Proof Theory of Finite-valued Logics Richard Zach September 21, 1993 Technical Report TUW–E185.2–Z.1–93 Technische Universität Wien Institut für Computersprachen E185.2 Abteilung für Anwendungen der formalen Logik Resselgasse 3/1, A–1040 Wien, Austria/Europe Phone: (+43 1) 588 01 x 4088 Fax: (+43 1) 504 1589 Proof Theory of Finite-valued Logics Diplomarbeit ausgeführt am Institut für Algebra und Diskrete Mathematik der Technischen Universität Wien unter Anleitung von Univ.-Doz. Dr. phil. Matthias Baaz durch Richard Zach Ada-Christen-Gasse 2/C/3, A-1100 Wien Wien, am 21. September 1993 Preface Many-valued logic is not much younger than the whole field of symbolic logic. It was introduced in the early twenties of this century by Lukasiewicz [1920] and Post [1921] and has since developed into a very large area of research. Most of the early work done has concentrated on problems of axiomatizability on the one hand, and algebraical/model theoretic investigations on the other. The proof theory of many-valued systems has not been investigated to any comparable extent. Proof theory requires appropriate formalisms, such as sequent calculus, natural deduction, and tableaux for classical (and intuitionistic) logic. Several people have, since the 1950's, proposed ways to generalize such formalisms from the classical to the manyvalued case. One particular method for systematically obtaining calculi for all finite-valued logics was invented independently by several researchers, with slight variations in design and presentation. (Section 3.1 contains a short overview of work done in this area). The main aim of this report is to develop the proof theory of finite-valued first order logics in a general way, and to present some of the more important results in this area. This report is actually a template, from which all results can be specialized to particular logics. This idea has also found its way into the notation: Throughout this report, we use V as denoting a set of m truth values, and 2 and Q as dummies representing n-ary connectives and quantifiers, respectively. Replace V by the set of values true and false, and 2 and Q by the usual logical connectives and quantifiers, and you have a treatise on proof theoretic systems for classical logic. Some of the material presented here has appeared in different forms elsewhere; The main innovations of this report are: the use of signed formula expressions and partial normal forms to provide a unifying framework in which clause translation calculi (Chapter 2), sequent calculi (Chapter 3), natural deduction (Chapter 4) and also tableaux can be represented; the recognition of a duality between two types- "negative" and "positive"-of resolution, sequents and tableaux (see also Baaz et al. [1993b]); bounds for partial normal forms for general (Section 1.6) and induced quantifiers (Section 1.7); and negative resolution (Section 2.6). The cut-elimination theorems extend previous results, the midsequent theorem is new, and Chapters 4 and 5 are entirely new. Material on refinements for resolution or on tableaux systems have purposely not been included. For that, see Baaz and Fermüller [1993] and Hähnle [1993a], respectively. I will refrain from trying to give a thorough motivation for many-valued logic in general, or even an answer to the question of Scott [1976], "Does many-valued logic have any use?" The interested reader will find ample material in introductory texts to the topic, e.g., Rescher [1969], Urquhart [1986], Gottwald [1989], or Bolc and Borowik [1992]. I would, however, like to remark that in recent years manyvalued logic has enjoyed a growth in interest from computer science, in such areas as Artificial Intelligence (see, e.g., Ginsberg [1988]) and circuit verification (see Hähnle and Kernig [1993]). The last chapter, on approximations, also tries to make a step in the direction of using many-valued logic-and the whole machinery of ii preface iii proof theoretic calculi and theorem proving systems-to approximate other logics, such as those originally devised for resoning applications. Acknowledgements. I would like to extend my gratitude to the following people: my family, for their unconditional and continuing support; my teachers Matthias Baaz, Chris Fermüller, Georg Gottlob, and Alexander Leitsch; the Dead Logicians Society, in particular Angelika Gängl-Ehrenwert, Peter Koppensteiner, Arie Ovtrucki, and Helmut Veith; Reiner Hähnle and Michel Parigot for comments on parts of this thesis; Elizabeth Dehning for checking my English; and many others too innumerable to mention. Contents Preface ii 1 Basic Concepts 1 1.1 Languages and Formulas . . . . . . . . . . . . . . . . . . . . . . . . . 1 1.2 Substitutions and Unification . . . . . . . . . . . . . . . . . . . . . . 3 1.3 Semantics of First Order Logics . . . . . . . . . . . . . . . . . . . . . 3 1.4 Signed Formula Expressions . . . . . . . . . . . . . . . . . . . . . . . 6 1.5 Partial Normal Forms . . . . . . . . . . . . . . . . . . . . . . . . . . 8 1.6 Bounds for Partial Normal Forms . . . . . . . . . . . . . . . . . . . . 12 1.7 Induced Quantifiers . . . . . . . . . . . . . . . . . . . . . . . . . . . . 15 2 Resolution 18 2.1 Introduction . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 18 2.2 Clauses and Herbrand Semantics . . . . . . . . . . . . . . . . . . . . 19 2.3 Clause Translation Calculi . . . . . . . . . . . . . . . . . . . . . . . . 20 2.4 Semantic Trees and Herbrand's Theorem . . . . . . . . . . . . . . . . 23 2.5 Soundness and Completeness . . . . . . . . . . . . . . . . . . . . . . 24 2.6 Negative Resolution . . . . . . . . . . . . . . . . . . . . . . . . . . . 27 3 Sequent Calculus 31 3.1 Introduction . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 31 3.2 Semantics of Sequents . . . . . . . . . . . . . . . . . . . . . . . . . . 32 3.3 Construction of Sequent Calculi . . . . . . . . . . . . . . . . . . . . . 34 3.4 Equivalent Formulations of Sequent Calculi . . . . . . . . . . . . . . 39 3.5 The Cut-elimination Theorem for PL . . . . . . . . . . . . . . . . . 42 3.6 The Cut-elimination Theorem for NL . . . . . . . . . . . . . . . . . 47 3.7 Analytical Properties of PL . . . . . . . . . . . . . . . . . . . . . . . 51 3.8 Interpolation . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 52 4 Natural Deduction 56 4.1 Introduction . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 56 4.2 Natural Deduction Systems . . . . . . . . . . . . . . . . . . . . . . . 57 4.3 Normal Derivations . . . . . . . . . . . . . . . . . . . . . . . . . . . . 61 5 Approximating Propositional Logics 64 5.1 Introduction . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . . 64 5.2 Propositional Logics . . . . . . . . . . . . . . . . . . . . . . . . . . . 64 5.3 Singular Approximations . . . . . . . . . . . . . . . . . . . . . . . . . 66 5.4 Sequential Approximations . . . . . . . . . . . . . . . . . . . . . . . 69 Bibliography 74 iv Chapter 1 Basic Concepts 1.1 Languages and Formulas In mathematical logic-and especially in proof theory-the syntactic level is very important. The way in which something is formalized may have many consequences, particularly for the statement of theorems and the technicalities of proofs. Sometimes even the validity of a proof theoretic argument depends essentially on the underlying language. One important point in our notion of (first order) logical syntax is the syntactical distinction between free and bound variables. In this, we follow Takeuti [1987] and, ultimately, Gentzen [1934]. 1.1.1. Definition A first-order logical language L consists of the following symbols: (1) Variables: (a) Free variables: a0, a1, a2, . . . , aj , . . . (j ∈ ω) (b) Bound variables: x0, x1, x2, . . . , aj , . . . (j ∈ ω) (2) Constants: (a) Function symbols of arity i (i ∈ ω): f i0, f i1, f i2, . . . , f ij , . . . (j ∈ ω). Nullary function symbols are called constants. (b) Predicate symbols of arity i (i ∈ ω): P i0, P i1, P i2, . . . , P ij , . . . (j ∈ ω) (3) Logical symbols: (a) Propositional connectives of arity nj : 2n00 , 2 n1 1 , . . . , 2 nr r (b) Quantifiers: Q0, Q1, . . . , Qq (4) Auxiliary symbols: (, ), and , (comma). 1.1.2. Definition Terms and semi-terms are inductively defined as follows: (1) Every individual constant is a term (semi-term). (2) Every free variable (and every bound variable) is a term (semi-term). (3) If fn is a function symbol of arity n, and t1, . . . , tn are terms (semi-terms), then fn(t1, . . . , tn) is a term (semi-term). 1.1.3. Definition Formulas, semi-formulas, and outermost logical symbols are inductively defined as follows: (1) If Pn is a predicate symbol of arity n, and t1, . . . , tn are terms (semi-terms), then Pn(t1, . . . , tn) is a formula (semi-formula). It is called atomic or an atom. It has no outermost logical symbol. 1 1.1. languages and formulas 2 (2) If A1, A2, . . . , An are formulas (semi-formulas) and 2n is a propositional connective of arity n, then 2n(A1, A2, . . . , An) is a formula (semi-formula) with outermost logical symbol 2n. (3) If A is a formula (semi-formula) not containing the bound variable x, a is a free variable and Q is a quantifier, then (Qx)A(x), where A(x) is obtained from A by replacing a by x at every occurrence of a in A, is a formula (semi-formula). Its outermost logical symbol is Q. A formula is called open, if it contains free variables, and closed otherwise. A formula without quantifiers is called quantifier-free. We denote the set of formulas of a language L by Frm(L). Of course, in defining specific logics, we shall introduce specific symbols for the propositional connectives and for the quantifiers. 1.1.4. Example The language L normally used for classical 2-valued logic and also for other m-valued logics contains the following propositional connectives: ¬ (not), ∧ (and), ∨ (or), ⊃ (implies), ≡ (equivalent) (of which ¬ is unary and all others binary), and the following two quantifiers: ∀ (for all) and ∃ (there exists). 1.1.5. Definition The complexity of a formula F , or its degree deg(F ), is inductively defined as follows: (1) If F is atomic, then deg(F ) = 1. (2) If F is of the form 2n(A1, . . . , An), then deg(F ) = 1 + ∑n i=1 deg(Ai). (3) If F is of the form (Qx)A(x), then deg(F ) = 1 + deg ( A(x) ) . In the ensuing chapters, we will talk about the object-level constructs, e.g., formulas, and later also sequences and sets of those, on a meta-level. For instance, a calculus may be given by a set of axiom and rule schemata, and every instance of such a schema is an axiom and an inference, respectively. Hence, we shall need meta-variables for the symbols of a language L: As a notational convention we use lowercase letters from the beginning of the alphabet (a, b, c, . . .) to denote free variables, letters from the middle of the alphabet (f , g, h, . . .) for function symbols and constants, letters from the end of the alphabet (x, y, z, . . .) for bound variables, symbols like 2 for connectives, and Q to denote quantifiers, all possibly indexed by subscripts. Furthermore, A, B, C, . . . will stand for formulas; Γ , ∆, Λ, . . . for sequences and sets of formulas, t and s for terms or semi-terms. We will write A(x) for a semi-formula possibly containing the bound variable x, and A(a) resp. A(t) for the formula obtained from A by replacing every occurrence of the variable x by the free variable a resp. the term t. Sometimes it will be convenient to handle this meta-notation more explicitly, particularly in defining and working with inference rules. Then we use α as a variable for free variables (an eigenvariable), and τ as a variable for terms (a term variable). A formula consisting of some formula variables, eigenvariables and term variables is called a schema. By a pre-instance A′ of a schema A we mean an actual formula from Frm(L) which contains occurrences of the eigenvariables and term variables of A. By an instance A′′ of A we mean a pre-instance A′ of A where the eigenand term variables have been replaced by free variables not occurring in A′, and terms, respectively. For instance, consider the schema A(α, τ) where α is an eigenvariable, and τ is a term variable. P 21 (α, a1) ∧ P 12 (τ) ⊃ P 21 (τ, α) is a pre-instance of A(α, τ), and P 21 (a0, a1) ∧ P 12 (f0(a1)) ⊃ P 21 (f0(a1), a0) is an instance. 1.2. substitutions and unification 3 1.2 Substitutions and Unification Substitutions are of great importance in proof theory and automated theorem proving. The resolution calculus is based on substitutions and the so-called unification computation. Unification means finding a substitution, such that two expressions become equal after the substitution has been applied. In this section, and when we refer to these concepts later on, we will use variable, term, and expression more loosely than defined previously: a variable may be free or bound, a term can refer to a term or a semi-term, and an expression can refer to (semi-)terms, formulas, tuples, sequences and sets of formulas, and similar constructs. 1.2.1. Definition A substitution is a mapping of variables to terms with finite domain. Substitutions are denoted by lowercase greek letters and are written as sets of expressions of the form t/x, where t is a (semi-)term and x is a variable. The application of a substitution σ to an expression E is denoted by juxtaposition. Eσ denotes the expression obtained from E by simultaneously replacing every occurrence of a variable x ∈ dom(σ) in E by its image σ(x). The composition σθ of two substitutions σ and θ is defined as follows: σθ(x) =( σ(x) ) θ for every variable x. 1.2.2. Definition A substitution σ is called a unifier of a set A of expressions, if Aσ is a singleton set. The substitution σ is called a most general unifier of A, if every unifier θ of A is an extension of σ, i.e., there is a substitution λ such that θ = σλ. For finite sets of expressions A, the problem of whether there exists a unifier for A is decidable. The unification algorithm (see Chang and Lee [1973]) calculates the most general unifier if it exists. 1.3 Semantics of First Order Logics The motivation to study many-valued logic does originally arise from semantical analysis of logical statements: One rejects the assumption that a statement must be either true or false: it may have some other "truth value". Different contexts have led to different interpretations for a third truth value, finitely or even infinitely many many truth values. This makes it necessary to introduce a basic set of semantic notations, although we shall almost exclusively deal with syntactical properties of finite-valued logics, and not with their model theories. We refer the interested reader to Rescher [1969] and Zinov'ev [1963] for more philosophical investigations into many-valued logic. A treatise on the model theory of many-valued logic can be found in Rasiowa [1974]. 1.3.1. Definition A matrix L for a language L is given by: (1) a nonempty set of truth values V of cardinality m, (2) a subset V + ⊆ V of designated truth values, (3) an algebra V with domain V of appropriate type: For every n-place connective 2 of L there is an associated truth function 2: V n → V , and (4) for every quantifier Q, an associated truth function Q: ℘(V ) \ {∅} → V The set of truth values of the logics considered here will always be finite, i.e., m is a finite number. The constructions given for calculi for many-valued logics can be extended to logics with an infinite set of truth values, although the constructions 1.3. semantics of first order logics 4 will then, in general, not be effective. For special cases, a meaningful proof theory can still be developed. Throughout this work we will give definitions and theorems relative to some given logic, and we will implicitly refer to this logic's language and matrix. In particular, V will always denote the set of truth values of the given logic, and m its cardinality. We shall also assume that V is equipped with a linear order, and write V = {v1, v2, . . . , vm} for the truth values in their given order. It is this order which we will refer to when we give the definition of a sequent. The semantical structure of a logic is effectively determined by its matrix; we use boldface type to denote a logic (with its associated language) and its matrix. For instance, we use CL to refer to classical two-valued first order logic and Lm to refer to m-valued quantificational Lukasiewicz logic. The intended meaning of a truth function for a propositional connective is analogous to the two-valued case: Given formulas A1, . . . , An, which take the truth values w1, . . . , wn, respectively, the truth value of 2(A1, . . . , An) is given by 2(w1, . . . , wn). A truth function for quantifiers is a mapping from nonempty sets of truth values to truth values: Given a quantified formula (Qx)F (x), such a set of truth values describes the situation where the ground instances of F take exactly the truth values in this set as values under a given interpretation. In other words, for a non-empty set M ⊆ V , (Qx)F (x) takes the truth value Q(M) if, for every truth value v ∈ V , it holds that v ∈ M iff there is a domain element d such that the truth value of F (d) is v (all relative to some interpretation). The set M is called the distribution of F . This generalization of quantifiers dates back at least to Mostowski [1963]. Quantifiers of this type have been dubbed distribution quantifiers by Carnielli [1987b]. In Carnielli [1987a], he has also investigated the problem of quantificational completeness, i.e., the problem of which sets of quantifiers can, together with a set of propositional connectives, express any distribution quantifier. 1.3.2. Example The matrix for classical two-valued logic CL is as follows: (1) The set of truth values V = {f, t} (false, true) (2) The set of designated truth values V + = {t} (3) The truth functions for classical two-valued connectives: ¬ f t t f ∧ f t f f f t f t ∨ f t f f t t t t ⊃ f t f t t t f t (4) The truth functions for the quantifiers ∀ and ∃: ∀ ( {t} ) = t ∀ ( {t, f} ) = ∀ ( {f} ) = f ∃ ( {f} ) = f ∃ ( {t} ) = ∃ ( {t, f} ) = t 1.3.3. Example The matrix for the three-valued Lukasiweicz logic L3 consists of: (1) The set of truth values V = {f, p, t} (false, possible, true) (2) The set of designated truth values V + = {t} (3) The truth functions for the connectives: ¬ f t p p t f ∧ f p t f f f f p f p p t f p t ∨ f p t f f p t p p p t t t t t ⊃ f p t f t t t p p t t t f p t 1.3. semantics of first order logics 5 (4) The truth functions for the quantifiers ∀ and ∃ (generalized ∧ and ∨): ∀ ( {t} ) = t ∃ ( {t} ) = t ∀ ( {t, p} ) = p ∃ ( {t, p} ) = t ∀ ( {t, f} ) = f ∃ ( {t, f} ) = t ∀ ( {t, p, f} ) = f ∃ ( {t, p, f} ) = t ∀ ( {p} ) = p ∃ ( {p} ) = p ∀ ( {p, f} ) = f ∃ ( {p, f} ) = p ∀ ( {f} ) = f ∃ ( {f} ) = f 1.3.4. Definition Let L be a first-order language, and let D be any set. The extended language L(D) is obtained by adjoining to the set of constant symbols of L the set D (D is assumed to be disjoint from L). 1.3.5. Definition A structure M = 〈D,Φ〉 for a language L (an L-structure) consists of the following: (1) A non-empty set D, called the domain (elements of D are called individuals). (2) A mapping Φ that satisfies the following: (a) Each n-ary function symbol f of L(D) is mapped to a function f : Dn → D if n > 0, or to an element of D if n = 0. If d ∈ D, then Φ(d) = d. (b) Each n-ary predicate symbol P of L(D) is mapped to a function P : Dn → V if n > 0, or to and element of V if n = 0. 1.3.6. Definition Let M be an L-structure. An assignment s is a mapping from the free variables of L to individuals. 1.3.7. Definition An interpretation I = 〈M, s〉 is an L-structure M = 〈D,Φ〉 together with an assignment s. 1.3.8. Definition Let I = 〈 〈D,Φ〉, s 〉 be an interpretation. The mapping Φ can be extended in the obvious way to a mapping ΦI from terms to individuals: (1) If t is a free variable, then ΦI(t) = s(t). (2) If t is of the form f(t1, . . . , t2), where f is a function symbol of arity n and t1, . . . , tn are terms, then ΦI(t) = Φ(f) ( ΦI(t1), . . . , ΦI(tn) ) . 1.3.9. Definition Given an interpretation I = 〈M, s〉, we define the valuation valI to be a mapping from formulas F of L(D) to truth values as follows: (1) If F is atomic, i.e., of the form P (t1, . . . , tn), where P is a predicate symbol of arity n and t1, . . . , tn are terms, then valI(F ) = Φ(P ) ( ΦI(t1), . . . , ΦI(tn) ) . (2) If the outermost logical symbol of F is a propositional connective 2 of arity n, i.e., F is of the form 2(F1, . . . , Fn), where F1, . . . , Fn are formulas, then valI(F ) = 2 ( valI(F1), . . . , valI(Fn) ) . (3) If the outermost logical symbol of F is a quantifier Q, i.e., F is of the form (Qx)G(x), then valI(F ) = Q ( ⋃ d∈D valIG(d) ) 1.4. signed formula expressions 6 Note that s maps variables to domain elements which are themselves constant symbols in the extended language. An assignment, if constrained to a finite number of variables, can then be viewed as a ground substitution. For a formula F it obviously does not make a difference if val is applied to F or to Fs. 1.3.10. Definition Let I = 〈M, s〉 be an interpretation. An interpretation I′ = 〈M, s′〉 is called a variant of I modulo a1, . . . , an, in symbols I ∼a1,...,an I′, iff s equals s′ except in the values of the free variables a1, . . . , an. 1.3.11. Proposition Let A be a formula, I = 〈M, s〉 be an interpretation, and s(a) = d ∈ D for the free variable a. Then valI(A) = valI′(A[d/a]), for all I′ ∼a I. 1.3.12. Definition A formula F is called a many-valued tautology iff, for every interpretation I, it holds that valI(F ) ∈ V +. It is called satisfiable iff there is an interpretation I s.t. valI(F ) ∈ V +, and unsatisfiable otherwise. Many-valued tautologies play an important rôle in the area of Hilbert-style axiomatizations of many-valued logics. We will not deal with Hilbert-style axiomatizations nor with many-valued tautologies here, and hence the set V + will not be very important in our considerations. We will, however, point out how to incorporate tautologies into our frameworks, e.g., in terms of provability. 