S, \Sigma, P)\), where

  1. \(V\) is the finite alphabet of the non-terminal symbols.
  2. \(S\) is the start non-terminal symbol and belongs to \(V\).
  3. \(\Sigma\) is the finite alphabet of the terminal symbols. \(V\) and \(\Sigma\) are disjoint.
  4. \(P\) is a finite set of grammar rules. \(P \subset V \times (V \cup \Sigma)^*\)

A grammar rule is usually written, with \(\alpha \in V\) and \(\beta_1, \beta_2, \ldots, \beta_k \in (V \cup \Sigma)\), as

\(\alpha \rightarrow \beta_1 \beta_2 \ldots \beta_k\).

In the above expression,

  • \(\alpha\) is a non-terminal.
  • \(\beta_1, \ldots \beta_k\) are either terminals or non-terminals.


Consider a grammar \(G_1\) with

  1. \[V = \{S, T\}\]
  2. \(S\) start symbol
  3. \[\Sigma = \{a,b,\#\}\]
  4. The following \(P\) set of grammar rules,
\[\begin{aligned} S &\rightarrow T\#\\ T &\rightarrow aTb \\ T &\rightarrow ab \\ \end{aligned}\]

Applying grammar rules

When applying a grammar rule, we replace inside a word, the left-hand-side symbol \(\alpha\) by the right-hand-side symbols \(\beta_1\ldots\beta_k\).

Let us have two words \(u,v \in (V \cup \Sigma)^*\). We will say that \(v\) is produced from \(u\), when \(u\) is obtained by replacing left-hand-side symbol in \(u\) by its right-hand-side counterpart in a grammar. We will note \(u \Rightarrow v\).

In case of one or more rule applications, \(u_1 \Rightarrow u_2 \Rightarrow \ldots v\), we will say that \(v\) derives from \(u\) and we will write \(u\stackrel{*}{\Rightarrow} v\).


Consider grammar \(G_1\) from previous example.

The word, \(aaabbb\#\) derives from \(S\), indeed \(\begin{aligned} S &\Rightarrow T\# &&\text{(by applying first rule)} \\ T\# &\Rightarrow aTb\# &&\text{(by applying second rule)} \\ aTb\# &\Rightarrow aaTbb\# &&\text{(by applying second rule)} \\ aaTbb\# &\Rightarrow aaabbb\# &&\text{(by applying third rule)} \\ \end{aligned}\)

Context Free Grammars are called context free because rules can be applied independently of the context, that is to say the neighboring symbols.

Language recognized by a grammar

We will say that the language \(L(G)\) recognized by a grammar \(G\) is the set of all words composed of terminal symbols that derive from \(S\).

\[L(G) = \{w \in \Sigma* | S \stackrel{*}{\Rightarrow} w \}\]

For instance, in the previous example, \(L(G_1) = L_{ab}\)

Parse Trees and ambiguous grammars

Given grammar \(G_2\),

  1. \[V = \{S, E\}\]
  2. \(S\) start symbol
  3. \[\Sigma = \{num, \#\}\]
  4. The following \(P\) set of grammar rules,
\[\begin{aligned} S &\rightarrow E\#\\ E &\rightarrow E + E\\ E &\rightarrow num \\ \end{aligned}\]

Let’s consider the word, \(num + num + num\). This word belongs to \(L(G_2)\) because \(S \Rightarrow E\# \Rightarrow E+E\# \\ \Rightarrow E+num\# \Rightarrow E+E+num\# \\ \Rightarrow E+num+num\# \Rightarrow num+num+num\#\)

This first derivation can also be represented as a Parse Tree,

Parse tree for first derivation
           |                       ((num+num)+num)
         E + E#
        /     \
      E + E   num
      |   |
     num num 

Interestingly, in this grammar a second derivation also produces this same word. \(S \Rightarrow E\# \Rightarrow E+E\# \\ \Rightarrow num+E\# \Rightarrow num+E+E\# \\ \Rightarrow num+num+E\# \Rightarrow num+num+num\#\)

Parse tree for second derivation
           |                       (num+(num+num))
         E + E#
        /     \
       num   E + E  
             |   |
            num num 

When the same word admits multiple parse trees, we will say that the grammar is ambiguous.

Backus-Naur Form (BNF)

A common notation to represent grammars is the Backus-Naur Form.
This notation will be used during lab 2.