For other uses, see Consistency (disambiguation).

In classical deductive logic, a consistent theory is one that does not contain a contradiction.[1][2] The lack of contradiction can be defined in either semantic or syntactic terms. The semantic definition states that a theory is consistent if and only if it has a model, i.e. there exists an interpretation under which all formulas in the theory are true. This is the sense used in traditional Aristotelian logic, although in contemporary mathematical logic the term satisfiable is used instead. The syntactic definition states that a theory is consistent if and only if there is no formula P such that both P and its negation are provable from the axioms of the theory under its associated deductive system.

If these semantic and syntactic definitions are equivalent for a particular deductive logic, the logic is complete. The completeness of the sentential calculus was proved by Paul Bernays in 1918[3] and Emil Post in 1921,[4] while the completeness of predicate calculus was proved by Kurt Gödel in 1930,[5] and consistency proofs for arithmetics restricted with respect to the induction axiom schema were proved by Ackermann (1924), von Neumann (1927) and Herbrand (1931).[6] Stronger logics, such as second-order logic, are not complete.

A consistency proof is a mathematical proof that a particular theory is consistent. The early development of mathematical proof theory was driven by the desire to provide finitary consistency proofs for all of mathematics as part of Hilbert's program. Hilbert's program was strongly impacted by incompleteness theorems, which showed that sufficiently strong proof theories cannot prove their own consistency (provided that they are in fact consistent).

Although consistency can be proved by means of model theory, it is often done in a purely syntactical way, without any need to reference some model of the logic. The cut-elimination (or equivalently the normalization of the underlying calculus if there is one) implies the consistency of the calculus: since there is obviously no cut-free proof of falsity, there is no contradiction in general.

Consistency and completeness in arithmetic and set theory

In theories of arithmetic, such as Peano arithmetic, there is an intricate relationship between the consistency of the theory and its completeness. A theory is complete if, for every formula φ in its language, at least one of φ or ¬ φ is a logical consequence of the theory.

Presburger arithmetic is an axiom system for the natural numbers under addition. It is both consistent and complete.

Gödel's incompleteness theorems show that any sufficiently strong effective theory of arithmetic cannot be both complete and consistent. Gödel's theorem applies to the theories of Peano arithmetic (PA) and Primitive recursive arithmetic (PRA), but not to Presburger arithmetic.

Moreover, Gödel's second incompleteness theorem shows that the consistency of sufficiently strong effective theories of arithmetic can be tested in a particular way. Such a theory is consistent if and only if it does not prove a particular sentence, called the Gödel sentence of the theory, which is a formalized statement of the claim that the theory is indeed consistent. Thus the consistency of a sufficiently strong, effective, consistent theory of arithmetic can never be proven in that system itself. The same result is true for effective theories that can describe a strong enough fragment of arithmetic – including set theories such as Zermelo–Fraenkel set theory. These set theories cannot prove their own Gödel sentences – provided that they are consistent, which is generally believed.

Because consistency of ZF is not provable in ZF, the weaker notion relative consistency is interesting in set theory (and in other sufficiently expressive axiomatic systems). If T is a theory and A is an additional axiom, T + A is said to be consistent relative to T (or simply that A is consistent with T) if it can be proved that if T is consistent then T + A is consistent. If both A and ¬A are consistent with T, then A is said to be independent of T.

First-order logic


⊢ (Turnstyle symbol) in the following context of Mathematical logic, means "provable from". That is, a ⊢ b reads: b is provable from a (in some specified formal system) -- see List of logic symbols) . In other cases, the turnstyle symbol may stand to mean infers; derived from. See: List of mathematical symbols.


A set of formulas \Phi in first-order logic is consistent (written Con\Phi) if and only if there is no formula \phi such that \Phi \vdash \phi and \Phi \vdash \lnot\phi. Otherwise \Phi is inconsistent and is written Inc\Phi.

\Phi is said to be simply consistent if and only if for no formula \phi of \Phi, both \phi and the negation of \phi are theorems of \Phi.

\Phi is said to be absolutely consistent or Post consistent if and only if at least one formula of \Phi is not a theorem of \Phi.

\Phi is said to be maximally consistent if and only if for every formula \phi, if Con (\Phi \cup \phi) then \phi \in \Phi.

\Phi is said to contain witnesses if and only if for every formula of the form \exists x \phi there exists a term t such that (\exists x \phi \to \phi {t \over x}) \in \Phi. See First-order logic.

Basic results

  1. The following are equivalent:
    1. Inc\Phi
    2. For all \phi,\; \Phi \vdash \phi.
  2. Every satisfiable set of formulas is consistent, where a set of formulas \Phi is satisfiable if and only if there exists a model \mathfrak{I} such that \mathfrak{I} \vDash \Phi .
  3. For all \Phi and \phi:
    1. if not \Phi \vdash \phi, then Con\left( \Phi \cup \{\lnot\phi\}\right);
    2. if Con \Phi and \Phi \vdash \phi, then Con \left(\Phi \cup \{\phi\}\right);
    3. if Con \Phi, then Con\left( \Phi \cup \{\phi\}\right) or Con\left( \Phi \cup \{\lnot \phi\}\right).
  4. Let \Phi be a maximally consistent set of formulas and contain witnesses. For all \phi and \psi :
    1. if \Phi \vdash \phi, then \phi \in \Phi,
    2. either \phi \in \Phi or \lnot \phi \in \Phi,
    3. (\phi \or \psi) \in \Phi if and only if \phi \in \Phi or \psi \in \Phi,
    4. if (\phi\to\psi) \in \Phi and \phi \in \Phi , then \psi \in \Phi,
    5. \exists x \phi \in \Phi if and only if there is a term t such that \phi{t \over x}\in\Phi.

Henkin's theorem

Let \Phi be a maximally consistent set of S-formulas containing witnesses.

Define a binary relation \sim on the set of S-terms such that t_0 \sim t_1 if and only if \; t_0 \equiv t_1 \in \Phi; and let \overline t \! denote the equivalence class of terms containing t \!; and let T_{\Phi} := \{ \; \overline t \; |\; t \in T^S \} where T^S \! is the set of terms based on the symbol set S \!.

Define the S-structure \mathfrak T_{\Phi} over T_{\Phi} \! the term-structure corresponding to \Phi by:

  1. for n-ary R \in S, R^{\mathfrak T_{\Phi}} \overline {t_0} \ldots \overline {t_{n-1}} if and only if \; R t_0 \ldots t_{n-1} \in \Phi;
  2. for n-ary f \in S, f^{\mathfrak T_{\Phi}} (\overline {t_0} \ldots \overline {t_{n-1}}) := \overline {f t_0 \ldots t_{n-1}};
  3. for c \in S, c^{\mathfrak T_{\Phi}}:= \overline c.

Let \mathfrak I_{\Phi} := (\mathfrak T_{\Phi},\beta_{\Phi}) be the term interpretation associated with \Phi, where \beta _{\Phi} (x) := \bar x.

For all \phi, \; \mathfrak I_{\Phi} \vDash \phi if and only if \; \phi \in \Phi.

Sketch of proof

There are several things to verify. First, that \sim is an equivalence relation. Then, it needs to be verified that (1), (2), and (3) are well defined. This falls out of the fact that \sim is an equivalence relation and also requires a proof that (1) and (2) are independent of the choice of t_0, \ldots ,t_{n-1} class representatives. Finally, \mathfrak I_{\Phi} \vDash \Phi can be verified by induction on formulas.

See also



External links

  • Chris Mortensen, Stanford Encyclopedia of Philosophy