1.4 Signed Formula Expressions It is a well-known and not too surprising fact that the truth of many-valued formulas can be reduced, in a sense, to the truth of formulas in two-valued classical logic. In the words of Rosser and Turquette [1952] (p. 1), two-valued logic is sufficient for the development of many-valued logic. The idea behind this "two-valisation" of many-valued logics is as follows: We introduce m two-valued predicates Av1 , . . . , Avm for every many-valued atom A. These predicates are called signed atomic formulas: Aw expresses that A takes the truth value w. This notation, and its corresponding theory, will serve as a convenient formalism for specifying manyvalued logics. In fact, the truth table specification can easily be translated into a specification by means of signed formula expressions, and it is this specification from which we will be able to obtain calculi for the logics under consideration in a completely systematic and mechanizable way. We will use signed formula expressions (sfes) to give meaning to syntactical constructs such as clauses and sequents, which will all have corresponding sfes of a particular form. Satisfaction, validity, etc. of, e.g., a sequent or a clause will be defined via satisfaction etc. of the corresponding sfes. On the other hand, sfe schemata provide a convenient means for specifying calculi by recurring on the semantics of sfes. Sfes of a certain kind will, e.g., be used to specify sequent calculi rules. 1.4.1. Definition A signed formula is an expression of the form Avi , where vi is a truth value, and A is a first order formula. A signed formula expression is a formula built up from signed formulas using ∧ , ∨ , ¬. A signed formula expression of the form Aw or ¬Aw is called a signed literal. It is called an atomic literal iff A is atomic. In that case, Aw is also called a signed atom. Signed formula expressions are Boolean expressions in the signed formulas, and every interpretation defines a Boolean truth value assignment to the signed formulas, and to the expression as a whole. We call an sfe ∆ valid iff ∆ is true under every interpretation via the induced valuation, and satisfiable iff there is an interpretation 1.4. signed formula expressions 7 s.t. ∆ is true in the induced valuation. Furthermore, assume that ¬ occurs in ∆ only immediately in front of signed formulas, if at all. Then ∆ is called positive iff all its signed literals are of the form Avi (i.e., ¬ does not occur at all), and negative iff they are all of the form ¬Avi . We shall use the following abbreviations: If A1, . . . , An are sfes, then ∨n i=1 Ai stands for A1 ∨ (A2 ∨ . . . (An−1 ∨ An) * * *), and ∧n i=1 for A1 ∧ (A2 ∧ . . . (An−1 ∧ An) * * *). If n = 0, we obtain the empty disjunction and the empty conjunction, respectively. By convention, these represent sfes which are always false or always true, respectively. If they occur in the context of other disjunctions or conjunctions, respectively, they may be deleted in the obvious way. Furthermore, if W ⊆ V , then AW denotes ∨ w∈W A w. The usual conventions for omitting parentheses apply: binary operators associate to the right, and precedence is given in the decreasing order ¬, ∧ , ∨ . Furthermore, it is easy to see that the well-known equivalences for classical logic also hold for sfes. So, e.g., associative, commutative, distributive and de Morgan's laws hold for sfes. 1.4.2. Definition Let I be an interpretation and F be a signed formula expression. We say I satisfies F , in symbols: I |= F , iff (1) F is of the form Aw, where A is an atomic formula, and valI(A) = w; (2) F is of the form ¬A and I 6|= A; (3) F is of the form A ∧ B and I |= A and I |= B; (4) F is of the form A ∨ B and I |= A or I |= B; Furthermore, F is called valid, in symbols |= F , iff every interpretation satisfies F ; it is called satisfiable iff there is an interpretation I s.t. I |= F , and unsatisfiable otherwise. 1.4.3. Proposition Let F be a formula. Then the following are equivalent: (1) F is a tautology. (2) The sfe ∨ w∈V + F w is valid. (3) The sfe ∧ w∈V \V + ¬Fw is valid. (4) The sfe ∨ w∈V \V + F w is unsatisfiable. (5) The sfe ∧ w∈V + ¬Fw is unsatisfiable. 1.4.4. Proposition Let F be a formula. Then the following are equivalent: (1) F is unsatisfiable. (2) The sfe ∨ w∈V + F w is unsatisfiable. (3) The sfe ∧ w∈V \V + ¬Fw is unsatisfiable. (4) The sfe ∨ w∈V \V + F w is valid. (5) The sfe ∧ w∈V + ¬Fw is valid. 1.4.5. Proposition Every sfe ∆ can be transformed to an equivalent positive sfe p(∆) (negative sfe n(∆)). Proof. Use de Morgan's laws to bring ¬ immediately before the signed formulas. Replace every signed literal ¬Avi (Avi) in ∆ by ∨ j 6=i A vj ( ∧ j 6=i ¬Avj ). 1.5. partial normal forms 8 1.4.6. Proposition Every sfe ∆ containing negation only immediately in front of signed formulas can be transformed to an equivalent positive (negative) conjunction of disjunctions of signed formulas (a conjunctive form of ∆). Proof. We prove the proposition by induction on the complexity of ∆: If ∆ is a signed atom Avi , then Avi is a positive and ∧ j 6=i ¬Avj a negative conjunction of disjunctions equivalent to ∆. If ∆ is of the form Ψ ∧ Φ, then by induction hypothesis, Ψ ′ is a conjunctive form of Ψ and Φ′ is a conjunctive form of Φ: Ψ ′ ∧ Φ′ is a conjunctive form of ∆. If ∆ is of the form Ψ ∨ Φ, then let Ψ = ∧ r i=1 ∨ki j=1 Ai,j and Φ = ∧s i′=1 ∨ li′ j=1 Bi′,j . We obtain a conjunctive form of ∆ by: r∧ i=1 s∧ i′=1 ( ki∨ j=1 Ai,j ∨ li′∨ j=1 Bi′,j ) 1.4.7. Proposition I |= ∆ iff I 6|= ¬∆. 1.4.8. Proposition |= ∆ iff ¬∆ is unsatisfiable. 1.5 Partial Normal Forms Partial normal forms were introduced by Rosser and Turquette [1952], p. 53ff. as a convenient means of specifying the truth-functional behaviour of many-valued connectives and quantifiers. For a given connective 2, the i-th partial normal form is a signed formula expression schema Fi equivalent to 2(A1, . . . , An) which contains only the (signed) formulas A1, . . . , An, and is in conjunctive normal form1. Partial normal forms are exhaustive and mutually exclusive, i.e., under any given interpretation exactly one of the m partial normal forms comes out true. They can easily be constructed from the truth tables of a given logic, also for the quantifiers of that logic, and in essence provide us directly with the rules to be used in our calculi. 1.5.1. Definition Let 2 be a propositional connective of arity n. An sfe schema F is called an i-th partial form of 2(A1, . . . , An) if the following hold: (1) The signed formulas in F are of the form Ajw (1 ≤ j ≤ n, w ∈ V ). (2) For every instance F ′ of F and every interpretation I and it holds that I |= F ′ iff valI ( 2(A1, . . . , An) ) = vi. It is called a partial normal form, if it is in conjunctive normal form. An i-th partial normal form for ¬2(A1, . . . , An) is defined as above, with = replaced by 6=. Thus, in other words, an i-th partial normal form for 2(A1, . . . , An) (or ¬2(A1, . . . , An)) is a conjunctive normal form in the Ajw which is equivalent to 2(A1, . . . , An) vi (or ¬2(A1, . . . , An)vi). Note that we are actually talking about schemata (our meta-notation of § 1.1) as objects, rather than using them to specify a class of, e.g., sfes. A partial form is a schema, not its collection of instances. This will be important in later chapters, as partial forms are used to define rules of inference in sequent calculus and natural deduction, which are rule schemata themselves. It should be noted again that from one partial normal form we will obtain a clause translation rule, a sequent calculus introduction rule, and a natural deduction rule (for a connective and a place). In this sense, the pnfs provide a relationship between all these calculi. 1Rosser and Turquette [1952] used disjunctive normal forms. 1.5. partial normal forms 9 1.5.2. Example Positive partial normal forms for the connectives of classical twovalued logic CL: (¬A)f = At (¬A)t = Af (A ∨B)f = Af ∧ Bf (A ∨B)t = At ∨ Bt (A ∧B)f = Af ∨ Bf (A ∧B)t = At ∧ Bt (A ⊃ B)f = At ∧ Bf (A ⊃ B)t = Af ∨ Bt Negative pnfs for their negations are given by: ¬(¬A)f = ¬At ¬(¬A)t = ¬Af ¬(A ∨B)f = ¬Af ∨ ¬Bf ¬(A ∨B)t = ¬At ∧ ¬Bt ¬(A ∧B)f = ¬Af ∧ ¬Bf ¬(A ∧B)t = ¬At ∨ ¬Bt ¬(A ⊃ B)f = ¬At ∨ ¬Bf ¬(A ⊃ B)t = ¬Af ∧ ¬Bt 1.5.3. Example Partial normal forms for the connectives of three-valued Lukasiewicz logic L3 are as follows: (¬A)f = At (¬A)p = Ap (¬A)t = Af (A ∨B)f = Af ∧ Bf (A ∨B)p = (Ap ∨ Bp) ∧ (Af ∨ Ap) ∧ (Bf ∨ Bp) (A ∨B)t = At ∨ Bt (A ∧B)f = Af ∨ Bf (A ∧B)p = (Ap ∨ Bp) ∧ (Ap ∨ At) ∧ (Bp ∨ Bt) (A ∧B)t = At ∧ Bt (A ⊃ B)f = At ∧ Bf (A ⊃ B)p = (Ap ∨ At) ∧ (Ap ∨ Bp) ∧ (At ∨ Bf ) or, (A ⊃ B)p = (Ap ∨ Bp) ∧ (At ∨ Bf ) ∧ (Bf ∨ Bp) (A ⊃ B)t = (Af ∨ Ap ∨ Bt) ∧ (Af ∨ Bp ∨ Bt) Negative pnfs for the negations are: ¬(¬A)f = ¬At ¬(¬A)p = ¬Ap ¬(¬A)t = ¬Af ¬(A ∨B)f = ¬Af ∨ ¬Bf ¬(A ∨B)p = (¬Ap ∨ ¬Bp) ∧ (¬Ap ∨ ¬Bf ) ∧ (¬Af ∨ ¬Bp) ¬(A ∨B)t = ¬At ∧ ¬Bt ¬(A ∧B)f = ¬Af ∧ ¬Bf ¬(A ∧B)p = (¬Ap ∨ ¬Bp) ∧ (¬Ap ∨ ¬Bt) ∧ (¬At ∨ ¬Bp) ¬(A ∧B)t = ¬At ∨ ¬Bt ¬(A ⊃ B)f = ¬At ∨ ¬Bf ¬(A ⊃ B)p = (¬Ap ∨ ¬Bf ) ∧ (¬At ∨ ¬Bp) ¬(A ⊃ B)t = ¬Af ∧ ¬Bt ∧ (¬Ap ∨ ¬Bp) We can immediately extract a partial normal form from the truth table in a way analogous to the method of obtaining complete conjunctive normal forms for two-valued functions: The idea is to look at the n-tuples of values for which 2 does not take the required truth value vi, to describe the negations of these situations and to combine these descriptions conjunctively. More precisely, let I = {(w1, . . . , wn) | 2(w1, . . . , wn) 6= vi}. 1.5. partial normal forms 10 Then 2(A1, . . . , An) vi =∧ (w1,...,wn)∈I ( ¬A1w1 ∨ ¬A2w2 ∨ . . . ∨ ¬Anwn ) is a negative i-th partial normal form for 2. We can obtain a positive partial normal form by replacing each ¬Aw by ∨ u∈V \{w} A u. 1.5.4. Example Consider ∧ in CL for the truth value t: Here, I = {(f, f), (f, t), (t, f)}. The corresponding negative partial normal form is: (A ∧B)t = (¬Af ∨ ¬Bf ) ∧ (¬Af ∨ ¬Bt) ∧ (¬At ∨ ¬Bf ). We obtain a positive normal form (A ∧B)t = (At ∨ Bt) ∧ (At ∨ Bf ) ∧ (Af ∨ Bt). A partial normal form constructed this way can have up to mn conjuncts (if 2 never takes the value vi), but standard methods for minimizing combinational function, such as the Quine-McCluskey procedure, can be used to find minimal positive or negative normal forms. Here, "minimal" is meant with respect to the number of conjuncts and the number of formulas per conjunct. For instance, the partial normal form for A ∧Bt given in Example 1.5.2 is minimal, compare this to the one obtained in Example 1.5.4. Nevertheless, there are connectives which admit no "simple" rule. This will be shown in the next section. An adaptation of the Quine-McCluskey procedure has been implemented and gives minimal positive pnfs in reasonable time for small m (see Baaz et al. [1993]). If a connective always takes the truth value vi, then the i-th partial normal form is the empty conjunction, which is always true. 1.5.5. Definition Let Q be a quantifier. An sfe schema F is called an i-th partial form of (Qx)A(x) if the following hold: (1) The signed atoms in F are among {A(τj)uji | 1 ≤ j ≤ p, 1 ≤ i ≤ m} ∪ {A(αj)wji | 1 ≤ j ≤ q, 1 ≤ i ≤ m} where the αi are eigenvariables and the τi term variables. (2) For every pre-instance F ′ of F and every interpretation I it holds that (a) If for all d1, . . . , dq ∈ D there are e1, . . . , ep ∈ D s.t. I |= F{e1/τ1, . . . , ep/τp, d1/α1, . . . , dq/αq} then valI ( (Qx)A′(x) ) = vi. (b) If for all e1, . . . , eq ∈ D there are d1, . . . , dp ∈ D s.t. I 6|= F{e1/τ1, . . . , ep/τp, d1/α1, . . . , dq/αq} then valI ( (Qx)A′(x) ) 6= vi. where A′ is the instance of A determined by F ′. It is called a partial normal form iff it is in conjunctive normal form. An i-th partial normal form for ¬(Qx)A(x) is defined as above, with = replaced by 6=. 1.5. partial normal forms 11 Term variables and eigenvariables express universal and existential conditions on formulas, in much the same way as the universal and existential quantifiers do in classical logic. It is therefore convenient to introduce quantifiers in sfes as a notational convention. The expression (∀ x)A(x) is then meant to stand for the sfe resulting from A(x) by replacing x with some eigenvariable α, and similarly, (∃ x)A(x) for A(τ). Different quantifiers correspond to different eigenvariables/term variables. It is easy to prove the familiar equivalences (distribution of ∃ over ∨ , of ∀ over ∧ , de Morgan's laws) for these quantifiers, allowing us to work with them just as in classical logic. For instance, if (∀ x)(A(x)∧B(x)) then (∀ x)A(x)∧ (∀ x)B(x). This is seen as follows: The first expression stands for F = A(α)∧ B(α), the second for G = A(α) ∧ B(β). Let I be an interpretation. Now if for each assignment of d ∈ D for α, I |= A(d/α) ∧ B(d/α), then also I |= A(d/α) and I |= A(d/α) for all d ∈ D separately. We may rename one α and obtain I |= A(d/α) for all d ∈ D and I |= A(d/β) for all d ∈ D. But this means that G is true in I. 1.5.6. Example Partial normal forms for the quantifiers of CL are as follows:( (∀x)A(x) )f = A(τ)f ((∀x)A(x))t = A(α)t( (∃x)A(x) )f = A(α)f ((∃x)A(x))t = A(τ)t Negative pnfs for their negations are given by: ¬ ( (∀x)A(x) )f = ¬A(α)f ¬((∀x)A(x))t = ¬A(τ)t ¬ ( (∃x)A(x) )f = ¬A(τ)f ¬((∃x)A(x))t = ¬A(α)t 1.5.7. Example Partial normal forms for the quantifiers of L3 are as follows:( (∀x)A(x) )f = A(τ)f ((∀x)A(x))t = A(α)t( (∀x)A(x) )p = A(τ)p ∧ (A(α)p ∨ A(α)t)( (∃x)A(x) )f = A(α)f ((∃x)A(x))t = A(τ)t( (∃x)A(x) )p = A(τ)p ∧ (A(α)f ∨ A(α)p) Negative pnfs for their negations are given by: ¬ ( (∀x)A(x) )f = ¬A(α)f ¬((∀x)A(x))t = ¬A(τ)t ¬ ( (∀x)A(x) )p = ¬A(τ)p ∧ (¬A(α)p ∨ A(τ)t) ¬ ( (∃x)A(x) )f = ¬A(τ)f ¬((∃x)A(x))t = ¬A(α)t ¬ ( (∃x)A(x) )p = (A(α)p ∨ A(τ)f ) ∧ ¬A(τ)p As in the case of the propositional connectives, truth tables for quantifiers yield partial normal forms. Let DistrI(A(x)) = {valI(A(d)) | d ∈ D}, the so called distribution of A(x) relative to I. The truth table for a quantifier Q maps distributions to truth values. To describe the conditions for (Qx)A(x) to take the truth value vi we proceed analogously to the case of a propositional connective: There, we looked at those combinations of truth values w1, . . . , wn for which 2(w1, . . . , wn) is different from the value vi and described these situations. The conjunctions over the negations of these descriptions yields the required conjunctive partial normal form. Accordingly, let W ⊂ V be a distribution s.t. Q(W ) 6= vi. For any particular interpretation I we have DistrI ( A(x) ) 6= W iff either there is some d ∈ D s.t. I |= A(d)V \W (i.e., DistrI ( A(x) ) contains a truth value not in W ), or there is a truth value w ∈ W s.t. for all e ∈ D, I |= A(e)V \{w} (i.e., there is a truth value in W which is never taken by A(x) under I). The conjunction of these conditions over all W s.t. Q(W ) 6= vi fully describes Q(W ) = vi. Let I = {W ⊆ V \ ∅ | Q(W ) 6= vi} 1.6. bounds for partial normal forms 12 The sfe expressing Q(W ) = vi is as follows∧ W∈I ( (∃ x ∈ D)A(x)V \W ∨ ∨ w∈W (∀ y)A(y)V \{w} ) Note that the quantifiers can be brought to the front in the order required by Definition 1.5.5. The corresponding sfe is as follows:∧ W∈I ( ∨ u∈V \W A(τ)u ∨ ∨ w∈W ∨ u∈V \{w} A(αw) u ) We have an upper bound on the number of conjuncts of 2m − 1: This is the case if Q never takes the value vi. If a quantifier never takes a particular truth value vi, then the i-th partial normal form is the empty conjunction, which is always true. 1.5.8. Example Consider the universal quantifier in CL at place t. We have I ={ {f}, {f, t} } . The corresponding pnf is( (∀x)A(x) )t = (A(τ)t ∨ A(α)t) ∧ (A(αt)f ∨ A(αf )t) The second conjunct says that A(x) is either false for all x or true for all x. The first expresses that there is an element e, such that A(e) is true (or for all elements, A(e) is true), hence the second alternative applies: A(x) is uniformly true. 1.5.9. Remark Rosser and Turquette [1952] use a very general quantifier definition where a quantifier may take several variables and formulas: A quantified formula then has the form (Qx1, . . . , xk) ( A1(x1, . . . , xk), . . . , Al(x1, . . . , xk) ) . The case where k = l = 1 is that of our distribution quantifiers. 1.6 Bounds for Partial Normal Forms In the last section we gave examples of pnfs for n-ary connectives and quantifiers having at most mn and 2m− 1 conjuncts, respectively. In this section we show how this can be improved to pnfs of at most mn−1 and 2m−1 conjuncts, respectively. It is also shown that these bounds are tight. The following proposition occurs as Lemma 1 in Rousseau [1967]. 1.6.1. Proposition Let 2 be an n-ary propositional connective. There is an i-th positive partial normal form for 2(A1, . . . , An) of at most mn−1 conjuncts. Proof. Consider the following disjunctive form:∨ (w1,...,wn−1) ( A1 w1 ∧ . . . ∧ An−1wn−1 ∧ ∨ wn 2(w1,...,wn−1,wn)6=vi An w ) This sfe is equivalent to ¬2(A1, . . . , An)vi : Suppose that Aiwi holds for 1 ≤ i ≤ n) and that 2(w1, . . . , wn−1, wn) 6= vi. But then the first part of the disjunct corresponding to (w1, . . . , wn−1) is true, and the second part is true since Anwn is true. Conversely, assume that the disjunct corresponding to (w1, . . . , wn−1) is true. Then Aiwi holds for 1 ≤ i ≤ n − 1. Furthermore, from the second part of the disjunct, there is a wn s.t. Anwn holds and that 2(w1, . . . , wn−1, wn) 6= vi. But this means that ¬2(A1, . . . , An)vi . 1.6. bounds for partial normal forms 13 By negating the whole sfe we obtain the following i-th partial normal form for 2(A1, . . . , An):∧ (w1,...,wn−1) ( ∨ u 6=w1 A1 u ∨ . . . ∨ ∨ u 6=wn−1 An−1 u ∨ ∨ wn 2(w1,...,wn−1,wn)=vi An wn ) This partial normal form has at most mn−1 conjuncts. Rousseau [1967] also stated that the bound given above is tight, although without detailed proof: 1.6.2. Proposition For every n, there is a propositional connective of arity n s.t. any i-th partial normal form for contains mn−1 conjuncts. Proof. Without loss of generality assume that the set of truth values V = {0, 1, . . . ,m − 1}, and let u = i − 1. Furthermore, let u′ ∈ V be a truth value s.t. u′ 6= u. Define as follows: (i1, . . . , in) = { u′ i1 + * * *+ in ≡ 0(m) u otherwise Obviously, ¬(A1, . . . , An)vi holds iff Aj takes the truth value ij (1 ≤ j ≤ n) and i1 + * * *+ in ≡ 0(m). We show that any given minimal normal form for has mn−1 conjuncts: Assume that F = ∨ (W1,...,Wn) ( ∨ w1∈W1 A1 w1 ∧ . . . ∧ ∨ wn∈Wn An wn ) is equivalent to ¬(A1, . . . , An)u and is minimal in the number of disjuncts. This is the case iff ∧ (W1,...,Wn) ( ∨ w1∈V \W1 A1 w1 ∨ . . . ∨ ∨ wn∈V \Wn An wn ) is a minimal conjunctive normal form for (A1, . . . , An)u. Consider an interpretation I making the form for ¬(*)u true, i.e. there is a tuple (W1, . . . ,Wn) such that the corresponding disjunct is true. We show that |Wj | = 1 for all Wj : Assume W1 ⊇ {k, l} and that valI(A1) = k. Then the interpretation I′ which is equal to I except that valI(A1) = l also satisfies F , and thus ¬(*)u. If valI(Aj) = valI(Aj) = ij for 2 ≤ j ≤ n, then we have that k + i2 + * * *+ in ≡ 0(m) l + i2 + * * *+ in ≡ 0(m) by the definition of . But this means that k− l ≡ 0(m) and hence, since 1 ≤ k, l < m that k = l. Hence |W1| = 1. Similarly, we show |Wj | = 1 for j ≥ 2. Thus F is of the form ∨ (i1,...,in) ( A1 i1 ∧ . . . ∧ Anin ) . Every disjunct corresponds to exactly one tuple (i1, . . . , in) s.t. i1 + * * *+ in ≡ 0(m). There are mn−1 such n-tuples. (In general, there are mn−1 tuples s.t. ∑n j=1 ij ≡ k(m) for any given k. Proof by induction: for n = 1, only (k) satisfies the condition. Assume the statement holds for n. For any choice of in+1 there are mn−1 choices for (i1, . . . , in) s.t. ∑n j=1 ij ≡ k − in+1(m).) 1.6. bounds for partial normal forms 14 1.6.3. Proposition For every quantifier Q there is an i-th partial normal form containing at most 2m−1 conjuncts. Proof. Consider the disjunctive form∨ X⊆V \{w} (∧ u∈X (∃ x)A(x)u ∧ [B] ∧ (∀ x) ∨ u∈X ( A(x)u ∨ [C] )) where B is (∃ x)A(x)v iff Q(X) = vi, and C is A(x)w iff Q(X ∪ {w}) 6= vi. This form is equivalent to ¬(Qx)A(x)vi : Let I be an interpretation satisfying ¬(Qx)A(x)vi . In other words, DistrI(A(x)) = Y s.t. Q(Y ) 6= vi. But then the disjunct corresponding to Y \{w} is true under I, since all the existential conjuncts are true (for every truth value u ∈ Y there is a d ∈ D s.t. A(d)u), and the universal conjunct is true as well (there is no d ∈ D s.t. A(d)u ′ for some u′ 6∈ Y ). Conversely, let I be an interpretation satisfying the disjunct corresponding to X ⊆ V \ {w}. We distinguish cases according to which of B and C are present in this disjunct: If B is present (not present) then X ∪ {w} ⊆ DistrI(A(x)) (X ⊆ DistrI(A(x))), since all existential conjuncts are satisfied. Since the universal conjunct is true as well, there can be no u ∈ DistrI(A(x)) s.t. u 6∈ X (u 6∈ X∪{w} if C is present). Hence, DistrI(A(x)) ⊆ X (DistrI(A(x)) ⊆ X ∪ {w}). Now, if neither C nor D are present, we have DistrI(A(x)) = X, and if both C and D are present, we have DistrI(A(x)) = X ∪{w}. In either case Q(X ∪{w}) 6= vi by the conditions on C and D. If only C is present, then X ⊆ DistrI(A(x)) ⊆ X ∪ {w}. But C alone is present iff Q(Y ) 6= vi for both Y = X and Y = X ∪ {w}. The case where B alone is present is impossible (the the universal conjunct cannot be true because of (∃ x)A(x)w). By negating the whole expression we obtain a conjunctive form∧ X⊆V \{w} (∨ u∈X (∀ x) ∨ u′∈V \{u,w} A(x)u ′ ∨ [B′] ∨ (∃ x) ∨ u∈V \(X∪{w}) A(x)u ∨ [C ′] ) where B′ is (∀ x) ∨ u′∈V \{w} A(x) u′ and C ′ is A(x)w. This conjunctive normal form contains at most 2m−1 conjuncts, the number of subsets of V \ {w}. 1.6.4. Proposition For every n, there is a distribution quantifier R s.t. any i-th partial normal form contains 2m−1 conjuncts. Proof. Let u 6= vi, and let R be defined by R(W ) = { vi |W | ≡ 0(2) u otherwise Let ∧ 〈V1,...,Vr,W 〉 (∀ x) ∨ w∈V1 A(x)w ∨ . . . ∨ (∀ x) ∨ w∈Vr A(x)w ∨ (∃ x) ∨ w∈W A(x)w ) be a minimal i-th partial normal form for (Rx)A(x) (Note that ∃ distributes over ∨ , hence any minimal pnf can be written in the above form). This gives a disjunctive form for ¬(Rx)A(x)vi which is minimal iff the above pnf is minimal and is of the form F = ∨ 〈V1,...,Vr,W 〉 ( (∃ x) ∨ w∈V1 A(x)w ∧ . . . ∧ (∃ x) ∨ w∈Vr A(x)w ∧ (∀ x) ∨ w∈W A(x)w ) 1.7. induced quantifiers 15 where Vj = V \ Vj and W = V \ W . Observe that from minimality it follows that no disjunct is redundant (i.e., for every disjunct there is an interpretation satisfying this disjunct and no other) and that no existential conjunct is redundant (i.e., for no k 6= l it holds that Vk ⊆ Vl). We show the following three properties for this form: (1) ⋃ Vj ⊆ W , (2) W ⊆ ⋃ Vj (and hence, W = ⋃ Vj), (3) |Vj | = 1. (1) If there were a j and a v ∈ Vj with v 6∈ W , then either v is redundant in the sense that the form with v deleted from Vj is also equivalent to ¬(Rx)A(x)vi , or there is some interpretation with domain D where A(d)v holds (d is the only witness for (∃ x) ∨ w∈Vj A(x) w). But then the universally quantified conjunct cannot be true, since v 6∈ W . (2) Assume that there were a w ∈ W with w 6∈ Vj for j = 1, . . . , r. Let I be some interpretation satisfying the disjunct under consideration, and let U = {w1, . . . , wr} be truth values (wj ∈ Vj) s.t. there are d1, . . . , dr ∈ D with A(dj)wj . Without loss of generality we can assume that D = {d1, . . . , dr}. It follows that R(U) 6= vi and hence that |U | ≡ 1(2). Consider the interpretation I′ with D′ = D ∪ {dw} and valI′(A(dw)) = w. Obviously, the above form is satisfied under I′, i.e., R(U∪{w}) 6= vi. But |U ∪ {w}| ≡ 0(2), and this contradicts the definition of R. (3) Assume there were some Vj with Vj = V ′j ∪ {u, v} where u 6= v. By way of contradiction, we show that there are two interpretations I′ and I′′ which both satisfy F , but I′ 6|= (Rx)A(x)vi and I′′ |= (Rx)A(x)vi . Let I be some interpretation satisfying the disjunct under consideration, and let U = {w1, . . . , wr} be truth values (wj ∈ Vj) s.t. there are d1, . . . , dr ∈ D with A(dj) wj . Without loss of generality we can assume that D = {d1, . . . , dr}. Assume wj = v. We have two cases according to whether u ∈ U : If u 6∈ U , then take I′ = I and U ′ = U . Otherwise, let k 6= j be such that A(dk)u and u ∈ Vk. Now there is some w′k in Vk with w ′ k 6∈ Vj (otherwise we would have Vk ⊆ Vj). Let I′ be as follows: D′ is D less all such dk for which A(dk) u holds, plus new domain elements d′k and let valI′(A(d ′ k)) = w ′ k. The distribution of A(x) under I ′ is U ′ = U \ {u}. We have that |U ′| ≡ 1(2). Now consider I′′ with D′′ = D′ ∪ {du} and valI′′(A(du)) = u. I′′ satisfies the same disjunct as I′. The distribution of A(x) under I′′ is U ′′ = U ′ ∪ {u}. But |U ′′| ≡ 0(2), a contradiction. In summary, we have that |Vj | = 1, that W = ⋃r j=1 Vj , and that Vk 6= Vl for k 6= l. We see that no conjunct can describe more than one situation. But there are 2|V |−1 subsets of V of odd cardinality (and thus non-empty), hence there are as many disjuncts. 1.7 Induced Quantifiers There is, however, a class of distribution quantifiers which have a lower complexity of their pnf's than the worst cases in the last section. These are the quantifiers induced by certain connectives, namely those which are idempotent, associative and commutative. For example, the usual existential and universal quantifiers are induced by ∨ and ∧, respectively. 1.7.1. Definition Let 2 be an idempotent, associative and commutative connective. The quantifier Q2 induced by 2 is defined as follows: Q2({w}) = w Q2({w1, . . . , wr}) = 2(w1, 2(w2, . . . , 2(wr−1, wr) * * *) 1.7. induced quantifiers 16 It is easily seen that the conditions of idempotence, associativity and commutativity are sufficient to make the induced quantifier well-defined. In fact, the induced quantifiers can be characterized as follows: 1.7.2. Proposition A quantifier Q is an induced quantifier iff the following hold: (1) Q({w}) = w and (2) Q(U ∪W ) = Q(U ∪ {Q(W )}) for all U,W 6= ∅ with U ∪W ⊆ V . The inducing connective is given by 2(u, w) = Q({u, w}). Proof. If: Define 2 as above. From the definition of 2 it is easy to see that it is idempotent and commutative. Furthermore, we have 2(u, 2(v, w)) = Q({u} ∪ Q{v, w}) = Q({u, v, w}) = = Q(Q({u, v} ∪ {w}) = 2(2(u, v), w). We prove Q(W ) = Q2(W ) by induction on |W |: Q({w}) = {w} = Q2({w}). Now, for the induction step we have Q({w1, . . . , wr+1) = Q({w1} ∪ Q({w2, . . . , wr+1)) = = Q({w1} ∪ Q2({w2, . . . , wr+1)) = = 2(w1, 2(w2, . . . , 2(wr, wr+1) * * *) = Q2(w1, . . . , wr+1). Only if: (1) and (2) immediately follow from Definition 1.7.1. There is a close connection between idempotent, associative and commutative connectives and upper semi-lattices over the set of truth values V : 1.7.3. Proposition Let 2 be an idempotent, associative and commutative connective. The relation v defined by u v w iff 2(u, w) = w is a partial order on V with the property that the least upper bound of {w1, . . . , wk} ⊆ V equals 2(w1,2(w2, . . . ,2(wk−1, wk) * * *). Proof. v is reflexive: 2 is idempotent, 2(u, u) = u, hence u v u for any u ∈ V . v is antisymmetric: If u v w and w v u, then 2(u, w) = w and 2(w, u) = u. But 2 is commutative, hence u = w. v is transitive: If u v v and v v w we have 2(u, v) = v and 2(v, w) = w. Thus 2(u, w) = 2(u, 2(v, w)) = 2(2(u, v), w) = 2(v, w) = w whence u v w. We prove that lub{w1, . . . , wk} = 2(w1, . . . ,2(wk−1, wk)) by induction on k: For k = 1 the equality holds since 2 is idempotent. Assume the statement holds for k > 1. Obviously, 2(w1, . . . ,2(wk, wk+1) * * *) is an upper bound for {w1, . . . , wk+1}. Now let u be an upper bound of {w1, . . . , wk+1}, in other words: 2(wj , u) = u. In particular, u is an upper bound for {w2, . . . , wk+1}, and by induction hypothesis 2(w2, . . . ,2(wk, wk+1) * * *) v u. But then 2(2(w1, . . . ,2(wk, wk+1) * * *), u) = = 2(w1,2(2(w2, . . . ,2(wk, wk+1) * * *), u)) = = 2(w1, u) = u. 1.7. induced quantifiers 17 Conversely, every such order defines an idempotent, associative and commutative connective by 2(u, w) = lub{u, w}. For finite V such upper semi-lattices can be pictured as (non-planar) labeled, rooted trees. Hence, there are mm−1 idempotent, associative and commutative connectives and as many induced quantifiers: Cayley's formula (see Cohen [1978], Theorem 98) gives the number of labeled trees as mm−2, and each of the m vertices can serve as the root. The induced quantifiers are obviously the supremum operators on the respective partial orders. 1.7.4. Proposition Let 2 be an idempotent, associative and commutative connective. There is an i-th partial normal form for Q2 with at most m− 1 conjuncts. Proof. We give a disjunctive normal form for ¬(Q2x)A(x)vi : (Q2x)A(x) does not take the value vi iff either there is some d s.t. A(d) u with u 6v vi holds, or for one of vi's lower neighbours, say w, we have that for all d, valI(A(d)) v w. This is expressed as follows: Let v(w) = {u | u v w} and let w1, . . . , wk be all of vi's lower neighbours. (∃ x) ∨ u∈V \v(vi) A(x)u ∨ k∨ j=1 (∀ x) ∨ u∈v(wj) A(x)u The negation of this yields an i-th pnf for Q2: (∀ x) ∨ u∈v(vi) A(x)u ∧ k∧ j=1 (∃ x) ∨ u∈V \v(wj) A(x)u This pnf has less than or equal to m− 1 conjuncts, except when k = m− 1. But in this situation, the first conjunct is redundant since v(vi) = V . Chapter 2 Resolution 2.1 Introduction With the rise of the digital computer in the 1950's, several logicians became interested in the possibilities of mechanizing proof search or theorem proving by computer. A major breakthrough in this area was the introduction of the resolution calculus for classical logic in Robinson [1965]. The resolution calculus is an elegant formalism, in that there is a single rule of inference, but no logical axioms. It uses a special notation for formulas: clause syntax. This notation, or actually the way one obtains clause syntax from ordinary formulas, hides much of the characteristics of predicate logic. Resolution is arguably the most widely used calculus for automated theorem proving, and also one of the most thoroughly investigated. Departing from the original formulation of Robinson, several refinements have been proposed and implemented. For an overview of resolution-based theorem proving see the classic textbooks of Chang and Lee [1973] and Loveland [1978]. A more recent account is Leitsch [1993]. See also Avron [1993] for a study of connections between resolution, sequent calculus and tableaux. Given its importance and ubiquity in the theorem proving arena, it is not surprising that resolution systems have also been developed for non-classical logics, among them many-valued logics. Probably the first such system was presented by Morgan [1976] for a class of many-valued logics introduced in Rosser and Turquette [1952]. Since then resolution formulations for several (classes of) finite and infinite-valued logics (propositional and first order) have been introduced. For a historical survey of theorem proving formalisms for many-valued logics see Hähnle [1993a]. The work of Stachniak et. al. should be mentioned here in particular, see, e.g., Stachniak and O'Hearn [1990]. A different approach, based on a truth-functional analysis of classical resolution, was proposed by Baaz [1984, 1989, 1992]. The resolution rule in this context takes the form C1 C2 (C1 \D1)σ ∪ (C2 \D2)σ where D1 ⊆ C1, D2 ⊆ C2, the literals in D1 and D2 are signed with truth values w1 and w2, respectively, w1 6= w2, and the atoms in D1 and D2 are unifiable with mgu σ. This is a straightforward generalization of the classical case, where w1 = t and w2 = f and the signs are expressed by absence or presence of ¬, respectively. One advantage of this approach is that several methods of refining the deductive system developed for the classical case can also be generalized to this framework. We shall not give such refinements here, but see Baaz and Fermüller [1992, 1993]. This chapter will be devoted to justifying the above rule. Some familiarity with classical resolution terminology is assumed. 18 2.2. clauses and herbrand semantics 19 2.2 Clauses and Herbrand Semantics 2.2.1. Definition A (many-valued) clause C = {A1w1 , . . . , Anwn} is a finite set of signed atoms. By 2 we denote the empty clause. The atom set at(C) of a clause C is the set of its atomic formulas: at(C) = {A1, . . . , An}. The clause syntax is a convenient notation for automated theorem proving. In the classical case, a clause denotes a disjunction of atoms or negations of atoms, and a set of clauses a conjunction of such disjunctions. Clause sets are implicitly considered as universally closed. By bringing a formula into prenex form and then introducing Skolem functions, we can always obtain a formula in this form which is satisfiable iff the original formula is. In this sense, clause syntax is sufficient to capture full classical first order logic. Similarly, for every set of many-valued clauses there is a corresponding sfe in conjunctive normal form: A clause C = {A1w1 , . . . , Apwp}, where Ai is an atomic formula and wi ∈ V , denotes ∨p i=1 Ai wi . A set of clauses C denotes the conjunction of the disjunctions corresponding to its members. The sfe in conjunctive normal form corresponding to C is uniquely determined up to associativity and commutativity of ∧ and ∨ . When we say, e.g., that C is true in an interpretation, we mean that the corresponding sfe is true. 2.2.2. Definition Let M = 〈D,Φ〉 be a structure. M universally satisfies a set of clauses C iff C is true in 〈M, s〉 for every assignment s,. 2.2.3. Definition The Herbrand universe H(C) of a set of clauses C is the set of all ground terms built up from the function symbols and constants in C, or from the function symbols in C and a new constant a, if C contains no constants. 2.2.4. Definition The Herbrand base A(C) of a set of clauses C is the set of all atomic formulas P (t1, . . . , tn) where P is a predicate symbol in C, and t1, . . . , tn ∈ H(C). 2.2.5. Definition An Herbrand interpretation (H-interpretation) H of a set of clauses C is a set of ground literals {Lφ(L) | L ∈ A(C)}, where φ is a function from A(C) to V . H H-satisfies C, iff for every ground instance C ′ of a clause C ∈ C it holds that C ′ ∩H 6= ∅. If no H-interpretation satisfies C, then C is said to be H-unsatisfiable. If H |= C, then H is called an Herbrand model for C. In general, if A is a set of atoms, then a set of the form {Lφ(L) | L ∈ A} is called an assignment to A. A subset of an assignment is a partial assignment to A. Hence, an H-interpretation is an assignment to A(C). Intuitively, it is clear that every H-interpretation also constitutes a structure in the sense of Definition 1.3.7, and that to every ordinary interpretation there is a corresponding H-interpretation. It is then not surprising that H-satisfiability and satisfiability proper coincide in the sense that a set of clauses is H-satisfiable iff it (considered as an sfe) is satisfiable. Since in the case of classical logic, the metanotation of sfes can be interpreted on the level of the object language, the following results subsume the corresponding results for classical logic as subcases. 2.2.6. Definition Let H be an H-interpretation of a set of clauses C. The structure MH corresponding to H is defined as follows: MH = 〈H(C), Φ〉, with Φ as follows: (1) function symbols: Φ(f) = f where f is the n-ary function mapping the arguments t1, . . . , tn ∈ H(C) to f(t1, . . . , tn) ∈ H(C), or Φ(f) = f if f is a constant symbol. 2.3. clause translation calculi 20 (2) predicate symbols: Φ(P ) = P where P is the function mapping 〈d1, . . . , dn〉 to φ ( P (d1, . . . , dn) ) . 2.2.7. Proposition If H H-satisfies a set of clauses C, then MH universally satisfies C. Proof. Let C be a clause in C and let s be any assignment for MH. This maps free variables to ground terms. C has a ground instance Cs obtained by replacing every variable in C by its image under s. By hypothesis, there is a literal Lw ∈ Cs ∩H. The way MH was constructed ensures that valMH(L) = w. But this means that C is true in 〈MH, s〉. 2.2.8. Definition Let M = 〈D,Φ〉 be a structure, and C a set of clauses. An H-interpretation HM corresponding to M is an H-interpretation satisfying the following condition: (C) If valM(F ′) = v, where F ′ is a ground instance of some formula F , then F ′ v ∈ HM. 2.2.9. Proposition If M = 〈D,Φ〉 universally satisfies C, then any corresponding HM H-satisfies C. Proof. Assume that there were a ground instance C ′ of a clause in C and an H-interpretation HM, s.t. C ′ ∩ HM = ∅. Let C ′ = Cσ, where σ = {t1/a1, . . . , tn/an} and a1, . . . , an are all the variables occurring in C. Let σ′ = {Φ′(t1)/a1, . . . , Φ′(tn)/an}, where Φ′ is Φ extended by a value for the new constant a in case C contained no constant symbols. By (C) we have for all A′i wi ∈ C ′ that valM(A′i) 6= wi (where A′i is an instance of a literal Ai ∈ C), and hence that the assignment σ′ of domain elements to the free variables of C is such that C is not satisfied in 〈M, σ′〉. This contradicts the assumption that M universally satisfies C. It is easy to see that H is an H-interpretation corresponding to MH for a given M. 2.3 Clause Translation Calculi In the two-valued case, there is a direct interpretation of a set of clauses as a firstorder formula, namely as the universal closure of a conjunction of disjunctions, where the literals in the disjunction are exactly the members of the individual clauses. Since any classical first-order formula can be written in this form by familiar equivalences of connectives, distribution rules, and skolemization, resolution becomes applicable as a refutational proof system. To establish the unsatisfiability of a formula, we have to test if the corresponding clause form derives the empty clause. In the many-valued case we also need to find such a translation mechanism in order for resolution to be of any practical value. It is not surprising that the partial normal forms for connectives and quantifiers provide us with such a translation mechanism. The conditions such a translation has to satisfy are soundness and completeness: A signed formula is universally unsatisfiable iff its clause translation is. By the last two propositions, this coincides with H-unsatisfiability on the clause level. The clause translation calculi presented here are language preserving. A different, and more effective way to translate formulas to clause forms is the structure preserving method (see Hähnle [1993b] and Baaz and Fermüller [1993]). 2.3.1. Definition An extended clause is a finite set of signed formulas. 2.3. clause translation calculi 21 Here again, an extended clause is read as the sfe consisting of the disjunction of its members. A set of extended clauses is the conjunction of its members. Every clause is also an extended clause, and every signed formula can be thought of as a unit extended clause. Universal satisfiability is defined just as in the case of ordinary clause sets. 2.3.2. Proposition Let F be a closed formula. (1) The extended clause {Fw | w ∈ W ⊆ V } is satisfiable iff it is universally satisfiable. (2) F is a many-valued tautology iff {Fw | w ∈ V \V +} is universally unsatisfiable. Proof. Immediate from Proposition 1.4.3, since F contains no free variables. Proof. By Proposition 1.4.3 2.3.3. Definition Let ∧ j∈I ∆2:i(j) be a positive i-th partial normal form of 2(A1, . . . , An), where 2 is of arity n, ∆2:i(j) is a disjunction of signed atoms of the form Aiw (1 ≤ i ≤ n), and let ∆′2:i(j) be the clause corresponding to ∆2:i(j) A translation rule for 2 at place i is a schema of the form: C ∪ { C ∪ {2(A1, . . . , An)vi} } C ∪ { C ∪∆′2:i(j) | j ∈ I} } 2:i 2.3.4. Theorem Let C, C, 2(A1, . . . , An)vi and ∆′2:i(j) be as in Definition 2.3.3. Then D = C ∪ { C ∪∆′2:i(j) | j ∈ I} } is satisfiable iff E = C ∪ { C ∪ {2(A1, . . . , An)vi} } is satisfiable. Proof. Let H be an H-interpretation satisfying D (E). MH universally satisfies D (E) by Proposition 2.2.7. Let s be an assignment to the free variables of D(resp., E). 〈MH, s〉 satisfies all clauses in C. If it does not satisfy C, then it has to satisfy all the ∆′2:i(j) (resp., 2(A1, . . . , An) vi). By the definition of a partial normal form for 2, 〈MH, s〉 then satisfies also 2(A1, . . . , An)vi (resp., all of ∆′2:i(j)). Hence MH universally satisfies E (resp., D). By Proposition 2.2.9, E (resp., D) is H-satisfied by H. 2.3.5. Definition Let ∧ j∈I ∆Q:i(j) be a positive i-th partial normal form of (Qx)A(x), where ∆Q:i(j) is a disjunction of signed atoms of the form A(τi) w (1 ≤ i ≤ p) respectively A(αi)w (1 ≤ i ≤ q), and let ∆′Q:i(j) be the clause corresponding to ∆Q:i(j). A translation rule for Q at place i is a schema of the form: E = C ∪ { C ∪ { ( (Qx)A(x) )vi}} D = C ∪ { C ∪∆′′Q:i(j) | j ∈ I} } 2:i where ∆′′Q:i(j) is obtained from ∆ ′ Q:i(j) by (1) replacing term variables τi by terms of the form fi(a1, . . . , ak), where fi are distinct new k-ary function symbols and a1, . . . , ak are all free variables in C ∪ { ( (Qx)A(x) )vi}, and by 2.3. clause translation calculi 22 (2) replacing eigenvariables αi by distinct new free variables bi (not occurring in E). 2.3.6. Theorem Let C, C, ( (Qx)A(x) )vi and ∆′′2:i(j) be as in Definition 2.3.5. Then D = C ∪ { C ∪∆′′Q:i(j) | j ∈ I} } is satisfiable iff E = C ∪ { C ∪ {(Qx)A(x)vi} } is satisfiable. Proof. If: Let H be an H-interpretation satisfying E . MH universally satisfies E by Proposition 2.2.7. Assume, by way of contradiction, that D were not universally satisfiable, i.e., for every structure M′ for D, there is an assignment s′ s.t. 〈M′, s′〉 does not satisfy D. In particular, any structure M′H extending MH by providing an interpretation for the new function symbols cannot universally satisfy D. We show that this would imply that MH does not universally satisfy E : Let M′H be such a structure extending MH. There is an assignment s′ s.t. 〈M′H, s′〉 does not satisfy D. If 〈M′H, s′〉 falsifies a clause in C, then so does 〈MH, s〉 since the new function symbols and variables do not occur in C. So assume that 〈M′H, s′〉 falsifies some ∆′′Q:i(j). That is, for the assignment of fi ( s′(a1), . . . , s′(ak) ) to τi and s′(bi) for αi we have 〈M′H, s′〉 6|= ∆Q:i(j). We know that this property holds for every structure extending MH, i.e., for every possible interpretation of the function symbols and every possible assignment to the variables, there are assignments to the new (eigen-) variables s.t. some ∆Q:i(j) is falsified. By Definition 1.5.5 (2b) of partial form for Q, valMH ( (Qx)A(x) ) 6= vi. Only if: If D is satisfiable, it is also universally satisfiable. Let M be a structure universally satisfying D, and let s be an assignment for M. In 〈M, s〉 either C is true or all of ∆′′Q:i(j) (j ∈ I). In the first case, 〈M, s〉 also satisfies E . Otherwise, let s′ be an assignment with s′ ∼b1,...,bq s. 〈M, s′〉 also satsifies all of ∆′′Q:i(j) (j ∈ I), since the bi do not occur in C. But this means that for all assignments to the eigenvariables (namely, all s′), there are assignments to the term variables (namely, Φ〈M,s′〉 ( fi(a1, . . . , ak) ) ) s.t. the conjunction of the ∆′′Q:i(j) is true in 〈M, s〉. By Definition 1.5.5 (2a) then 〈M, s〉 |= (Qx)A(x)vi . The assignment s was arbitrary, so E is universally satisfiable, and, by Proposition 2.2.9, H-satisfiable. Using the translation rules defined above, any set of extended clauses E can be transformed to a set of clauses C which is H-unsatisfiable iff E is universally unsatisfiable. In case E contains no free variables, this is just the case iff E is unsatisfiable. If we want to prove that a given closed formula F is unsatisfiable, Proposition 2.3.2 gives us a set of extended formulas which is unsatisfiable iff F is. Using the translation, we can apply resolution to obtain a proof. 2.3.7. Example The following are the well known translation rules for classical logic: C ∪ { C ∪ {¬Af} } C ∪ { C ∪ {At} } ¬:f C ∪ {C ∪ {¬At}} C ∪ { C ∪ {Af} } ¬:t C ∪ { C ∪ {A ∧Bf} } C ∪ { C ∪ {Af , Bf} } ∧:f C ∪ {C ∪ {A ∧Bt}} C ∪ { C ∪ {At}, C ∪ {Bt} } ∧:t C ∪ { C ∪ {A ⊃ Bf} } C ∪ { C ∪ {At}, C ∪ {Bf} } ⊃:f C ∪ {C ∪ {A ⊃ Bt}} C ∪ { C ∪ {Af , Bt} } ⊃:t C ∪ { C ∪ {(∀x)A(x)f} } C ∪ { C ∪ {A(f(a1, . . . , ak))f} } ∀:f C ∪ {C ∪ {(∀x)A(x)t}} C ∪ { C ∪ {A(b)t} } ∀:t 2.4. semantic trees and herbrand's theorem 23 2.4 Semantic Trees and Herbrand's Theorem 2.4.1. Definition Let C be a set of clauses, and K ⊆ A(C). Let B(K) = {Avi | A ∈ K, vi ∈ V } be the union of all assignments to K. A semantic tree S for K is a finitary, rooted, downward directed tree with labeled edges s.t. the following hold: (1) Every edge e is labeled by a subset l(e) ⊆ B(K). (2) If e1, . . . , ek are all the edges leaving some node X, then {l(e1), . . . , l(ek)} is H-unsatisfiable. (3) For every branch B in S the following hold: Let B consist of the edges e1, e2, . . . There is exactly one assignment AK to K s.t. AK ∩ ⋃ i≥1 l(ei) = ∅. We say B omits AK . (4) For every assignment AK to K there is a branch B in S which omits AK . By I(X) we denote the union of labels on the unique path from the root to X. I(X) is called the refutation set of X. If B is a branch in S then I(B) = ⋃ X∈B I(X). 2.4.2. Remark Note that the concept of semantic tree as defined here is dual to the usual one (see Kowalski and Hayes [1969]). We see from the above definition that, given a semantic tree S for A(C), every H-interpretation is omitted by a branch in S and every branch in S omits an Hinterpretation. 2.4.3. Example We construct a special binary semantic tree for K were every edge is labeled by a singleton. For this, let A be an atom. By T (A) denote the tree constructed as follows: (1) Stage 0: T (A)0 contains a single node X. (2) Stage k + 1: Let Xs be a leaf of T (A)k, and let V (X) = {w | Aw ∈ I(Xs)} be the set of all truth value indices labeling the edges on the path from X to Xs. Let i < j be minimal indices s.t. vi /∈ V (X), vj /∈ V (X). Append to Xs two nodes Xs,1 and Xs,2, and label the edges leading to them {Avi} and {Avj}, respectively. (3) Stop at stage m− 1. The following holds for T (A): First, all branches of T (A) have length m−1 and the labels on one branch are distinct. Hence, every branch omits exactly one literal. Furthermore, every literal is omitted by some branch in T (A). Proof by induction on m: For m = 2 this is evident. Now consider the tree T (A)m−2: The maximal index on a branch is m − 1. In fact, T (A)m−2 equals the tree T ′(A) for V ′ = V \{vm}. By induction hypothesis, some branch B in T (A)m−2 omits a given literal Au. Whatever the vi and vj are in the construction of the successors of the last node on B, one of them is certainly distinct from u. The respective branch in T (A) omits Au. Now let K be enumerated (without repetitions) by A1, A2, A3, . . . Define T in stages as follows: (1) Stage 1: Write down T (A1). (2) Stage k+1. Replace all leaves of the tree constructed in stage k by T (Ak+1). It is easy to see that T is indeed a semantic tree. 2.4.4. Definition Let C be a clause and S a semantic tree. C fails at a node X in S iff there is a ground instance C ′ of C s.t. C ′ ⊆ I(X). A node X in S falsifies a set of clauses C iff some C ∈ C fails at X. If there is no node above X with this property, then X is called a failure node w.r.t. C. If all successors of a node Y are failure nodes for C, then Y is called an inference node. If every branch in S contains a failure node, then S is said to be closed w.r.t. C. 2.5. soundness and completeness 24 2.4.5. Proposition A set of clauses C is unsatisfiable iff every semantic tree for A(C) is closed w.r.t. C. Proof. If: Let S be a semantic tree for A(C) and let H be an H-interpretation for C. By definition of a semantic tree, H is omitted by some branch B in S. S is closed, so there is a failure node X on B. For a ground instance C ′ of a clause C ∈ C we have C ′ ⊆ I(X). If H would H-satisfy C, then C ′ ∩ H 6= ∅ and consequently I(X) ∩H 6= ∅. But this contradicts the assumption that B omits H. Only if: Let S be a semantic tree for A(C). Every branch B of S omits an H-interpretation H. Since C is H-unsatisfiable, we have that for a ground instance C ′ of some clause C ∈ C we have C ′ ∩H = ∅. From this it follows that C ′ ⊆ I(B) (since B omits H, and hence B(A(C)) \H = I(B)), and C fails at some node X on B. So S is closed w.r.t. C. 2.4.6. Proposition The tree R(S) resulting from pruning a closed semantic tree S at all its failure nodes is finite. Proof. Assume R(S) were infinite. By König's Lemma, it would then contain an infinite branch. This contradicts the assumption that S is closed. 2.4.7. Theorem (Herbrand's Theorem) A set of clauses C is unsatisfiable iff there is a finite unsatisfiable set of ground instances of clauses in C. Proof. If: This follows immediately from the definition of H-satisfaction. Only if: By Propositions 2.4.5 and 2.4.6, there is a closed semantic tree R, all of whose (finitely many) failure nodes Xi falsify ground instances C ′i of clauses in C. If X1, . . . , Xn are all failure nodes of R, then {C ′1, . . . , C ′n} is a finite unsatisfiable set of ground instances of clauses in C. 2.5 Soundness and Completeness 2.5.1. Definition Let C1 and C2 be clauses s.t. var(C1) ∩ var(C2) = ∅, and let D1 ⊆ C1, D2 ⊆ C2 be such that every literal in D1 is signed with a truth value v, every literal in D2 with a truth value w, and v 6= w. If A = at(D1 ∪D2) is unifiable with mgu σ, then 〈D1, D2, A〉 is called a key triple with mgu σ for C1, C2. Resolution of two clauses C ′1 and C ′ 2 requires that they be variable disjoint, i.e., that var(C ′1)∩var(C ′2) = ∅. We call C1 and C2 variable disjoint renamings of C ′1 and C ′2, if C1 = C ′ 1λ1 and C2 = C ′ 2λ2, where λi leaves the variables in var(Ci) \ var(Cj) intact, and maps var(Ci)∩var(C2) to different new variables not in ran(λj) (i = 1, 2; j = 2, 1). 2.5.2. Definition (Resolvent) Let C ′1 and C ′ 2 be clauses, let C1 and C2 be variable disjoint renamings, respectively, and let 〈D1, D2, A〉 be a key triple with mgu σ for C1, C2. Then the clause (C1 \D1)σ ∪ (C2 \D2)σ is a resolvent of C ′1, C ′2. 2.5.3. Definition A resolution deduction of a clause D from a set of clauses C is a finite sequence of clauses R1, R2, . . . , Rs = D s.t. either Ri ∈ C or Ri is a resolvent of clauses Rj , Rj′ , j, j′ < i. In that case we write C ` D. 2.5.4. Lemma Let R be a resolvent of C1 and C2. If a model H satisfies C1 and C2, then it also satisfies R. 2.5. soundness and completeness 25 Proof. Clearly, H satisfies C1 and C2 iff it satisfies their variable disjoint renamings. W.l.o.g. we can therefore assume that var(C1) ∩ var(C2) = ∅ and hence var(R) ⊆ var(C1)∪var(C2). Assume, by way of contradiction, that H would not satisfy R, i.e., that Rλ∩H = ∅ for some ground substitution λ. If a literal Lvσλ occurs in Rλ, then Lv occurs in C1 or C2 (σ is the mgu the resolution). The substitution σλ can be extended to a ground substitution η with ran(η) = var(C1)∪ var(C2) s.t. η | var(R) = σλ. Now (C1 \D1)η∩H = ∅. D1η is a singleton set, say, {Lv}, and since H satisfies C1 we have that {Lv} ∩ H 6= ∅. In other words, Lv ∈ H. Similarly, we obtain D2η = {Lw} ⊆ H. But v 6= w and this contradicts the definition of Herbrand model. 2.5.5. Theorem (Soundness) If C ` 2 then C is unsatisfiable. Proof. If C were satisfiable, then a model of C would also be a model of the resolvents R1, . . . , Rs in the corresponding resolution deduction by Lemma 2.5.4. But Rs = 2 and 2 is unsatisfiable. 2.5.6. Theorem (Lifting Lemma) Let C ′1 and C ′ 2 be ground instances of clauses C1 and C2, respectively, and let R ′ be a resolvent of C ′1 and C ′ 2. Then there is a resolvent R of C1 and C2 s.t. R ′ is a ground instance of R. Proof. Let C ′1 = C1λ1 and C ′ 2 = C2λ2, and 〈D′1, D′2, A〉 be the key triple for R with D′1 ⊆ C ′1 and D′2 ⊆ C ′2. D′1 and D′2 are ground, so the corresponding mgu is the empty substitution, D′1 = {L′ v} and D′2 = {L′ w}. Let D1 and D2 be maximal subsets of C1 and C2, respectively, s.t. D1λ1 = D′1 and D2λ2 = D ′ 2. Since λ1λ2 is a unifier of at(D1 ∪D2), we have a key triple 〈D1, D2, at(D1 ∪D2)〉 with mgu σ. Since σ is more general than λ1λ2 we have that R′ is an instance of R = (C1 \D1)σ ∪ (C2 \D2)σ. 2.5.7. Remark Note that the proof of the lifting lemma given in Chang and Lee [1973] is flawed, as has been pointed out by Leitsch [1989]. This is also the reason why we have followed Robinson's [1965] original formulation of resolution with key triples and implicit factoring as the starting point for the generalization to the many-valued case. 2.5.8. Theorem (Completeness) If C is unsatisfiable then C ` 2. Proof. By Proposition 2.4.5, the semantic tree T for A(C) of Example 2.4.3 is closed w.r.t. C, and hence (by Proposition 2.4.6), the corresponding reduced tree R(T ) is finite. We prove, by induction on the number of nodes in R(T ), that there is a resolution deduction of 2 from C. If R(T ) has only one node, then 2 must be in C, since no other clause is falsified at the root of a semantic tree. Otherwise R(T ) contains a node X, s.t. its immediate successors X1, X2 are failure nodes. The edges leaving X are labeled by {Au1} and {Au2}, u1 6= u2. There are ground instances C ′1 and C ′2 of clauses C1 and C2 in C s.t. I(X1) falsifies C ′1 and I(X2) falsifies C ′2, but I(X) does not falsify either . Therefore, C ′1 must contain a literal A u1 and C ′2 must contain a literal A u2 . Now 〈{Au1}, {Au2}, {A}〉 is a key triple for C ′1 and C ′2 with mgu id . We obtain the resolvent R′ = (C ′1\{Au1})∪(C ′2\{Au2}). R′ fails at X, since otherwise there would be some literal L ∈ R′ with L /∈ I(X). Assume L ∈ C ′1 \ {Au1}. Then L /∈ I(X1), contradicting the assumption that C ′1 fails at X1. Similarly for L ∈ C ′2 \ {Au2}. By the lifting lemma, there is a resolvent R of Ci and Cj , s.t. R′ is an instance of R. T is also closed w.r.t. C ∪ {R}, and the tree R(T )′ obtained from R(T ) by 2.5. soundness and completeness 26 pruning at the first node that falsifies R′ (at worst, this is X) is a proper subtree of R(T ) containing all failure nodes. By induction hypothesis, there is a resolution deduction of 2 from C ∪ {R}. 2.5.9. Proposition Any resolution deduction R from C can be rewritten as a resolution deduction R′ in tree form where no variable renaming is necessary in the resolution steps, possibly using renamed copies (variants) of clauses in C. Proof. By induction on the length l of R = C1, . . . , Cl: l = 1: C1 ∈ C. This proof is in tree form, R′ = R. l > 1: Cl is the resolvent of clauses Ci, Cj where i, j < l. The deductions of Ci and Cj from C have length < l. By induction hypothesis, they have tree like deductions R1, R2. Then the deduction R′ .... R1 C1 .... R2 C2 Cl is in tree form. If we take vaiable-disjoint renamings of all topmost clauses and carry the renamings through to the conclusion, then all resolved clauses are variable-disjoint. Being in tree form means that every clause occurence in R is used only once in a resolution step. R can therefore be written as a tree, where the nodes are the clauses in R. This tree has a well-defined height. We call resolution proofs of the form given in the preceding proposition regular. If R contains only ground clauses, then the corresponding proof in tree form uses only literal copies of the clauses in R. 2.5.10. Definition The cumulative substitution ρ(R) of a regular resolution proof is defined by induction on the height h of R: h = 1: ρ(R) = id . h > 1: R ends in a resolution of the form .... R1 C1 .... R2 C2 C with mgu σ. C1 and C2 are variable-disjoint. ρ(R) = ( ρ(R1)ρ(R2) ) σ = ( ρ(R1) ∪ ρ(R2) ) σ. Regular resolution proofs have the property that they can be grounded, as the following lemma shows. This property will be used in the next section. 2.5.11. Lemma Let R be a regular resolution deduction of C from C. There is a substitution γ which grounds R. That is, γ is s.t. Rγ contains only ground clauses and Rγ is a resolution deduction of Cγ from Cγ. Proof. Let var(R) be the set of all variables occurring in R. We define a partition of var(R) = V1 ∪ . . . ∪ Vp by the equivalence relation x ∼ y ⇐⇒ xρ(R) = yρ(R). Let c1, . . . , cp be p new, distinct, constant symbols, and λ the substitution mapping members of Vi to ci. Then γ = ρ(R)λ grounds R, if Rρ(R) is a deduction of Cρ(R) from Cρ(R). But this follows from the definition of ρ(R). 2.5.12. Remark The so-called ground projection is a standard notion in the resolution literature. Cumulative substitutions exist for all (not only regular) resolution proofs which have ground projections. See Baaz and Leitsch [1992] for results relating proof complexity and ground projections. Their Theorem 3.6 also shows that the assumption of regularity in Lemma 2.5.11 is necessary. 2.6. negative resolution 27 2.6 Negative Resolution 2.6.1. Definition A negative clause C is a set of negations of signed atoms. The positive translation C+ of a set of negative clauses C is the set of clauses C obtained by replacing every negative clause {¬A1w1 , . . . ,¬Anwn} by ⋃n j=1{Aj vi | vi 6= wj}. If D = C+, then we define C = D−. The sfe corresponding to a negative clause and universal satisfiability are defined analogous to the positive case. It is obvious that a set of negative clauses is satisfiable iff its positive translation is. For this notion of clause the appropriate notion of resolution is as follows: 2.6.2. Definition Let C1, C2, . . . , Cm be negative clauses s.t. ⋂m i=1 var(Ci) = ∅, and let ¬Di ⊆ Ci be such that every literal in ¬Di is signed with the truth value vi, If A = at( ⋃m i=1 ¬Di) is unifiable with mgu σ, then 〈¬D1, . . . ,¬Dm, A〉 is called an negative key tuple (n-key tuple) with mgu σ for C1, C2, . . . , Cm. 2.6.3. Definition Let C ′1, C ′ 2, . . . , C ′ m be clauses, let C1, C2, . . . , Cm be variable disjoint renamings, respectively, and let 〈D1, . . . , Dm, A〉 be an n-key tuple with mgu σ for C1, . . . , Cm. Then the clause (C1 \D1)σ∪ . . .∪ (Cm \Dm)σ is a resolvent of C ′1, . . . , C ′ m. Negative resolution will play a part in the cut-elimination theorem for negative sequent calculus in the next chapter. We need the following theorem: 2.6.4. Theorem (Soundness) If C ` 2 then C is unsatisfiable. Proof. We prove that for every interpretation H satisfying C, if C `n C then H satisfies C. By induction on the height h of the derivation tree of C from C: If h = 1, then C ∈ C and hence is satisfied by every interpretation which satisfies C. If h > 1, then let C1 C2 . . . Cm C be the last resolution in the derivation, and let 〈¬D1, . . . ,¬Dm, A〉 be its n-key tuple with mgu σ. Since D1, . . . , Dm are unifiable, we can simulate the step by a sequence of positive resolutions on the positive translations: Let ¬Ci = ¬Ei ∪¬Di, and let Ci, Ēi and Di be the positive translations of ¬Ci, ¬Ei and ¬Di, respectively. By Dvi we denote those elements of Di with sign v. We show that there is a positive resolution derivation of C from {C1, . . . , Cm} or, equivalently, a deduction of 2 from {D1, . . . , Dm}. In general, call a set of clauses F = {F1, . . . , Fl+1} contractible iff F is unifiable, for every Fi there is a set Wi = {w1, . . . , wm−l} of truth values s.t. wj ∈ Wi iff F wj i is empty, and for every 1 ≤ i ≤ l + 1 there are truth values wi, wi ∈ Wi (not necessarily distinct) s.t. for all j > i, wi /∈ Wj and for all k < i, wi /∈ Wk. It is easy to see that {D1, . . . , Dm} is contractible with l = m − 1, Wi = {vi}, wi = wi = vi. We show by induction on l that 2 is derivable by positive resolution from a contractible set F : l = 0: This means F = {2}. (Note that {2} is contractible.) l > 0: We have F = {F1, . . . , Fl+1}. Let 1 ≤ i ≤ l and consider Fi and Fi+1. F wi+1 i and F wi i+1 are nonempty. Hence we have that Fi and Fi+1 have a resolvent Gi with key triple 〈Fwi+1i , F wi i+1, at(F wi+1 i ∪F ui i+1)〉. The set of clauses G = {G1, . . . , Gl} is contractible: For Gi, the corresponding set W ′i is Wi ∪ Wi+1 having m − l + 1 elements, u′i = ui and w′i = wi+1. 2.6. negative resolution 28 The following diagram illustrates the case of m = 4. l m− l ¬A1 ¬A2 ¬A3 ¬A4 3 1 A2, A3, A4 A1, A3, A4 A1, A2, A4 A1, A2, A3 1 2 3 4 2 2 A3, A4 A1, A4 A1, A2 1, 2 2, 3 3, 4 1 3 A4 A1 1, 2, 3 2, 3, 4 0 4 2 1, 2, 3, 4 Note how the lists of signed A's form a tree-like resolution deduction of the empty clause. 2.6.5. Lemma (Lifting Lemma) Let C ′1, C ′ 2, . . . , C ′ m be ground instances of clauses C1, C2, . . . , Cm, respectively, and let C ′ be a negative resolvent of C ′1, C ′ 2, . . . , C ′ m. Then there is a negative resolvent C of C1, C2, . . . , Cn s.t. C ′ is a ground instance of C. Proof. Analogous to the positive case. 2.6.6. Lemma Let R′ be a ground proof of C ′ from C′, where C′ is a ground instance of C. Then there is a proof R of C from C, s.t. C ′ is a ground instance of C. Proof. By induction on the height h of R′: h = 1: Then C ′ ∈ C′, and hence R consist only of the clause C. h > 1: R′ ends in a resolution step from the clauses C ′1, . . . , C ′ m. If C ′ i is a topmost clause, then it is a ground instance of some clause Ci ∈ C. Otherwise, C ′i is the last clause in a resolution deduction R′i. By induction hypothesis, there is a resolution deduction Ri of Ci from C where C ′i is a ground instance of Ci. From the lifting lemma, we know that there is some C of which C ′ is a ground instance, s.t. C is a resolvent of C1, . . . , Cm. 2.6.7. Theorem (Completeness) If C is unsatisfiable then C ` 2. Proof. Since C+ is unsatisfiable, there is a regular resolution deduction R of 2 from C+ of length l. By Lemma 2.5.11, R can be grounded, resulting in a ground proof R′ of 2 from a ground instance C′+. We translate this deduction into a negative deduction ν(R′) of 2 by induction on the height h or R′: h = 1: Then 2 ∈ C. h > 1: We proceed in three stages: (1) First, we mark those literals throughout the proof, which result in the last literals resolved upon (in a sense to be made more precise). This results in a proof where some groups of literals in the initial clauses are marked. Every such group is the positive translation of one negative literal in the original set of negative clauses. (2) The proof is pruned by omitting each of these groups from the initial clauses, one at a time. This results in several proofs of 2 of height less than h from subclauses of the initial clauses (all of which are positive translations of subclauses of the original negative clauses). (3) The induction hypothesis is applied to these proofs, yielding negative proofs of 2 from subclauses of the original negative clauses. (4) We add the negative translations of the groups that were left out in step (2), obtaining several negative proofs of clauses of the form {¬P v}. The set of all these clauses is inconsistent, since there is a subproof of R′ of 2 from their positive translations. Hence there is one such 2.6. negative resolution 29 clause for each v ∈ V , and furthermore, P is of the same form in all of these clauses. (5) By resolving, we obtain a negative ground proof of 2 from ground instances of C. By Lemma 2.6.6, there is a negative resolution proof of 2 from C. (1) Let R′ be as follows: .... R ′ 1 {Avi} .... R ′ 2 {Avj} 2 We introduce an indexing system of literals in R′. The initial clauses C′+ are the positive translations of negative clauses in C′. We provide every literal in C′ by an index s.t. no two occurrences of literals in C′ have the same index. If ¬Bw ∈ C ′ ∈ C′ and the index of ¬Bw is x, then we write the indexed literal as ¬xBw. The indexes are extended to literals in the positive translation of C: If∨ v∈V \w B v is the positive translation of ¬xBw, then all the Bv are indexed by x, too. So, if {¬x1A w1 1 , . . . ,¬xnAwnn } ∈ C′, then {x1Av1 | v ∈ V \ w1} ∪ . . . ∪ {xnAvn | v ∈ V \ wn} ∈ C′+ This way, all literals in R′ are marked. Define the operator u(S) on indexed ground proofs S inductively as follows: If S is only an initial clause, then u(S) = S. Otherwise, S ends in a resolution step: .... S1 D1 ∪ {xAw1} .... S2 D2 ∪ {yAw2} D1 ∪D2 let D′i ∪ {x′Avi} be the last clause of u(Si) (i = 1, 2). Then u(S) is .... u(S1) D′1 ∪ {x′Aw1} .... u(S2)[x/y] D′2[x/y] ∪ {x′Aw2} D′1 ∪D′2[x/y] where [x/y] means that the index y is replaced everywhere by the index x. Furthermore, we introduce an operator d(S, C, X) which is defined for an indexed ground proof u(S), a set of initial clauses C and a set of indices X as follows: h = 1: u(S) is only the initial clause D: if D ∈ C, then d(S, X) = D∗, where D∗ is the clause obtained from D by omitting every literal indexed by some index x ∈ X. Otherwise d(S, D, X) = D. h > 1: u(S) is of the form .... S1 D1 ∪ {yAw1} .... S2 D2 ∪ {yAw2} D1 ∪D2 where C = D1∪D2. Let D∗i be the last clause in d(Si, C, X) (i = 1, 2). If yAw1 /∈ D∗1 , then d(S, C, X) = d(S1, C, X). If yAw2 /∈ D∗2 , then d(S, C, X) = d(S2, C, X). Otherwise, let d(S, X) be .... d(S1, C, X) D∗1 .... d(S2, C, X) D∗2 D∗1 \ {yAw1} ∪D∗2 \ {yAw2} Intuitively, u specifies a relation among literal occurrences in the indexed proof S where all literals are either related by having the same negative literal as their 2.6. negative resolution 30 common ancestor, or are related to literals which are resolved upon in the proof. The operator s then is the pruning operator, parametric on a set of inices X. Given an indexed proof S, d(S, C, X) is the proof resulting from S by deleting all literals indexed by members of X from the initial clauses C and all resolutions acting on such literals. (2) We have indexed the proofs R′1 and R ′ 2, resulting in indexed proofs of {xA vi} and {xAvj}, respectively. Let C1,1, . . . , C1,p and C2,1, . . . , C2q be the initial clauses of R1 and R2 respectively. By applying d, we obtain proofs d(R′1, {C1,i}, {x}) (1 ≤ i ≤ p) and d(R′2, {C2,j}, {x}) (1 ≤ j ≤ q) with 2 as their last clause. (3) These proofs are of length < h, hence the induction hypothesis applies. This yields negative resolution proofs N1,i and N2,j of 2 from initial clauses among {D−1,i}i and {D−2,j}j , respectively. In Nr,s, the initial clause D − r,k is the negative translation of Cr,k if k 6= s and the negative translation of d(Cr,k, Cr,k, {x}) otherwise. (4) By adding the negative literals ¬Lvr,sr,s omitted from clause Cr,s in Nr,s again we obtain proofs N ′r,s of {¬L vr,s r,s }. Now consider the original proof R′. Its pruning d(R′, C, I \ {x}) is a proof of 2 from the positive translations of the clauses in N = { {¬Lvr,sr,s } | r, s } . This ensures that all Lr,s are actually equal (i.e., unifiably with unifier id). Furthermore, the set of all vr,s is the whole set of truth values V . For if there were some w ∈ V which is not among the vr,s, the structure which makes Lr,s take the value w satisfies N+. But this is impossible, since N is unsatisfiable. (5) Thus, we obtain a negative resolution proof R′− of 2 from C′. By Lemma 2.6.6, there is a negative resolution proof of 2 from C. Chapter 3 Sequent Calculus 3.1 Introduction Sequent calculus was introduced by Gentzen [1934] for classical and intuitionistic logic and has since proved to be an important formalism for proof theoretic studies of formal systems. Gentzen's sequents are expressions of the form Γ → ∆, where Γ and ∆ are finite sequences of formulas. The full calculus consists of axioms of the form A → A, rules for introducing propositional connectives resp. quantifiers in the left and right sides of the sequent, as well as structural rules (among them the cut rule). The sequent notation can be interpreted in two different ways: firstly, as expressing entailment. A sequent A → B stands for: A entails B, or more generally: Γ → ∆ stands for: the conjunction of the formulas in Γ entails the disjunction of the formulas in ∆. Secondly, the sequent may be interpreted truth-functionally: either one of the formulas in Γ is false or one of the formulas in ∆ is true. For classical logic, these two interpretations coincide, since the truth-functional characterization is equivalent to the derivability of ∧ Γ ⊃ ∨ ∆, and this, by the deduction theorem, to ∧ Γ ` ∨ ∆. For many-valued logics, however, these notions are in general distinct, and we have two fundamental alternatives as to what we mean by "sequent calculus for a many-valued logic": Firstly, we can seek to formalize the entailment relation (or, if we have a complete axiomatization of a given many-valued logic, of derivability) as a sequent calculus. Avron [1992] has developed a general theory of consequence relations for logics, and has also applied this to the study of 3-valued logics in his [1991], where he gives a calculus of hypersequents for Lukasiewicz logic. Here, we will take the second approach via truth-functionality, a property every many-valued logic has. For this, we need a sequent with one place for every truth value in the logic: Γ1 | Γ2 | . . . | Γm, where m is the number of truth values. We interpret this sequent as true under a given interpretation iff one of the Γi contains a formula that takes the value vi. From this point of view, we then see axioms as expressing the fact that every formula must take at least one truth value, and introduction rules as characterizing the truth value of a composite formula in terms of its subformulas. Yet we may also take another, in a sense dual standpoint: A sequent is true under an interpretation, if one of the Γi contains a formula that does not take the value vi. Then, an axiom comes to state the fact that a formula may take at most one truth value. An introduction rule can be viewed as a rule for how to falsify a composite formula by falsifying a subformula. This second interpretation corresponds, in the classical case, exactly to the semantic tableaux of Beth [1955]. We call the first interpretation positive, the second negative. These two perspectives give rise to two different, but closely related, sets of calculi for a many-valued logic-not only two sequent calculi, but also two tableau systems, two natural deduction systems, and two clause translation calculi (for pos31 3.2. semantics of sequents 32 itive and negative resolution, respectively). This correspondence has been described in Baaz et al. [1993a]. This approach has often been used in the literature, and people have reinvented the wheel over and over again. The works differ in notation, and in degrees of generality. Some have used the positive, some the negative approach. The first to introduce a calculus of sequents for many-valued logics (for the propositional case) seems to have been Schröter [1955]. He used the negative approach. His method of axiomatization has also been given treatment in the textbook of Gottwald [1989]. Seemingly independent, Rousseau [1967] treated sequent calculus for first-order finitely-valued logics in full generality and also gave soundness and completeness proofs. In Rousseau [1970] he investigated generalizations of intuitionistic sequent calculus LJ to the many-valued case (with a concept of sequent distinct from the one in his [1967]). Takahashi [1967a] considers a sequent calculus for first-order manyvalued logics with universal and existential quantifiers. Interpolation and definability theorems are proven using this calculus by Miyama [1974], and Hanazawa and Takano [1985]. In Takahashi [1970] a very general Gentzen-style proof theory for continuous logics, i.e., many-valued logics over a topological space as the set of truth values, has been developed. The model theory of continuous logics has been worked out in Chang and Keisler [1966]. Ohya [1967] extended the formalism of Takahashi [1967a] to a system of simple type theory, for which he also shows cut-elimination along the lines of Takahashi [1967b]. After Schröter's work, the negative approach has been taken also by Borowik [1985] for the propositional case only. Earlier, Becchio and Pabion [1977] have outlined a negative sequent calculus for three-valued Lukasiewicz logic. A tableau system for m-valued Lukasiewicz logic was given in Suchoń [1974]. Similar work has been done on equivalent tableau systems starting with Surma [1977] and Carnielli [1987b]. Recently, Carnielli [1991] has presented a notational variant of our negative sequent calculi, based on his work on tableau calculi for many-valued logics in [1987b]. Hähnle [1992, 1993a] presents an in-depth study of many-valued tableaux systems. He introduces several refinements such as sets-as-signs, and investigates classes of many-valued logics (e.g., so-called normal logics) where rules can be efficiently and uniformly represented. Several other papers deal with sequent calculi for specific many-valued logics, in particular Post logic. Kirin [1966] has given sequent calculi based on semantic considerations in Post algebras (see also his [1968]). Similar work was done by Rasiowa [1972], Saloni [1972], and Przymusińska [1980a]. 3.2 Semantics of Sequents 3.2.1. Definition An (m-valued) sequent Γ is an m-tuple Γ1 | . . . | Γmof finite sequences Γi of formulas. If Γ is a sequent, then Γi denotes the i-th component of Γ . If ∆ is a sequence of formulas and I ⊆ M = {1, . . . ,m} (or W ⊆ V ), we denote by [I: ∆] ([W : ∆]) the sequent whose i-th component is ∆ if i ∈ I (vi ∈ W ), and is empty otherwise. For [{i1, . . . , ik}: ∆] ([{w1, . . . , wk}: ∆]) we write [i1, . . . , ik: ∆] ([w1, . . . , wk: ∆]). If Γ and Γ ′ are sequents, then we write Γ, Γ ′ for the sequent Γ1, Γ ′ 1 | . . . | Γm, Γ ′m. 3.2.2. Definition Let I be an interpretation. I p-satisfies (n-satisfies) a sequent Γ iff there is an i (1 ≤ i ≤ m) s.t., for some formula F ∈ Γi, valI(F ) = vi (valI(F ) 6= vi). I is called a p-model (n-model) of Γ and we write I |=p Γ (I |=n Γ ). Γ is called p-satisfiable (n-satisfiable) iff there is an interpretation I s.t. I |=p Γ (I |=n Γ ), and p-valid (n-valid) iff for every interpretation I, I |=p Γ (I |=n Γ ). 3.2. semantics of sequents 33 Evidently, these two notions of satisfiability resp. validity are distinct: an interpretation may be a p-model of a sequent, but not an n-model; a sequent may be p-satisfiable but not n-satisfiable; it may be p-valid but not n-valid (or vice versa). 3.2.3. Example To see how the notion of many-valued sequent is a generalization of the classical case, consider PL with V = {f, t}: A sequent takes the form Γ | ∆ (usually written Γ → ∆). It is p-satisfied if some formula in Γ is false or some formula in ∆ is true. It is n-satisfied if some formula in Γ is true or some formula in ∆ is false. The calculus resulting from p-satisfaction is the calculus given by Gentzen [1934], the one from n-satisfaction corresponds to Beth's 1955 tableaux. If we exchange t and f , then we also obtain Gentzen's system from n-sequents (cf. Carnielli [1991], p. 66). However, in the general case the correspondence is not one of "trivial" duality as in the classical case: you cannot in general pass from the positive to the negative system just by switching the sides of the sequents. We can employ the metaformalism of sfes introduced in Chapter 1 to study the connection between pand n-sequent calculi. The notions of satisfaction and validity of pand n-sequents can be straightforwardly translated by interpreting sequents as sfes of a certain form: A sequent Γ corresponds to the sfe psfe(Γ ) = ∨ A∈Γ1 Av1 ∨ . . . ∨ ∨ A∈Γm Avm modulo p-satisfiability and to nsfe(Γ ) = ∨ A∈Γ1 ¬Av1 ∨ . . . ∨ ∨ A∈Γm ¬Avm modulo n-satisfiability. Note how the first sfe-translation of a sequent corresponds to a positive clause as in Definition 2.2.1, and of the second to a negative clause as in Definition 2.6.1. We can now apply the metatheory of sfes to sequents. P-satisfiability corresponds to reading a sequent as a positive disjunction, and n-satisfiability to reading it as a negative disjunction of signed formulas. Call the translation of a sequent to a disjunction of positive signed formulas a p-sequent, and the translation to a negative disjunction an n-sequent. In light of Proposition 1.4.6, any sfe can be expressed as a conjunction of por n-sequents, and a conjunction of p-sequents (n-sequents) can be transformed to an equivalent conjunction of n-sequents (p-sequents) by Proposition 1.4.5. The negation of a p-sequent (n-sequent) is equivalent to a conjunction of n-sequents (p-sequents). Proposition 1.4.8 provides us with a relation between validity of a sequent and the unsatisfiability of its negation. For the special case where the sequent contains only one formula i: A we obtain the following: 3.2.4. Proposition A sequent [i: A] is p-unsatisfiable (n-unsatisfiable) iff it is nvalid (p-valid). Proof. The negation of the p-sequent Avi is the n-sequent ¬Avi . The claim follows from Proposition 1.4.8. On the other hand, the n-sequent ¬Avi can also be written as a p-sequent∨ j 6=i A vj , and hence the p-unsatisfiability of [i: A] can be established by proving [V \ {vi}: A] p-valid. Proposition 3.2.4 gives us as an alternative to proving the latter sequent in order to establish the unsatisfiability of [i: A]. Instead, we can use a complete (w.r.t. n-validity) calculus for proving [i: A] n-valid and, hence, punsatisfiable. This is of value especially if there is only one designated truth value: 3.2.5. Proposition Let F be a formula. Then the following are equivalent: 3.3. construction of sequent calculi 34 (1) F is a tautology (2) The sequent [V +: F ] is p-valid (3) The sequents [j: F ], where j ∈ V \ V +, are all n-valid. 3.2.6. Proposition Let F be a formula. Then the following are equivalent: (1) F is a unsatisfiable (2) The sequent [V \ V +: F ] is p-valid (3) The sequents [j: F ], where j ∈ V +, are all n-valid. 3.3 Construction of Sequent Calculi 3.3.1. Definition Let ∧ j∈I ∆2:i(j) be a positive (negative) i-th partial normal form of 2(A1, . . . , An) (¬2(A1, . . . , An)), where 2 is of arity n, ∆2:i(j) is a disjunction of signed atoms (literals) of the form Aiw (¬Aiw), where 1 ≤ i ≤ n, and let ∆′2:i(j) be the sequent corresponding to ∆2:i(j). A p- (n-) introduction rule for 2 at place i is a schema of the form: 〈Γ,∆′2:i(j)〉j∈I Γ, [i: 2(A1, . . . , An)] 2:i 3.3.2. Example Consider the implication in three-valued Lukasiewicz logic L3 given in Example 1.3.3. The partial normal forms from Example 1.5.3 yield the following introduction rules for P L3: Γ | ∆ | Π,A Γ, B | ∆ | Π Γ, A ⊃ B | ∆ | Π ⊃:f Γ | ∆, A | Π,A Γ | ∆, A,B | Π Γ,B | ∆ | Π,A Γ | ∆, A ⊃ B | Π ⊃:p Γ,A | ∆, A | Π,B Γ, A | ∆, B | Π,B Γ | ∆ | Π,A ⊃ B ⊃:t and the following rules for N L3: Γ | ∆ | Π,A Γ, B | ∆ | Π,A Γ,A ⊃ B | ∆ | Π ⊃:f Γ,A | ∆, A | Π Γ | ∆, B | Π,A Γ | ∆, A ⊃ B | Π ⊃:p Γ,A | ∆ | Π Γ | ∆ | Π,B Γ | ∆, A,B | Π Γ | ∆ | Π,A ⊃ B ⊃:t 3.3.3. Definition Let ∧ j∈I ∆Q:i(j) be a positive (negative) i-th partial normal form of (Qx)A(x) (¬(Qx)A(x)), where ∆Q:i(j) is a disjunction of signed atoms (literals) of the form A(α)w or A(τ)w (¬A(α)w or ¬A(τ)w), and let ∆′Q:i(j) be the sequent corresponding to ∆Q:i(j) A p- (n-) introduction rule for Q at place i is a schema of the form:〈 Γ,∆′Q:i(j) 〉 j∈I Γ, [i: (Qx)A(x)] Q:i where the free variables α occurring in the ∆′Q:i(j) satisfy the eigenvariable condition: No α occurs in the lower sequent. 3.3. construction of sequent calculi 35 In an introduction rule as above, the formula-2(A1, . . . , An) or (Qx)A(x)- being introduced is called the main formula, the formulas in ∆f :i(j) are called the auxiliary formulas and the formulas in Γ are called the side formulas. The sequents above the line are called the upper sequents or premises, the sequent below the line is called the lower sequent or the conclusion of this rule. 3.3.4. Example Consider the universal quantifier ∀ in the three-valued Lukasiewicz logic L3. From the pnfs in Example 1.5.7 we obtain the following rules: Γ,A(τ) | ∆ | Π,A Γ, (∀x)A(x) | ∆ | Π ∀:f Γ | ∆ | Π,A(α) Γ | ∆ | Π, (∀x)A(x) ∀:t Γ | ∆, A(α) | Π,A(α) Γ | ∆, A(τ) | Π Γ | ∆, (∀x)A(x) | Π ∀:p 3.3.5. Definition A p-sequent calculus PL for a logic L is given by: (1) axiom schemas of the form: [V : A], (2) for every connective 2 and every truth value vi a p-introduction rule 2:i, (3) for every quantifier Q and every truth value vi a p-introduction rule Q:i, (4) weakening rules for every place i: Γ Γ, [i: A] w:i (5) exchange rules for every place i: Γ, [i: B,A],∆ Γ, [i: A,B],∆ x:i (6) contraction rules for every place i: Γ, [i: A,A] Γ, [i: A] c:i (7) cut rules for every two i 6= j: Γ, [i: A] ∆, [j: A] Γ,∆ cut:ij (2) and (3) are called logical rules, (4)–(7) are called structural rules. 3.3.6. Definition An n-sequent calculus NL for a logic L is given by: (1) axiom schemas of the form: [i, j: A], where i 6= j, (2) for every connective 2 and every truth value vi an n-introduction rule 2:i, (3) for every quantifier Q and every truth value vi an n-introduction rule Q:i, (4) weakening rules for every place i: Γ Γ, [i: A] w:i 3.3. construction of sequent calculi 36 (5) exchange rules for every place i: Γ, [i: B,A],∆ Γ, [i: A,B],∆ x:i (6) contraction rules for every place i: Γ, [i: A,A] Γ, [i: A] c:i (7) the cut rule: Γ1, [1: A] . . . Γm, [m: A] Γ1, . . . , Γm cut (2) and (3) are called logical rules, (4)–(7) are called structural rules. 3.3.7. Definition An upward tree P of sequents is called a proof in a sequent calculus SL iff every leaf is an instance of an axiom in SL, and all other sequents in it are obtained from the ones standing immediately above it by an application of one of the rules of SL. The sequent at the root of P is called its end-sequent. A sequent Γ is called provable in SL (in symbols: `SL Γ ) iff it is the end-sequent of some proof (in SL). In the case of PL (NL), we say that Γ is p-provable (nprovable). 3.3.8. Theorem (Soundness for PL) If a sequent is p-provable, then it is p-valid. Proof. By induction on the length l of a proof of a sequent Π. l = 1: Π is an axiom. Since every formula takes some truth value, the disjunction over all truth values is always true. l > 1: Π is the conclusion of a rule of inference. Let I be an interpretation. The induction hypothesis is that all upper sequents are p-valid, and in particular that each is satisfied by I. We distinguish cases according to the rule J applied last: (1) J is a weakening: I |=p Π since A implies A or B. (2) J is an exchange: I |=p Π since disjunction is commutative. (3) J is a contraction: I |=p Π since A or A implies A. (4) J is a cut: The upper sequents are Γ, [i: A] and Γ ′, [j: A] (i 6= j), and Π = Γ, Γ ′. Now either I |=p Γ , in which case also I |=p Γ, Γ ′, or I |=p [i: A]. But then I 6|=p [j: A] and hence I |=p Γ ′. (5) J is an introduction rule for a propositional connective 2 at place i: The upper sequents are instances Γ,∆′′2:i(j) and Π = Γ, [i: 2(A1, . . . , An)]. If I |=p Γ , then also I |=p Π. Otherwise, I |= ∧ psfe ( ∆′′2:i(j) ) . Hence, by Definitions 3.3.1 and 1.5.1 we have that I |=p [i: 2(A1, . . . , An)]. (6) J is an introduction rule for a quantifier Q at place i: The upper sequents are instances Γ,∆′′Q:i(j), and Π = Γ, [i: (Qx)A(x)]. If I |=p Γ , then also I |=p Π. Otherwise, I |= ∧ psfe ( ∆′′Q:i(j) ) . Note that in ∆′′Q:i(j) (and psfe ( ∆′′Q:i(j) ) ) terms t1, . . . , tp take the place of the term variables and free variables a1, . . . , aq take the place of the eigenvariables. Since a1, . . . , aq do not occur in Γ , we have that for all d1, . . . , dq ∈ D, there are e1, . . . , ep ∈ D- namely ΦI(t1[d1/a1, . . . , dq/aq]), . . . , ΦI(tp[d1/a1, . . . , dq/aq])-such that I |=∧ psfe(∆′′Q:i(j)[d1/a1, . . . , dq/aq, e1/t1, . . . , ep/tp]) (Note that the ei may actually depend on the dj , since the terms may contain eigenvariables). Hence, by Definitions 3.3.3 and 1.5.5, we have that I |=p [i: (Qx)A(x)]. 3.3. construction of sequent calculi 37 3.3.9. Theorem (Soundness for NL) If a sequent is n-provable, then it is n-valid. Proof. The proof is analogous to the case for p-sequent calculus, simply replace |=p by |=n and psfe(*) by nsfe(*) throughout the proof. The cases of axioms and cut rules are handled as follows: Π is an axiom, say [i, j: A]. Let I be an interpretation. If I |= ¬Avi or I |= ¬Avj then I |=n [i, j: A]. Since A can take only one truth value, these are the only possible cases. (4′) J is a cut: The upper sequents are Γi, [i: A] 1 ≤ i ≤ m, and Π = Γ1, . . . , Γm. Now either I |=n Γ1, in which case also I |=n Π. Otherwise I |=n [1: A]. But then I 6|=n [i: A] for all 2 ≤ i ≤ m, hence I |=n Γi and a fortiori I |=n Π. 3.3.10. Theorem (Completeness for PL) If a sequent is p-valid, then it is pprovable without cuts from atomic axioms. Proof. We use the method of reduction trees, due to Schütte [1956] (see also Takeuti [1987], Ch. 1, § 8). We show that every sequent Γ is either provable in the sequent calculus or has a counter-model. Let E be an enumeration of all tuples of terms over L. Call a free variable available at stage k iff it occurs in the tree constructed before stage k (if there is no such variable, pick any and call it available) and new otherwise. A p-tuple t of terms is available for the reduction of a formula F at place i with eigenvariables a1, . . . , aq at stage k on a branch B iff (1) t contains only variables which are available at stage k or which are among a1, . . . , aq, and either (2a) t has not been used at all for a reduction of F at place i on B in a stage before k, or (2b) the pre-instance of the premise lying on B of a reduction of F at place i in a stage before k where t has been used did not contain any term variables. A reduction tree TΓ is an upward, rooted tree of sequents constructed from Γ in stages as follows: Stage 0: Write Γ at the root of the tree. Stage k: If the topmost sequent Γ ′ of a branch contains an atomic formula A s.t. A ∈ ⋂ j∈I Γ ′ j then stop the reduction for this branch. Call a branch open if it does not have this property. Repeat the following reduction steps for every formula F occurring at place i in the topmost sequent Γ ′ of an open branch, which has neither already been reduced at place i on this branch in this stage, nor is the result of a reduction at this stage: (1) F ≡ 2(A′1, . . . , A′n): Replace Γ ′ in the reduction tree by: 〈Γ ′,∆′′2:i(j)〉j∈I Γ ′ where ∆′′2:i(j) is an instance of ∆ ′ 2:i(j) in the rule 2:i introducing F as in Definition 3.3.1, obtained by instantiating A1, . . . , An with A′1, . . . , A ′ n, respectively. 3.3. construction of sequent calculi 38 (2) F ≡ (Qx)A′(x): Let a1, . . . , ap be all eigenvariables and t1, . . . , tq be all term variables in the premises of the rule schema Q:i. Replace Γ ′ in the reduction tree by: 〈 Γ ′,∆′′Q:i(j) 〉 j∈I Γ ′ where ∆′′Q:i(j) is an instance of ∆ ′ Q:i(j) in Q:i introducing F as in Definition 3.3.3, obtained by instantiating A with A′, the eigenvariables α1, . . . , αq with the first q-tuple a1, . . . , aq of new free variables in the enumeration E, and the term variables τ1, . . . , τq with the first (w.r.t. E) p-tuple t1, . . . , tp of terms which is available for the reduction of F at place i with eigenvariables a1, . . . , aq at stage k on B. Observe that F ∈ Γ ′i and thus occurs in all upper sequents. If TΓ is finite, then every topmost sequent contains an atomic formula that occurs at each place in that sequent. A cut-free proof of Γ from axioms containing these formulas is easily constructed by inserting weakenings and exchanges. If TΓ is infinite it has an infinite branch B by König's Lemma. For every atomic formula P (t1, . . . , tn) in B, there is an l, 1 ≤ l ≤ m, s.t. P (t1, . . . , tn) never occurs at position l in any sequent on B. We construct an interpretation I as follows: the domain is the set of terms, ΦI(t) = t (t a term), and ΦI ( P (t1, . . . , tn) ) = vl, where vl is the truth value corresponding to the place l. If F is a formula occurring in B, and F occurs at place i anywhere in B, then valI(F ) 6= vi. This is seen by induction on the complexity of F : (1) F is atomic: valI(F ) 6= vi by the construction of I. (2) F ≡ 2(A1, . . . , An): F is reduced somewhere in B. Let ∆′′2:i(j) be the main formulas in the corresponding premises. By induction hypothesis none of the Al in the premise that belongs to B evaluates to the truth value corresponding to its place in the premise. In other words, if Γ ′,∆′′2:i(k) is the premise on B, the induction hypothesis says that I 6|=p ∆′′2:i(k), and hence I 6|= ∧ j psfe ( ∆′′2:i(j) ) . By Definitions 3.3.1 and 1.5.1, I(F ) 6= vi. (3) F ≡ (Qx)A(x): F is reduced somewhere in B. Let ∆′′′Q:i(j) be the pre-instances of the main formulas in the premises, corresponding to ∆′′Q:i(j). In particular, let ∆′′′Q:i(k) be the one on B, and let α1, . . . , αr be the eigenvariables and τ1, . . . , τs be the term variables occurring in it. Now consider the substitution instances of A(x) occurring on B above the lowermost reduction of F : F is reduced in every stage above this one in B. We want to show that for all terms (i.e., domain elements) t1, . . . , tp there are terms t′1, . . . , t ′ q s.t. I 6|=p ∆′′′Q:i(l)[t1/τ1, . . . , tp/τp, t′1/α1, . . . , t′q/αq]. Obviously, we need only consider the termand eigenvariables actually occurring in ∆′′′Q:i(l). If ∆ ′′′ Q:i(k) contains an eigenvariable α, then in ∆ ′′ Q:i(k) this α is replaced by a free variable a, and by induction hypothesis, A(a) does not take the truth value at which it stands (in ∆′′′Q:i(l)). If ∆ ′′′ Q:i(l) contains a term variable τ , we have two cases: (a) There are infinitely many reductions of F on B s.t. the pre-instance of the corresponding premise on B contains a term variable. Hence, B contains infinitely many occurrences of instances of A[t/τ ] where t is any term, and in fact all such instances (This follows from the way available tuples of terms are chosen). None of these instances take the truth values corresponding to the place at which they stand (in ∆′′′Q:i(j)). Together with what has been said above for the eigenvariables, it follows that I 6|=p 3.4. equivalent formulations of sequent calculi 39 ∆′′′Q:i(l)[t1/τ1, . . . , tp/τp, t ′ 1/α1, . . . , t ′ q/αq] and hence that valI(F ) 6= vi by Definition 1.5.5. (b) The reductions of F on B yield only a finite number (including zero) of instances A[t/τj ] on B because there are only finitely many reductions of F on B at place i whose pre-instances of premises on B contain term variables. Since F is reduced infinitely often on B, there is a first reduction where the pre-instance of the corresponding premise on B contains no term variables and such that all reductions above it share this property. Let ∆′′′Q:i(l ′) be the corresponding pre-instance of the premise on B. ∆′′′Q:i(l ′) contains only eigenvariables, and we have that I 6|=p ∆′′Q:i(l′) for the corresponding instance (by induction hypothesis) and hence that valI(F ) 6= vi by Definition 1.5.5. In particular, no formula in Γ evaluates to the truth value corresponding to the position at which it stands. Hence I does not p-satisfy Γ . 3.3.11. Corollary Let A be any formula. [V : A] is provable in PL without cuts from atomic axioms. Proof. Since [V : A] is valid, the construction of T[V :A] terminates and gives a cut-free proof with atomic axioms. 3.3.12. Theorem (Completeness for NL) If a sequent is n-valid, then it is nprovable without cuts from atomic axioms. Proof. The construction of the reduction tree for a sequent Γ is as in the positive case, only that a branch with topmost sequent Γ ′ is closed, if there is a formula A s.t. A ∈ Γ ′i and A ∈ Γ ′j for i 6= j. If TΓ is infinite it has, again, an infinite branch B by König's Lemma. For every atomic formula P (t1, . . . , tn) there is at most one position i where it ever appears in the sequents of B. We construct an interpretation I as follows: the domain is the set of terms, ΦI(t) = t (t a term) and ΦI ( P (t1, . . . , tn) ) = vi, where vi is the truth value corresponding to the position i. If F is a formula occurring in B, and F occurs at place i anywhere in B, then valI(F ) = vi. This is proved as in the positive case, mutatis mutandis. In particular, every formula in Γ evaluates to the truth value corresponding to the position at which it stands. Hence I does not n-satisfy Γ . 3.4 Equivalent Formulations of Sequent Calculi The way we defined sequent calculus is the one closest to Gentzen's original definition. There are several minor modifications to the definitions that can be made, concerning the structural rules and the sequents as sequences of formulas, according to taste and convenience. First of all, it is possible to divide the side formulas of introductions similar to the way the cuts are defined, as follows: 3.4.1. Definition Let f be 2 or Q, let F be 2(A1, . . . , An) or (Qx)A(x), respectively, and let ∆′f :i(j) be as in Definition 3.3.1 or 3.3.3, respectively. Then the following is called a combinational rule:〈 Γj ,∆ ′ f :i(j) 〉 j∈I Γ, [i: F ] f :i 3.4. equivalent formulations of sequent calculi 40 where Γ = 〈Γj | j ∈ I〉. Furthermore, we allow an inference to be made, even if only subsequents of ∆′f :i(j) are present in the premises. The resulting calculus is called the combinational sequent calculus. 3.4.2. Proposition If Γ is provable, then a subsequent of Γ is provable without weakenings. Proof. We inductively translate a proof P of Γ to a proof η(P ) of a subsequent of Γ ′: If Γ is an axiom, we take η(P ) = P . Otherwise, distinguish cases according to the last rule J in P : (1) J is an exchange: Γ is of the form Λ, [i: B,A],Π. By induction hypothesis we have a proof η(P ′) of Λ′,∆,Π ′, where Λ′ and Π ′ are subsequents of Λ and Π, respectively, and ∆ is either empty, [i: A], [i: B], or [i: A,B]. In the former cases, we are done (η(P ) = η(P ′)). In the latter case we get η(P ) as .... η(P ′) Λ′, [i: A,B],Π ′ Λ′, [i: B,A],Π ′ x:i (2) J is a contraction: Γ is of the form Λ, [i: A,A]. By induction hypothesis we have a proof η(P ′) of Λ′,∆, where Λ′ is a subsequent of Λ and ∆ is either empty, [i: A] or [i: A,A]. In the former cases, we are done (η(P ) = η(P ′)). In the latter case we get η(P ) as .... η(P ′) Λ′, [i: A,A] Λ′, [i: A] c:i (3) J is a weakening: Γ is of the form Λ, [i: A]. By induction hypothesis we have a proof η(P ′) of Λ′, where Λ′ is a subsequent of Λ. We take η(P ) = η(P ′), which contains no weakening. (4) J is a cut: Γ is of the form Λ1, . . . , Λk. By induction hypothesis we have proofs η(Pj) of Λ′j ,∆ ′ j , where Λ ′ j is a subsequent of Λj and ∆j is either empty or [ij : A] (In the positive case, k = 2, in the negative case k = m and ij = j). If one of the ∆j is empty, we take η(P ) = η(Pj). Otherwise, we get η(P ) as .... η(P1) Λ′1, [i1: A] . . . .... η(Pk) Λ′k, [ik: A] Λ′1, . . . , Λ ′ j cut (5) J is an introduction of f at place i: Π is of the form Γ, [i: F ] (F is 2(A1, . . . , An) or (∀x)A(x), according to whether f is 2 or Q). By induction hypothesis we have proofs η(Pj) of Γ ′j ,∆ ′′ f :i(j), where Γ ′ j and ∆ ′′ f :i(j) are subsequents of Γj and ∆′f :i(j). If one of the ∆ ′′ f :i(j) is empty, we take η(P ) = η(Pj). Otherwise, we get η(P ) as 〈 .... η(Pj) Γ ′j ,∆ ′′ f :i(j) 〉 j∈I Γ ′, F f :i 3.4. equivalent formulations of sequent calculi 41 Note that by allowing inferences on premises with subsets of the auxiliary formulas present, we can in effect simulate the weakening rule. In the case of Gentzen's LK, this allowance need not be made to obtain the previous result. This is because the rules of LK have the property that every premise contains only one occurrence of an auxiliary formula. By inspection of the proof of the preceding proposition, we see that this is also sufficient (see also Kleene [1952], § 80). In the many-valued sequent calculus, it is in fact possible to give an equivalent formulation with this property, but only for the propositional case. Consider an introduction rule for 2 at place i, where ∆′2:i(j) contains more than one formula, i.e., ∆ ′ 2:i(k) = ∆, [l: Ap]. Then we can split the rule into two as follows:〈 Γj ,∆ ′ f :i(j) 〉 j∈I\k Γk,∆ Γ, [i: 2(A1, . . . , An)] 2:i′ 〈 Γj ,∆ ′ f :i(j) 〉 j∈I\k Γk, [l: Ap] Γ, [i: 2(A1, . . . , An)] 2:i′′ It is easily seen by induction that we can obtain for any propositional rule a set of rules with only one auxiliary formula per premise. The resulting system is called the split-rule calculus. This system is equivalent to the combinational sequent calculus: Every instance of a split rule can be made into an instance of the original rule by means of a series of weakenings. On the other hand, assume we have a deduction in the combinational calculus, 〈 .... Pj Γj ,∆ ′ f :i(j) 〉 j∈I Γ, [i: 2(A1, . . . , An)] 2:i Then we can give a deduction in the split-rule calculus. It suffices to give the case of a rule split into two, the general case is by induction: 〈 .... Pj Γj ,∆ ′ f :i(j) 〉 j∈I\k 〈 .... Pj Γj ,∆ ′ f :i(j) 〉 j∈I\k .... Pk Γk,∆, [l: Ap] Γk, [l: Ap],∆ x Γ, [l: Ap], [i: 2(A1, . . . , An)] 2:i′ Γ, [i: 2(A1, . . . , An)], [l: Ap] x Γ ′, [i: 2(A1, . . . , An),2(A1, . . . , An)] 2:i′′ Γ, [i: 2(A1, . . . , An)] xc Nota bene that the method of splitting rules does in general not work for quantifier rules. Consider for instance the sequent [i, j: A(α)], which might be a premise for a quantifier introduction. It expresses the condition that, for all α, A(α) either takes the value vi or the value vj . The corresponding split sequent, however, says that either for all α, A(α) takes the value vi or, for all α, A(α) takes the value vj . In a word, ∀ does not distribute over ∨ . We could also define a sequent to be a tuple of sets (instead of sequences) of formulas. We call such sequents set-sequents. If sequent calculus formulated with set-sequents, the contraction and exchange rules become superfluous. It should be pointed out that the study of sequent systems where combinational and ordinary rules are combined while some or all structural rules are left out is a difficult and highly interesting topic. Considerations along these lines lead to the investigation of so-called substructural logics, for instance Girard's Linear Logic, or Relevance Logic. For a short survey see Troelstra [1992], § 2.8. 3.5. the cut-elimination theorem for pl 42 3.5 The Cut-elimination Theorem for PL The cut-elimination theorem, or Hauptsatz, was first proved for classical logic in Gentzen [1934]. It is one of the most important theorems in proof theory. Its underlying idea is to eliminate structure from a given proof to extract information from it. This is an idea often used, and cut-elimination corresponds closely to methods used in other calculi to obtain similar results: the 2nd ε-Theorem in Hilbert and Bernays [1939] in ε-calculus, the normalization theorem of Prawitz [1971] in natural deduction, and the normalization results for free deduction obtained by Parigot [1992]. Important consequences of the Hauptsatz for classical (and intuitionistic) logic include the Midsequent Theorem, Herbrand's Theorem, Craig's Interpolation Theorem, and Beth's Definability Lemma (see Takeuti [1987], ch. 1, § 6). The proof of the cut-elimination theorem for the family PL of many-valued sequent calculi is analogous to the proof of the classical case given by Gentzen. It proceeds by reducing cuts to cuts on formulas of smaller complexity and moving the cuts in a given proof upwards until they only involve axioms; such cuts are easily removed. The most important prerequisite for the proof is the ability to transform a cut on a composite formula to a derivation using only cuts acting on subformulas of the original cut-formula. In the classical case, this is established by case distinction; we, however, have to establish this uniformly by the lemmas below. Note how the way Takahashi [1967a] defined his sequent calculus makes similar lemmas unnecessary for his proof of cut-elimination. In the proof of the cut-elimination theorem, we replace the cut rule, for technical reasons, by an equivalent rule called mix: Π Λ Π(i), Λ(j) (A, i, j) where A occurs in Πi and Λj, and Π(i) (Λ(j)) is obtained from Π (Λ) by deleting every occurrence of A in Πi (Λj). A cut can be replaced by a mix followed by a sequence of weakenings, and a mix can be emulated by a series of cuts. Call the calculus obtained from PL by replacing the cut rule with the mix rule PL′. 3.5.1. Lemma Consider the derivation schema 〈Π,∆′2:i(r)〉r Π, [i: 2(A1, . . . , An)] 〈Λ, ∆′2:j(s)〉s Λ, [j: 2(A1, . . . , An)] Π,Λ (2(A1, . . . , An), i, j) where neither Πi nor Λj contains 2(A1, . . . , An). Then there is a proof of Π,Λ from the sequents Π,∆′2:i(r) and Λ, ∆ ′ 2:j(s) using only mixes on A1, . . . , An. Proof. It suffices to show the lemma for Π = Γ = ∅: If we have a proof of the empty sequent from the sequents ∆2:i(r) and ∆2:j(s), then we can immediately obtain (using exchanges and contractions) a proof of Π,Λ by writing Π or Λ left of every sequent in the deduction, as appropriate. The conjunction ∧ r psfe ( ∆′2:i(r) ) ∧ ∧ s psfe ( ∆′2:j(s) ) of the sfes corresponding to the sequents ∆′2:i(r) and ∆ ′ 2:j(s) is clearly unsatisfiable, since ∧ r psfe ( ∆′2:i(r) ) is equivalent to 2(*)vi and ∧ s psfe ( ∆′2:j(s) ) is equivalent to 2(*)vj , where i 6= j. We translate this conjunction to a set of clauses, which is also unsatisfiable: Let C = ⋃ j∈I C2:i(j) be the set of clauses where C2:i(j) = {PAvk | A ∈ ∆′2:i(j)k, 1 ≤ k ≤ m} C is unsatisfiable, since the conjunction of the ∆'s is. By the completeness of many-valued resolution (Theorem 2.5.8), there is a resolution deduction R of 2 3.5. the cut-elimination theorem for pl 43 from C. Note that the literals in C are all ground, and hence, in every resolution step, there is only one literal that is resolved upon. Without loss of generality, assume that the resolution proof is in tree form (cf. Proposition 2.5.9). Such a proof can immediately be translated to a PL′ deduction as follows: Let C1 ∪ {Aλvk} C2 ∪ {Aλvl} C1 ∪ C2 be the last resolution step in R, where A is a of one of A1, . . . , An. The corresponding mix is: ∆1, [k: A, . . . , A] ∆2, [l: A, . . . , A] ∆1,∆2 (A, k, l) where ∆1 and ∆2 are the sequents corresponding to C1 and C2, respectively. 3.5.2. Lemma Consider the derivation schema 〈Π,∆2:i(r)〉r Π, i: (Qx)A(x) 〈Λ, ∆Q:j(s)〉s Λ, j: (Qx)A(x) Π,Λ ((Qx)A(x), i, j) where neither Πi nor Λj contains (Qx)A(x), and the appropriate eigenvariable conditions are satisfied. Then there is a proof of Π,Λ from the sequents Π,∆Q:i(l) and Λ, ∆Q:j(k) using only mixes on instances Aσ of formulas A occurring in the ∆'s, where dom(σ) contains only eigenvariables. Proof. Again, it suffices to show the lemma for Π = Γ = ∅. The conjunction ∧ r psfe ( ∆′Q:i(r) ) ∧ ∧ s psfe ( ∆′Q:j(s) ) of the sfes corresponding to the sequents ∆Q:i(r) and ∆Q:j(s) is universally unsatisfiable (where universal satisfiability for sfes is defined analogous to Definition 2.2.2. We translate this conjunction to a set of clauses, which is also universally unsatisfiable: Let C = ⋃ j∈I CQ:i(j) be the set of clauses where CQ:i(j) = {PA(cτ )vk | A(τ) ∈ ∆′Q:i(j)k, 1 ≤ k ≤ m} ∪ ∪{PA(aτ )vk | A(α) ∈ ∆′Q:i(j)k, 1 ≤ k ≤ m} where the aα are distinct free variables and the cτ are distinct constant symbols. By the completeness of many-valued resolution (Theorem 2.5.8), there is a resolution deduction R of 2 from C, which is (w.l.o.g.) in tree form. We recursively translate Rρ(R) (ρ the cumulative substitution of Definition 2.5.10) into a PL′-derivation as follows: Let .... R1 C1 .... R2 C2 (C1 \ {A})} ∪ (C2 \ {A}} be the last resolution step in Rρ(R), where 〈{A}, {A}, at({A})〉 is a key triple (with mgu id), A contains an atom of the form PA(a) or PA(c) (a is a variable and c a constant symbol). Recall that the cumulative substitution has been applied to R, so resolution steps actually do take this special, ground-like, form. Let P1 (P2) be the result of the recursive translation of R1 (R2). The corresponding deduction in PL′ is: .... P1 ∆ (k) 1 , [k: A] .... P2 ∆ (l) 2 , [l: A] ∆1,∆2 (A, r, s) 3.5. the cut-elimination theorem for pl 44 where ∆1 and ∆2 are the sequents corresponding to C1 and C2, respectively (where free variables are replaced by eigenvariables and constants by term variables). The corresponding σ is easily obtained from ρ(R). Note that here, too, we can add Π and Λ to the sequents without changing the term structure of formulas other than the ones in the ∆'s, since only eigenvariables (which must not occur in Π or Λ) are substituted into. It is clear that σ will in general substitute into eigenvariables. This corresponds to the case in the cut-elimination theorem for LK where a cut on (∀x)A(x) is reduced to a cut on A(t). The eigenvariable above the strong premise of the cut is replaced throughout the whole proof with the term occurring in the premise of the weak ∀-introduction above the right premise of the cut. We are now ready to state and prove the Hauptsatz for PL: 3.5.3. Theorem A PL-proof with end sequent Λ can be transformed into a cut-free PL-proof of Λ. Proof. It suffices to show that mixes can be eliminated from PL′-proofs. If every PL′-proof with only one mix as its last inference can be transformed to a PL′-proof without a mix, then the mixes can be eliminated from all PL′-proofs. This is seen by induction on the number c of mixes in the proof: For c = 1 this is the hypothesis. If c > 1, then consider a mix s.t. the subproof ending in the mix contains only this mix as its last inference. By hypothesis, this mix can be eliminated, and the subproof replaced by a mix-free proof of the same endsequent. The resulting proof has c − 1 mixes, and the induction hypothesis applies. Hence, it suffices to show the following: 3.5.4. Lemma Let P be a PL′-proof containing only one mix (A, i, j) as its last inference. Then P can be transformed to a mix-free PL′-proof P ′ of the same end-sequent. Let P be a PL′-proof containing only one mix (A, i, j) which occurs as the last inference in P . The degree of P , denoted d(P ), is the complexity deg(A) of the mix formula A. We call a thread in P containing the left (right) upper sequent of the mix a left (right) thread. The rank of a left (right) thread is the number of consecutive sequents counting upwards from the left (right) upper sequent of the mix which contain the mix formula at place i (j). The left (right) rank of P , denoted rl(P ) (rr(P )) is the maximum of the ranks of its left (right) threads. The rank of P , denoted r(P ), is the sum of its left and right rank: r(P ) = rl(P ) + rr(P ). The proof is by double induction on the rank and degree of P : (1) r = 2, i.e., left and right rank of P equal 1. We distinguish cases according to the type of the inferences above the mix: (a) Π is an axiom [V : A]: P is of the form [V : A] .... P1 Λ [V \ vi: A], Λ(j) (A, i, j) We can derive Π(i), Λ(j) without a mix as follows: .... P1 Λ [j: A, . . . , A], Λ(j) x [j: A], Λ(j) c [V \ vi: A], Λ(j) w 3.5. the cut-elimination theorem for pl 45 (b) Λ is an axiom. Similarly. (c) Π is the conclusion of a structural inference. Since the left rank is 1, this inference must be a weakening at place i: .... P1 Π(i) Π(i), [i: A] w: i .... P2 Λ Π(i), Λ(j) (A, i, j) where Π is Π ′, [i: A]. We obtain Π(i), Λ(j) without a mix as follows: .... P1 Π(i) Π(i), Λ(j) w (d) Λ is the conclusion of a structural inference. Similarly. (e) Both Π and Λ are conclusions of introduction rules 2:i and 2:j for the connective 2: 〈Π,∆2:i(r)〉r Π, i: 2(* * *) 〈Λ, ∆2:j(s)〉s Λ, j: 2(* * *) Π,Λ (2(* * *), i, j) By Lemma 3.5.1, there is a derivation of Π,Λ from Π,∆2:i(r) and Λ, ∆2:j(s), using only mixes, exchanges, and contractions. Since the formulas in the ∆'s are subformulas of 2(* * *), the degrees of all the resulting mixes are less than the degree of the original mix. We iteratively apply the induction hypothesis to the new mixes and obtain a derivation without mixes of Π,Λ. (f) Both Π and Λ are conclusions of introduction rules Q:i and Q:j for the quantifier Q: 〈Π,∆Q:i(r)〉r Π, i: (Qx)A(x) 〈Λ, ∆Q:j(s)〉s Λ, j: (Qx)A(x) Π,Λ ((Qx)A(x), i, j) By Lemma 3.5.2, there is a derivation of Π,Λ from Π,∆Q:i(r)σ and Λ, ∆Q:j(s)σ, using only mixes, exchanges, and contractions, where σ only substitutes into eigenvariables of the two quantifier introductions. Since the formulas in the ∆'s are subformulas of (Qx)A(x), the degrees of all the resulting mixes are less than the degree of the original mix. We iteratively apply the induction hypothesis to the new mixes and obtain a derivation without mixes of Π,Λ. (2) rr(P ) > 1: Again, we distinguish cases: (a) Λi contains A: We obtain the following mix-free proof: .... P1 Π Π(i), [i: A, . . . , A] e Π(i), [i: A] c Π(i), Λ(j) w 3.5. the cut-elimination theorem for pl 46 (b) Πj contains A: We obtain the following mix-free proof: .... P2 Λ Λ(j), [j: A, . . . , A] e Λ(j), [j: A] c Π(i), Λ(j) w (c) Λ is the consequence of an inference J2, which is either structural (but not cut), or a logical inference not introducing A at place j. P is of the form .... P1 Π .... P (1) Ψ(1) . . . .... P (p) Ψ(p) Λ J2 Π(i), Λ(j) (A, i, j) Let j1, . . . , js, 1 ≤ jk ≤ p, be all indices s.t. Ψ(jk) contains A (There is at least one such jk, otherwise the right rank of P would equal 1). Consider the proofs P (jk)′: .... Π .... Ψ(jk) Π(i), Ψ(jk)(j) (A, i, j) In P (jk)′, rl(P (jk)′) = rl(P ) and rr(P (jk)′) ≤ rr(P ) − 1, and in sum r(P (jk)′) ≤ r(P ) − 1. Hence the induction hypothesis applies and we have mix-free proofs P (jk)′′ of Π(i), Ψ(jk)(j). For indices l not occurring in the above list, we have that Ψ(l) equals Ψ(l)(j), and we define P (l)′′ as .... P (l) Ψ(l) Π(i), Ψ(l) w If J2 is a (w:j) (and consequently, p = 1 and Ψ(1)(j) = Λ(j)), then P (1)′′ serves as our transformed proof. Otherwise, construct a proof as follows: .... P (1) ′′ Π(i), Ψ(1)(j) . . . .... P (p) ′′ Π(i), Ψ(p)(j) Π(i), Λ(j) J2 Π(i), Λ(j) x (d) Λ is the consequence of a logical inference J2 introducing A at place j. P is of the form .... P1 Π .... P (1) Λ, ∆(1) . . . .... P (p) Λ, ∆(p) Λ, [j: A] J2 Π(i), Λ(j) (A, i, j) Consider the proofs P (k)′ (Note that ∆(k) does not contain A-only proper subformulas of A-and hence ∆(k)(j) equals ∆(k)): .... P1 Π .... P (k) Λ, ∆(k) Π(i), Λ(j),∆(k) (A, i, j) 3.6. the cut-elimination theorem for nl 47 In P (k), rl(P (k)′) = rl(P ), rr(P (k)′) ≤ rr(P )−1 and in sum r(P (k)′) ≤ r(P )−1. Hence, the induction hypothesis applies and we obtain mix-free proofs P (k)′′ of Π(i), Λ(j),∆(k). Construct a proof P ′ as follows: .... P1 Π .... P (1) ′′ Π(i), Λ(j),∆(1) . . . .... P (p) ′′ Π(i), Λ(j),∆(p) Π(i), Λ(j), [j: A] J2 Π(i),Π(i), Λ(j) (A, i, j) Note that A does not occur at place j in Π(i), since otherwise case (2)(b) would have applied, hence rr(P ′) = 1. With rl(P ′) = rl(P ) we have that r(P ′) < r(P ) and the induction hypothesis yields a mix-free proof P ′′ of Π(i),Π(i), Λ(j). We obtain a mix-free proof: .... P ′ Π(i),Π(i), Λ(j) Π(i), Λ(j) xc (3) rr(P ) = 1 and rl(P ) > 1: This case is dealt with in the same way as (2) above, mutatis mutandis. This completes the proof of the cut-elimination theorem. 3.6 The Cut-elimination Theorem for NL The cut elimination theorem also holds for the negative sequent calculus. Carnielli [1991] gives a proof of cut-elimination for his (n-)sequent calculus via tableaux. The induction proof itself is again very similar to the proof for the positive case. As above, we first have to ensure that cuts on composite formulas are decomposable. The mix in the negative sequent calculus takes the form Π1 . . . Πm Π (1) 1 , . . . ,Π (m) m (A) where A occurs in Πii (1 ≤ i ≤ m), and Π (i) i is obtained from Πi by deleting every occurrence of A in Πii. Call the calculus obtained from NL by replacing the cut rule with the mix rule NL′. 3.6.1. Lemma Consider the derivation schema 〈Π1,∆2:1(r1)〉r1 Π1, [1: 2(A1, . . . , An)] . . . 〈Πm,∆2:m(rm)〉rm Πm, [m: 2(A1, . . . , An)] Π1, . . . ,Πm ( 2(A1, . . . , An) ) where Πii does not contain 2(A1, . . . , An). Then there is a proof of Π1, . . . ,Πm from the sequents Πi,∆2:i(ri) using only mixes on A1, . . . , An. Proof. Completely analogous to the positive case, using negative resolution deductions. 3.6. the cut-elimination theorem for nl 48 3.6.2. Lemma Consider the derivation schema 〈Π1,∆Q:1(r1)〉r1 Π1, [1: (Qx)A(x)] . . . 〈Πm,∆Q:m(rm)〉rm Πm, [m: (Qx)A(x)] Π1, . . . ,Πm ( (Qx)A(x) ) where Πii does not contain (Qx)A(x), and the appropriate eigenvariable conditions are satisfied. Then there is a proof of Π1, . . . ,Πm from the sequents Πi,∆Q:i(ri) using only mixes on instances Aσ of formulas A occurring in the ∆'s, where dom(σ) contains only eigenvariables. Proof. Completely analogous to the positive case, using negative resolution deductions. 3.6.3. Theorem A NL-proof with end sequent Λ can be transformed into a cut-free NL-proof of Λ. Proof. It suffices to show the following: Let P be a NL′-proof containing only one mix A as its last inference. Then P can be transformed to a mix-free NL′-proof P ′ of the same end-sequent. Let P be a NL′-proof containing only one mix A which occurs as the last inference in P . The degree of P , denoted d(P ), is the complexity of the mix formula A. We call a thread in P containing the i-th upper sequent of the mix an i-th thread. The rank of an i-th thread is the number of consecutive sequents counting upwards from the i-th upper sequent of the mix which contain the mix formula at place i. The i-th rank of P , denoted ri(P ) is the maximum of the ranks of its i-th threads. The rank of P , denoted r(P ), is the sum of its i-th ranks rank: r(P ) = ∑m i=1 ri(P ). The proof is by double induction on the rank and degree of P : (1) r = m, i.e., all i-th ranks of P equal 1. We distinguish cases according to the type of the inferences above the mix: (a) Πi is an axiom [i, j: A]: P is of the form .... P1 Π1 . . . [i, j: A] .... Pm Πm Π (1) 1 , . . . , [j: A], . . . ,Π (m) m (A) We can derive the conclusion without a mix as follows: .... Pj Πj [j: A, . . . , A],Π(j)j x [j: A],Π(j)j c Π (1) 1 , . . . , [j: A], . . . ,Π (m) m wx (b) Πi is the conclusion of a structural inference. Since the i-th rank is 1, this inference must be a weakening at place i: .... P1 Π1 . . . .... Pi Π (i) i Π (i) i , [i: A] w: i . . . .... Pm Πm Π (1) 1 , . . . ,Π (i) i , . . . ,Π (m) m (A) 3.6. the cut-elimination theorem for nl 49 where Πi is Π (i) i , [i: A]. We obtain the conclusion without a mix as follows: .... Pi Π (i) i Π (1) 1 , . . . ,Π (i) i , . . . ,Π (m) m wx (c) All the Πi are conclusions of introduction rules 2:i for the connective 2: 〈Π1,∆2:1(r1)〉r1 Π1, [1: 2(* * *)] . . . 〈Πm,∆2:m(rm)〉rm Πm, [m: 2(* * *)] Π1, . . . ,Πm ( 2(* * *) ) By Lemma 3.6.1, there is a derivation of Π1, . . . ,Πm from Πi,∆2:i(ri) (1 ≤ i ≤ m), using only mixes, exchanges, and contractions. Since the formulas in the ∆'s are subformulas of 2(* * *), the degrees of all the resulting mixes are less than the degree of the original mix. We iteratively apply the induction hypothesis to the new mixes and obtain a derivation without mixes of Π1, . . . ,Πm (d) All the Πi are conclusions of introduction rules Q:i for the quantifier Q: 〈Π1,∆Q:1(r1)〉r1 Π1, [1: (Qx)A(x)] . . . 〈Πm,∆Q:m(rm)〉rm Πm, [m: (Qx)A(x)] Π1, . . . ,Πm ( (Qx)A(x) ) By Lemma 3.6.2, there is a derivation of Π1, . . . ,Πm from Πi,∆Q:i(ri)σ (1 ≤ i ≤ m), using only mixes, exchanges, and contractions, where σ only substitutes into eigenvariables of the two quantifier introductions. Since the formulas in the ∆'s are subformulas of (Qx)A(x), the degrees of all the resulting mixes are less than the degree of the original mix. We iteratively apply the induction hypothesis to the new mixes and obtain a derivation without mixes of Π1, . . . ,Πm. (2) ri(P ) > 1 and rj(P ) = 1 for all j < i: Again, we distinguish cases: (a) Πkl (k 6= l) contains A: We obtain the following mix-free proof: .... Pl Πl Π (l) l , [l: A, . . . , A] e Π (l) l , [l: A] c Π (1) 1 , . . . ,Π (m) m wx (b) Πi is the consequence of an inference J2, which is either structural, or a logical inference not introducing A at place i. P is of the form .... P1 Π1 . . . .... P (1) Ψ(1) . . . .... P (p) Ψ(p) Πi J2 . . . .... Pm Πm Π (1) 1 , . . . ,Π (m) m A 3.6. the cut-elimination theorem for nl 50 Let j1, . . . , js, 1 ≤ jk ≤ p, be all indices s.t. Ψ(jk) contains A (There is at least one such jk, otherwise the i-th rank of P would equal 1). Consider the proofs P (jk)′: .... Π1 . . . .... Ψ(jk) . . . .... Πm Π (1) 1 , . . . , Ψ(jk) (i), . . . ,Π (m) m A In P (jk)′, rj(P (jk)′) = rj(P ) (j 6= i) and ri(P (jk)′) ≤ ri(P ) − 1, and in sum r(P (jk)′) ≤ r(P ) − 1. Hence the induction hypothesis applies and we have mix-free proofs P (jk)′′ of Π (1) 1 , . . . , Ψ(jk) (i), . . . ,Π (m) m . For indices l not occurring in the above list, we have that Ψ(l) equals Ψ(l)(i), and we define P (l)′′ as .... P (l) Ψ(l) Π (1) 1 , . . . , Ψ(l), . . . ,Π (m) m w If J2 is a (w:i) (and consequently, p = 1 and Ψ(1)(i) = Π (i) i ), then P (1) ′′ serves as our transformed proof. Otherwise, construct a proof as follows: .... P (1) ′′ Π (1) 1 , . . . , Ψ(1) (i), . . . ,Π (m) m Π (1) 1 , . . . ,Π (m) m , Ψ(1)(i) x . . . .... P (p) ′′ Π (1) 1 , . . . , Ψ(p) (i), . . . ,Π (m) m Π (1) 1 , . . . ,Π (m) m , Ψ(p)(i) x Π (1) 1 , . . . ,Π (m) m ,Π (i) i J2 Π (1) 1 , . . . ,Π (m) m x (c) Πi is the consequence of a logical inference J2 introducing A at place i. P is of the form .... P1 Π1 . . . .... P (1) Πi,∆(1) . . . .... P (p) Πi,∆(p) Πi, [i: A] J2 . . . .... Pm Πm Π (1) 1 , . . . ,Π (m) m (A) Consider the proofs P (k)′ (Note that ∆(k) does not contain A-only proper subformulas of A-and hence ∆(k)(i) equals ∆(k)): .... P1 Π1 . . . .... P (k) Πi,∆(k) . . . .... Pm Πm Π (1) 1 , . . . ,Π (i) i ,∆(k), . . . ,Π (m) m (A) In P (k), rj(P (k)′) = rj(P ) (j 6= i), ri(P (k)′) ≤ ri(P ) − 1 and in sum r(P (k)′) ≤ r(P )−1. Hence, the induction hypothesis applies and we obtain mix-free proofs P (k)′′ of Π(1)1 , . . . ,Π (i) i ,∆(k), . . . ,Π (m) m . Construct a proof P ′ as follows: .... P1 Π1 . . . .... P (1) ′′ Π (1) 1 , . . . , Π (i) i , ∆(1), . . . , Π (m) m Π (1) 1 , . . . , Π (m) m , ∆(1) . . . .... P (p) ′′ Π (1) 1 , . . . , Π (i) i , ∆(p), . . . , Π (m) m Π (1) 1 , . . . , Π (m) m , ∆(p) Π (1) 1 , . . . , Π (m) m , [i: A] J2 . . . .... Pm Πm Π (1) 1 , . . . , Π (i−1) i−1 , Π (1) 1 , . . . , Π (m) m , Π (i+1) i+1 , . . . Π (m) m (A) 3.7. analytical properties of pl 51 Note that A does not occur at place i in Π(j)j (j 6= i), since otherwise case (2) would have applied, hence ri(P ′) = 1. With rj(P ′) = rj(P ) (j 6= i) we have that r(P ′) < r(P ) and the induction hypothesis yields a mix-free proof P ′′, whence we have: .... P ′′ Π (1) 1 , . . . ,Π (i−1) i−1 ,Π (1) 1 , . . . ,Π (m) m ,Π (i+1) i+1 , . . . Π (m) m Π (1) 1 , . . . ,Π (m) m xc This completes the proof of the cut-elimination theorem. 3.7 Analytical Properties of PL The cut-elimination theorem is often invoked to show certain simple properties of sequent calculus, such as the subformula property, the decidability of propositional logic, or consistency. In fact, these properties do not depend on the cut-elimination theorem itself, i.e., on the availability of an algorithm for eliminating cuts from a given proof, but on the completeness of the cut-free fragment. However, we may want to extract information from cut-free proofs. The cut-elimination algorithm makes the constructive extraction of such information possible, given an arbitrary proof (containing cuts). The following subformula property says that a proof of a statement (i.e., a sequent) need only consist of subformulas of the formulas in the given sequent. This makes the notion precise that a sequent can be proved without "detours,"or having to use information unrelated to the end sequent. 3.7.1. Proposition (Subformula Property) A cut-free PL-proof contains only formulas which are subformulas of the end-sequent. Proof. By induction on the length of proofs: All rules except the cut rule have the property that every formula occurring in the premises also occurs in the conclusion. The midsequent theorem is what Gentzen called his Verschärfter Hauptsatz, meaning a strengthening of the cut-elimination theorem. It shows the existence of proofs which are not only cut-free, but also have all quantifier inferences below all propositional inferences. 3.7.2. Theorem (Midsequent Theorem) Let Π be a sequent consisting only of prenex formulas which is provable in PL. Then there is a cut-free proof of Π containing sequents Σ1, . . . , Σp, s.t. (1) Σj is quantifier-free, (2) every inference above Σj is either structural or propositional, and (3) every inference below Σj is either structural or a quantifier inference. Proof. By the cut-elimination theorem and Corollary 3.3.11, there is a cut-free proof P of Π from atomic axiom sequents. The order of a quantifier introduction J in P is defined as the number of propositional inferences below J . The order o(P ) of P is the sum of the orders of all quantifier inferences occurring in P . We prove the theorem by induction on the order of P : 3.8. interpolation 52 (1) o(P ) = 0: There is no propositional inference occurring below any quantifier inferences. Let B be a branch in P . If there is a propositional inference on B, then let ΣB be the conclusion of the lowermost such, otherwise let ΣB be the (atomic) axiom sequent on B. If ΣB contains a quantified formula F , then F has to be introduced by weakenings. To see this, recall that Π contains only prenex formulas: By the subformula property, ΣB is a subformula of a formula in Π, hence no propositional inferences apply to it. Eliminate F and all inferences applying to it from the part of B above ΣB , and add appropriate weakenings and exchanges directly below ΣB . The (finite) set of all ΣB serves as the set of Σi in the statement of the theorem. (2) o(P ) > 0: Then there is a quantifier inference I with the following property: The topmost logical inference below I is a propositional inference, say J . The part of P between I and J takes the following form: .... Γ1 . . . 〈 .... Γ ′,∆Q:i(k) 〉 k Γ ′, [i: (Qx)A(x)] I .... ∗ Γ ′′ . . . .... Γn Γ J where the part denoted by ∗ contains only structural inferences, and Γ and Γ ′′ contains (Qx)A(x) as a sequent-formula. We can now lower the order by exchanging the positions of I and J : / ∖ .... Γ1 ∆Q:i(k), Γ1 . . . .... Γ ′,∆Q:i(k) Γ ′,∆Q:i(k), [i: (Qx)A(x)] w ∆Q:i(k), Γ ′, [i: (Qx)A(x)] e .... ∗ ∆Q:i(k), Γ ′′ . . . .... Γn ∆Q:i(k), Γn ∆, Γ J Γ,∆Q:i(k) e ∖ / k Γ, [i: (Qx)A(x)] I Γ ec In classical logic, the quantifier rules all have only one premise each. From this it follows that there always is one midsequent. This midsequent contains, in essence, an Herbrand disjunction of the prenex formula in the end-sequent. If the many-valued logic under consideration has a similar structure, i.e., the quantifiers considered have only one premise, then an analogue of Herbrand's Theorem holds for this logic. 3.8 Interpolation Interpolation is an interesting problem for logics with something resembling an implication connective. The question asked is, for two given formulas A and B where A ⊃ B, is there a formula C, called an interpolant of A and B, s.t. (1) A ⊃ C and C ⊃ B, (2) C contains only predicate symbols common to A and B and 3.8. interpolation 53 (3) C contains only free variables common to A and B. For classical logic, the interpolation theorem was first proved by Craig [1957]. Maehara [1960] showed the interpolation theorem in a proof theoretical fashion by making use of what has become known as Maehara's Lemma (see also Takeuti [1987], p. 31ff). It also enabled Schütte [1962] to extend the theorem to intuitionistic predicate logic. Maehara's Lemma states that, if a sequent Γ1, Γ2 → ∆1,∆2 is provable, then there is an interpolant such that Γ1,→ ∆1, C and C,Γ2 → ∆2 are provable and the analogues of (2) and (3) hold. We prove a version of Maehara's Lemma in the framework of the p-sequent calculus for many-valued logic. This was first done by Miyama [1974] and more recently in a slightly more general form by Hanazawa and Takano [1985] for the formalism of Takahashi [1967a]. Gill [1970] was the first to consider interpolation in three-valued logic. The interpolation theorem for classical logic follows immediately from Maehara's Lemma since the sequent arrow and the implication are equivalent in the sense that A ⊃ B is valid iff A → B is. The theorem is obtained from the lemma simply by taking Γ1 = A, ∆2 = B, and Γ2 = ∆1 = ∅. In the many-valued case of course, no such correspondence holds in general. In fact, Maehara's Lemma is true in any many-valued logic, while the interpolation theorem need not hold. For instance, none of the Lukasiewicz logics (other than two-valued) interpolate, as has been shown by Krzystek and Zachorowski [1977]. Assume the logic under consideration contains the following truth functions (either as primitive connectives or representable): (1) For every truth value v the constant truth function Cv (2) Every unary truth function GUv,w, where GUv,w(v) = v if v ∈ V \ U and = w otherwise. (Note that these are all unary truth function taking only two truth values.) (3) Every n-ary truth function Hnu,w, where Hnv,w(v1, . . . , vn) = w if v1 = . . . = vn = w and = v otherwise. (4) Every quantifier Kv,w, where Kv,w(V ) = v if V = {v} and = w otherwise. We denote by Ci, GUv,w, H n v,w and Kv,w the corresponding formula schemata, respectively. 3.8.1. Theorem (Maehara's Lemma) Let Γ,Λ be a p-provable (and hence p-valid) sequent where Γ ∩ Λ = ∅. Then there is, for every pair v 6= w of truth values, a formula C, called an interpolant of Γ and Λ, with the following properties: (1) Γ, [v: C] and Λ, [w: C] are both provable. (2) Every predicate symbol of C occurs in both Γ and Λ. (3) Every free variable of C occurs in both Γ and Λ. Proof. The proof is by induction on the height h of the proof of Γ,Λ: h = 0: Γ,Λ is an axiom. We distinguish cases: (1) Γ = ∅: Take for C the constant formula Cv. Since [u: Cu] is p-valid, Γ, [v: Cv] is provable. Λ is provable by hypothesis, hence Λ, [w: Cv] is provable by weakening. (2) Λ = ∅: Similarly, with C = Cw. 3.8. interpolation 54 (3) Γ = [U : A] and Λ = [V \ U : A]: Take for C the formula GUv,w(A). Let I be some interpretation. By the definition of GUv,w, C takes the value w if A takes a value from U , and takes v otherwise. So when I p-satisfies Γ , i.e., does not p-satisfy Λ, then C takes the value w. Conversely, if Γ is not p-satisfied by I, then C takes the value v. From this it follows that Γ, [v: C] and Λ, [w: C] are both p-valid (and hence provable by completeness). h > 0: Γ, Λ is the conclusion of a rule of inference J . We distinguish cases according to this last rule in the proof: (1) J is a weakening at place i: Assume the weakening formula A belongs to Γ (the other case is symmetric). J is of the form: Γ ′, Λ Γ ′, [i: A], Λ w By induction hypothesis, there is an interpolant C of Γ ′ and Λ. C is also an interpolant of Γ and Λ, since Γ, [v: C] is provable from Γ ′, [v: C] by weakening, and C contains no predicate symbols or free variables of A which do not occur in Λ. (2) J is a introduction rule for 2 at place i: Assume the introduced formula 2(A1, . . . , An) belongs to Γ (the other case is symmetric). The proof ends in Γ ′,∆2:i(1), Λ . . . Γ ′,∆2:i(p), Λ Γ ′, [i: 2(A1, . . . , An)], Λ 2:i The induction hypothesis yields interpolants Cj of Γ ′,∆2:i(j) and Λ (j = 1, . . . , p). Let C equal Hpv,w(C1, . . . , Cp). C is an interpolant of Γ and Λ: First, it contains no predicate symbols or free variables not common to Γ and Λ. Furthermore, Γ, [v: C] and Λ, [w: C] are both p-valid. Let I be an interpretation. I p-satisfies all of Γ ′,∆2:i(j), [v: Cj ] (j = 1, . . . , p). This is the case if either (a) I p-satisfies Γ ′, (b) I p-satisfies all of ∆2:i(j), or (c) I p-satisfies at least one [v: Cj ]. In case (a), Γ ′, [i: 2(A1, . . . , An)], [u: C] is p-satisfied, since Γ ′ is p-satisfied. In case (b), we know that 2(A1, . . . , An) takes the value vi by the definition of partial normal forms. In case (c), C takes the value v, since not all of the arguments to Hpv,w equal w. In any case, then, we have that I p-satisfies Γ, [v: C]. On the other hand, if I does not p-satisfy Λ, it p-satisfies all of [w: Cj ] (j = 1, . . . , p), since Λ, [w: Cj ] are all p-valid. But then all the arguments to Hpv,w in C equal w, hence C takes the value w. Hence, Λ, [w: C] is also p-valid. (3) J is an introduction rule for Q at place i: We will again only treat the case where the introduced formula [i: (Qx)A(x)] belongs to Γ . J takes the form Γ ′,∆Q:i(1), Λ . . . Γ ′,∆Q:i(p), Λ Γ ′, [i: (Qx)A(x)], Λ Q:i There are interpolants Cj for the premises Γ ′,∆Q:i(j) and Λ (j = 1, . . . , p), by induction hypothesis. As in the last case, D = Hpv,w(C1, . . . , Cp) is such that Γ, [v: D] and Λ, [w: D] are p-valid. Note that the Cj and hence D do not contain eigenvariables of the inference J . To see this, recall that Cj contains only free variables common to Γ ′,∆Q:i(J) and Λ. But eigenvariables cannot have this property, since they must not appear in Λ. However, since the Cj are interpolants of the premises, they may contain free variables which occur only in terms being replaced by the bound variable x 3.8. interpolation 55 and do not occur in Γ ′. These variables occur only in Λ and no longer in Γ ′, [vi: (Qx)A(x)]. The truth or falsity of Γ, [v: D] can obviously not depend on the assignment of such a variable in a given interpretation. We can then bind all these variables by the quantifier K. To be more precise, let a1, . . . , aq be all free variables occurring in D which do not occur in Γ . Let I be an interpretation. If Γ, [v: D] is p-satisfied by I′, then every interpretation I′ ∼a1,...,aq I p-satisfies Γ, [v: D] as well. In particular, this holds for every I′′ ∼a1 I. But this means that DistrI ( D(x, a2, . . . , aq) ) = {u} where u = valI ( D(a1, . . . , aq) ) . Hence, if I p-satisfies [v: D] (i.e., u = v), then also I |=p [v: (Kx)D(x, a2, . . . , aq)]. By induction, this holds also for C = (Kx1) . . . (Kxq)D(x1, . . . , xq). In other words, Γ, [v: C] is p-valid. On the other hand, the truth of Λ, [w: D(a1, . . . , aq)] may depend on the assignment of a1, . . . , aq, since they occur in Λ. But if there is an interpretation I′ ∼a1 I s.t. I |=p Λ and I′ 6|=p Λ but I′ |=p [w: D(a1, . . . , aq)] this means that DistrI ( D(x, a2, . . . , aq) ) 6= {v}. By the definition of K it holds that valI ( (Kx)D(x, a2, . . . , aq) ) = w. By induction we have that any given interpretation I either p-satisfies Λ or [w: C], i.e., Λ, [w: C] is p-valid. In contrast to the proofs of Maehara [1960] and Miyama [1974], the above proof is not entirely proof theoretical in that the p-validity of certain sequents was shown by semantical considerations. This makes it necessary to invoke the completeness theorem. However, this is only due to the fact that the formulas/functions Cv, GUv,w, H n v,w and Kv,w are, to a certain extent, left unspecified. One could, for instance, require that they be given as primitive connectives and quantifiers in the logic, and fix certain introduction rules for them. If this is done, it seems probable that the demonstrations of p-validity of certain sequents in the proof can be replaced by giving derivation schemata of these sequents. For instance, it should be possible to give a uniform way of deriving Γ ′, [vi: 2(A1, . . . , An)], [v: C] from Γ ′,∆2:i(j), [v: Cj ], like in the proofs of Maehara and Miyama. This has the advantage of not only giving a constructive method of obtaining the interpolant (this is also accomplished by our proof), but also of obtaining proofs of Γ, [v: C] and Λ, [w: C]. Deeper investigations are also needed for exploring the connections between Maehara's Lemma and interpolation as a property of implication. For instance, what does an implication function have to look like so that Maehara's Lemma implies the interpolation property? Can one find simpler, or more natural, families of functions so that Maehara's Lemma can be proved? Chapter 4 Natural Deduction 4.1 Introduction Gentzen [1934] formulated natural deduction for intuitionistic logic as the system NJ. In correspondence with the intuitionistic sequent calculus LJ, where the right side of a sequent is restricted to at most one formula, NJ deals with inference patterns ("Schlussweisen") of one conclusion from a set of assumptions. At the application of rules, assumptions of a certain form can be discharged in parts of the proof. A proof of a formula is a deduction tree where all assumptions have been discharged. In NJ, the symbol "⊥" is used to denote falsehood, or equivalently, an empty conclusion. Gentzen gives introduction and elimination rules for the connectives and quantifiers, as well as the weakening rule for NJ: ⊥ A Natural deduction for classical logic NK is then obtained by adding tertium non datur to NJ. Alternatively, one can drop the restriction to one formula in the conclusion and allow sets of formulas. For a discussion of the resulting systems see for instance Ungar [1992], Chapters 4 and 5, or Girard [1987], § 2. Parigot [1992] has given computational meaning to the multi-conclusion natural deduction calculus for classical logic by normalization and Curry-Howard isomorphism (see also his [1993]). A general construction of sound and complete natural deduction calculi leads to an adequate syntactical (proof-theoretic) characterization of many-valued logics for which one wants to emphasize the rôle of a particular (set of) truth value(s). We generalize the classical multi-conclusion system of natural deduction to the m-valued case. This is done in essence by splitting the many-valued (positive) sequent into the part corresponding to the designated truth values and the part corresponding to the non-designated ones. The former then corresponds to facts derived, the latter to facts assumed. We give, in a systematic way, introduction rules for every connective or quantifier for every designated position, and elimination rules for every non-designated position. This system, with the set of designated truth values restricted to {vm}, has been presented in Baaz et al. [1993a]. In contrast to the many different (more or less equivalent) approaches in the literature to sequent calculi and tableaux calculi, natural deduction has not been treated in this fashion before (to the best of the author's knowledge). For the three-valued logic of Lukasiwewicz, however, natural deduction systems have been developed, e.g., the calculus of White [1980]. The approach of Becchio and Pabion [1977] is very similar to the one used here. In fact, in the same paper they first consider three-part sequents to formalize L3. This is their starting point for a natural deduction system, which uses the operator M to indicate that a formula 56 4.2. natural deduction systems 57 takes the intermediate truth value. Still, their formulation is closer to the original NJ. Natural deduction is concerned with validity, or with valid patterns of inference. Hence, we make use of the theory of positive sequent calculi only. This seems more natural, as the positive sequent calculus is more closely related to the notion of validity, in the same way as the negative system is more closely related to unsatisfiability. It is obvious how to adapt the presentation for a dual system of negative natural deduction. All the results of this chapter remain valid for the negative case. 4.2 Natural Deduction Systems Let Γ be a (set) sequent, V + ⊆ V a set of designated truth values. The set of non-designated truth values is then V − = V \ V +. We divide the sequent Γ into its designated part Γ+ and its non-designated part Γ− in the obvious way: Γ+ := 〈Γi | i = 1, . . . ,m; vi ∈ V +〉 Γ− := 〈Γi | i = 1, . . . ,m; vi ∈ V −〉 In natural deduction, we deal with inferences from certain assumptions Γ− to certain conclusions Γ+. Such an inference is written thus: Γ− Γ+ The semantics employed here are the positive semantics for sequents. More precisely, an interpretation I makes an inference pattern true, if I p-satisfies Γ . In other words, whenever all formulas in Γ− take a truth value not corresponding to the place at which they stand, then some formula in Γ+ does take a truth value corresponding to its place. In the two-valued case we have the following situation: V = {t, f}, V + = {t}. An inference schema is of the form Γ ∆ The meaning of this is that if everything in Γ is not false (i.e., true), then something in ∆ is true. Hence, to say that it means "if Γ is true, then ∆ is true" is misleading, since we are dealing with a conjunction over Γ but a disjunction over ∆. The disanalogy is even more apparent in the many-valued case. Of course, in intuitionistic natural deduction, ∆ is a singleton, so the problem only arises in the classical multiple-deduction calculus. The natural deduction calculus acts on such inferences by giving rules how some of these inferences can be combined or altered to give new inferences. This will be made precise later. First, we give the deduction rules. These are based on the introduction rules of the positive sequent calculus, modified as described in Section 3.4 to deal with set-sequents and combinational rules. 4.2.1. Definition Let f be either 2 or Q, and let F be 2(A1, . . . , An) or (Qx)A(x), respectively. Let an introduction rule for f at place i ∈ V + be given as in Definition 3.4.1. The (natural deduction) introduction rule f :Ii for f at place i ∈ V + is given by: 〈 Γ−j , d∆′f :i(j)−e Γ+j ,∆ ′ f :i(j) + 〉 j∈I Γ+, [i: F ] The formulas in d. . .e are those which can be discharged at this inference. If f is Q, then the rule has to satisfy the appropriate eigenvariable conditions: The eigenvariables in ∆′f :i(j) (j ∈ I) must not occur in Γ or in F = (Qx)A(x). 4.2. natural deduction systems 58 The partial sequents Γ+j ,∆ ′ f :i(j) are called the premises of the rule; Γ +, [i: F ] is called the conclusion. 4.2.2. Definition Let f be either 2 or Q, and let F be 2(A1, . . . , An) or (Qx)A(x), respectively. Let an introduction rule for f at place i ∈ V − be given as in Definition 3.4.1. The elimination rule f :Ei for f at place i ∈ V − is given by: Γ−0 , d[V − \ vi: F ]e Γ+0 , [V +: F ] 〈 Γ−j , d∆′f :i(j)−e Γ+j ,∆ ′ f :i(j) + 〉 j∈I Γ+0 , Γ + The formulas in square brackets are those which can be discharged at this inference. If f is Q, then the rule has to satisfy the appropriate eigenvariable conditions: The eigenvariables in ∆′f :i(j) (j ∈ I) must not occur in Γ0, Γ or in F = (Qx)A(x). The partial sequent Γ+0 , [V +: F ] is called the major premise of the rule, the partial sequents Γ+j ,∆f :i(j) are called the minor premises of the rule, and Γ + 0 Γ + is called the conclusion. 4.2.3. Example The introduction rule for ⊃ in the logic L3 is: Γ, dAe | ∆, dAe Π,B Γ ′, dAe | ∆′, dBe Π ′, B Π,Π ′, A ⊃ B The elimination rule at place p is: Γ, dA ⊃ Be | ∆ Π,A ⊃ B Γ ′′ | ∆′′, dAe Π ′, A Γ ′′ | ∆′′, dA,Be Π ′′ Γ ′′′, dAe | ∆′′′] Π ′′′, A Π,Π ′,Π ′′,Π ′′′ The elimination rule at place f is: Γ | ∆, dA ⊃ Be Π,A ⊃ B Γ ′′ | ∆′′ Π ′, A Γ ′′, dBe | ∆′′ Π ′′ Π,Π ′,Π ′′ 4.2.4. Remark Note that some of Gentzen's original rules for LK are different from the those as obtained by Definition 3.3.1, in that a rule for a given place is split into two, which together give a complete characterization of the connective. These rules can also be translated into natural deduction rules as above. 4.2.5. Definition A natural deduction system for a logic L is given by: (1) Assumptions of the form [V −: A] where A is any formula, (2) For every connective 2 and every truth value vi an introduction rule 2:Ii (if vi ∈ V +) or an elimination rule 2:Ei (if vi ∈ V −); (3) For every quantifier Q and every truth value vi an introduction rule Q:Ii (if vi ∈ V +) or an elimination rule Q:Ei (if vi ∈ V −); (4) The weakening rule for all vi ∈ V +: Γ− Γ+ Γ+, [i: A] w: i Weakenings are considered as introductions. 4.2. natural deduction systems 59 4.2.6. Definition A natural deduction derivation is defined inductively as follows: (1) Let A be any formula. Then [V −: A] [V +: A] is a derivation of A from the assumption [V −: A] (an initial derivation). (2) If Dk are derivations of Γ+k ,∆ + k from the assumptions Γ − k , ∆ − k , and〈 Γ−k , d∆ − k e Γ+k ,∆ + k 〉 k∈K Π+ is an instance of a deduction rule with ∆−k a subsequent of ∆ − k , and all eigenvariable conditions are satisfied, then 〈Dk〉k∈K Π+ is a derivation of Π+ from the assumptions ⋃ k∈K Γ − k . The formulas in ∆ − k which do not occur in ⋃ k∈K Γ − k are said to be discharged at this inference. 4.2.7. Definition We call a formula occurrence A (1) the conclusion formula of an introduction, if it is the formula being introduced, i.e., it is F in the conclusion [i: F ]; (2) a premise formula of an introduction, if it is one of the formulas in ∆′f :i(j) + in that introduction; (3) a major premise formula of an elimination, if it is among the formula being eliminated, i.e., in the major premise [V +: F ]; (4) a minor premise formula of an elimination, if it is among the formulas in ∆′f :i(j) + in that elimination, (5) a discharged assumption formula of an elimination, if it stands immediately below an assumption which contains the formulas in ∆′f :i(j) − being discharged at that elimination. A formula occurrence A is said to follow A′, if both are of the same form and A′ stands immediately above A at the same position. 4.2.8. Theorem (Soundness) If a partial sequent Γ+ can be derived from the assumptions Γ−, then the following holds for every interpretation M: If no formula in iΓ− (vi ∈ V −) evaluates to the truth value vi, then there a vj ∈ V + and a formula in Γ+j that evaluates to vj . Proof. The statement of the theorem is obviously equivalent to: If Γ+ can be derived from assumptions Γ−, then Γ is p-valid. We prove this by inductively translating every derivation D of Γ+ from Γ− to a positive sequent calculus proof of Γ : (1) D is an initial derivation: [V −: A] [V +: A] The translated proof π(D) is [V : A] which is an axiom. 4.2. natural deduction systems 60 (2) D ends in an introduction rule at place i:〈 Γ−j , d∆′f :i(j)−e.... Dj Γj ,∆ ′ f :i(j) 〉 j∈I Γ+, F f :Ii The corresponding sequent calculus proof π(D) is:〈 .... π(Dj) Γj ,∆ ′ f :i(j) 〉 j∈I Γ, [i: F ] f :i (3) D ends in an elimination rule at place i: Γ−0 , d[V − \ vi: F ]e.... D0 Γ+0 , [V +: F ] 〈 Γ−j , d∆′f :i(j)−e.... Dj Γ+j ,∆ ′ f :i(j) + 〉 j∈I Γ+0 , Γ + f :Ei Now let S be the following sequent deduction:〈 .... π(Dj) Γj ,∆ ′ f :i(j) 〉 j∈I Γ, [i: F ] f :i The corresponding sequent calculus proof π(D) is: .... π(D0) Γ0, [V \ {vi}: F ] .... S Γ, [i: F ] Γ0, Γ, [V \ {v1, vi}: F ] cut:1i .... S Γ, [i: F ] Γ0, Γ, [V \ {v1, v2, vi}: F ] cut:2i .... Γ0, Γ, [vm: F ] .... S Γ, [i: F ] Γ0, Γ cut:mi (4) D ends in a weakening at place i: Add a weakening at place i to the translated sequent calculus proof. Note that eigenvariable conditions remain satisfied in π(D). 4.2.9. Remark Translating sequent rules for two-valued logic yield natural deduction elimination rules which differ from those given by Gentzen. However, Gentzen's rules can be obtained in a systematic way by a simplification of the constructed rules. The resulting schema falls outside of our definition of natural deduction rules. We demonstrate this simplification pars pro toto for the ∀-elimination rule. The classical version as given by Parigot [1992] is: Γ ∆, (∀x)A(x) ∆, A(t) 4.3. normal derivations 61 The constructed rule is: Γ ∆, (∀x)A(x) Γ ′, dA(t)e ∆′ ∆, ∆′ Taking {A(t)} for ∆′ and ∅ for Γ ′, we obtain Parigot's rule by disregarding the redundant right premise. 4.3 Normal Derivations A maximum segment in the intuitionistic natural deduction calculus NJ is a sequence of formulas in a derivation that starts with an introduction and end with an elimination. In the classical multi-conclusion system, it is a sequence starting with an introduction of a formula and ending in an elimination acting on the same formula. A maximum segment constitutes a redundancy in the proof. In NJ, and also in multi-valued natural deduction, there are always proofs without such redundancies (see Prawitz [1971]). 4.3.1. Definition A sequence A1, . . . , Ar of occurrences of one and the same formula is called a maximum segment, if A1 is the conclusion formula of an introduction, Aj+1 stands immediately below Aj , and Ar is the the major premise formula in an elimination. 4.3.2. Definition A normal derivation is a natural deduction derivation where no major premise of an elimination stands below an introduction. 4.3.3. Proposition A normal derivation contains no maximum segments. 4.3.4. Theorem Every cut-free sequent calculus proof of a sequent Λ can be translated into a normal natural deduction derivation of Λ+ from the assumptions Λ−, where Λ− ⊆ Λ−. Proof. By induction on the height h of the proof P of Λ. h = 1: Then Λ is an axiom of the form [V : A]. The corresponding natural deduction proof δ(P ) is the initial derivation [V −: A] [V +: A] h > 1: We distinguish cases according to the last rule in P : (1) P ends in an introduction rule for f at place vi ∈ V +. Λ is of the form Γ, [i: F ]:〈 .... Pj Γj ,∆ ′ f :i(j) 〉 j∈I Γ, [i: F ] f :i Construct a natural deduction proof δ(P ) of Γ+, [i: F ] from the assumptions Γ− as follows: 〈 Γ−j , d∆′f :i(j)−e.... δ(Pj) Γ+j ,∆ ′ f :i(j) + 〉 j∈I Γ+, [i: F ] f :Ii This only adds an introduction at the end of the derivation, hence δ(P ) is normal. 4.3. normal derivations 62 (2) P ends in an introduction rule at place vi ∈ V −: Λ is of the form Γ, [i: F ]:〈 .... Pj Γj ,∆ ′ f :i(j) 〉 j∈I Γ, [i: F ] f :i Construct a natural deduction proof δ(P ) of Γ+ from the assumptions Γ−, [i: F ] as follows: d[V \ vi: F ]e, [vi: F ] [V +: F ] 〈 Γ−j , d∆′f :i(j)−e.... δ(Pj) Γ+j ,∆ ′ f :i(j) + 〉 j∈I Γm f :Ei This only adds an elimination at the beginning of a normal derivation, hence δ(P ) is normal. (3) P ends in a weakening at place vi ∈ V +: Append a weakening at vi to the natural deduction proof. (4) P ends in a weakening at place vi ∈ V −: Do nothing. Note that eigenvariable conditions remain satisfied in the translated proof. 4.3.5. Corollary (Completeness) Natural deduction systems are complete. Proof. By Theorem 3.3.10, cut-free sequent calculus is complete, hence every valid sequent Λ has a cut-free sequent calculus proof. The translation of this proof yields a natural deduction derivation of Λ+ from the assumptions Λ− (Λ− ⊆ Λ−). Every valid inference schema is of this form. 4.3.6. Corollary (Normal Form Property) For every natural deduction derivation, there exists a normal natural deduction derivation of the same set of formulas from a subset of the assumptions. Proof. If there is a derivation of Λ+ from Λ− then by Theorem 4.2.8 there is a cut-free sequent calculus proof of Λ, whose translation yields a normal derivation. 4.3.7. Definition A path in a natural deduction derivation is a sequence of occurrences of formulas A1, . . . , Ar s.t. (1) A1 is either (a) a formula standing immediately below an assumption or (b) is the conclusion formula of an introduction without premise formulas (e.g., weakenings); (2) Ar is either (a) an end formula of the derivation or (b) a minor premise formula of an elimination or (c) a major premise formula of an elimination without discharged assumption formulas, and 4.3. normal derivations 63 (3) Aj+1 (1 ≤ j ≤ r − 1) is either (a) a discharged assumption formula of an elimination rule, if Aj is the major premise formula of that elimination, or (b) the conclusion formula of an introduction if Aj is a premise formula of that rule, or (c) follows Aj . 4.3.8. Proposition A path in a normal derivation can be divided into three (possibly empty) parts: (1) The analytical part A1, . . . , Ap, where each formula is the major premise formula of an elimination and stands immediately below an assumption; Aj is a subformula of Aj−1 (2 ≤ j ≤ p). (2) The minimum part Ap+1, . . . , Aq; Aj is equal to Aj+1 (p ≤ j ≤ q). (3) The synthetical part Aq+1, . . . , Ar; Aq+1 is the conclusion formula of an introduction with premise formula Aq; Aj−1 is a subformula of Aj (q + 1 ≤ j ≤ r). 4.3.9. Remark If a cut-free sequent calculus proof with atomic axiom sequents is translated as in the proof of Theorem 4.3.4, the minimum segment in a path with non-empty analytical and synthetical part is atomic. 4.3.10. Remark It remains to be investigated for which collections of operators one can achieve strong normalization (i.e. normal form transformations with ChurchRosser property) according to some reasonable definition. (See Parigot [1992] for positive and Zucker [1974] for negative results in the two-valued case.) Chapter 5 Approximating Propositional Logics 5.1 Introduction It is clear by what has been said so far that many-valued logics are, from the perspective of proof theory, very close to classical logic. In particular, procedures for algorithmic proof search are very similar to those developed for classical two valued logic, both in the concepts used and in their complexity. Furthermore, the structural proximity to classical logic allows the translation or adaptation of many strategies and heuristics which have been developed for classical logic. This is especially true of propositional many-valued logic. The resolution method can be carried much further than in Chapter 2, see for instance Baaz and Fermüller [1992] or Baaz and Fermüller [1993]. The method of tableaux, a calculus very similar to the (negative) sequent calculus, can also be used for mechanical theorem proving in many-valued logics, see Hähnle [1993a]. This approach has in fact produced an efficient tautology checker for many-valued logics (see Hähnle et al. [1992]). Although satisfiability in many-valued propositional logics is (as in classical logic) NP-complete (see Mundici [1987]), this is still much better than several other propositional formalisms currently under investigation. Especially in the area of logics for Artificial Intelligence, the corresponding decision problems are hard or complete for classes higher up in the polynomial hierarchy (see, e.g., Eiter and Gottlob [1992]). From a practical point of view it seems natural to ask for other logics, which in a sense approximate these more complex logics to a satisfactory degree, but are not as complex themselves. The remarks above suggest that manyvalued logic might be an appropriate candidate for this task. Of course, the results in this chapter are far from being applicable for real applications. But they are, if not a first step in a promising direction, at least theoretically appealing. 5.2 Propositional Logics 5.2.1. Definition A propositional language L consists of the following: (1) propositional variables: X0, X1, X2, . . . , Xj , . . . (j ∈ ω) (2) propositional connectives of arity nj : 2n00 , 2 n1 1 , . . . , 2 nr r . If nj = 0, then 2j is called a propositional constant. (3) Auxiliary symbols: (, ), and , (comma). Formulas are defined as usual. We denote the set of formulas over a language L by Frm(L), or by simply by Frm if the language is understood. A propositional many-valued logic M is given by a set of truth values V (M) = {1, 2, . . . ,m}, the set of designated truth values V +(M), and a set of truth functions 2i: V (M)i → V (M) for all connectives 2i. 64 5.2. propositional logics 65 5.2.2. Example Examples are provided by classical propositional logic and propositional three-valued Lukasiewicz logic. These are given by the truth functions for the connectives given in Examples 1.3.2 and 1.3.3. Furthermore, we have the sequence of m-valued Gödel logics Gm given by V (Gm) = {0, 1, . . . ,m − 1}, the designated values V +(Gm) = {m− 1}, and the following truth functions: ¬Gm(v) = { m− 1 for v = 0 0 for v 6= 0 ∨Gm(v, w) = max(a, b) ∧Gm(v, w) = min(a, b) ⊃Gm(v, w) = { m− 1 for v ≤ w w for v > w This sequence of logics was used in Gödel [1932] to show that intuitionistic logic can not be characterized by a finite matrix. 5.2.3. Definition A valuation I is a mapping from the set of propositional variables into V (M). A valuation I can be extended in the standard way to a function from formulas to truth values. I satisfies a formula F , in symbols: I |=M F , if I(F ) ∈ V +(M). In that case, I is called a model of F , otherwise a countermodel. A formula F is a tautology of M iff it is satisfied by every valuation. Then we write M |= F . We denote the set of tautologies of M by Taut(M). 5.2.4. Definition A propositional logic L in the language L is given by a finite set of propositional rule schemas R(L) being a subset of Frm∗. Rules of length 1 are called axioms. A formula F is a theorem of L, if there is a derivation of F from R(L), i.e., a finite sequence F1, F2, . . . , Fn = F of formulas, s.t. for each Fi there is a rule 〈A1, . . . , An〉 ∈ R(L) where Fi is a substitution instance of An, and there are k1, . . . , kn−1 < i s.t. Fkj is a substitution instance of Aj (1 ≤ j ≤ n − 1). In that case we write L ` F . The set of theorems of L is denoted by Thm(L). 5.2.5. Remark Propositional logics as defined here are always finitely axiomatized and have the substitution property. Many well-known propositional logics are finitely axiomatizable, e.g., classical and intuitionistic propositional logic, as well as many modal logics. The assumption of finite axiomatizability is used essentially in some proofs. 5.2.6. Example Intuitionistic propositional logic H is given by the following axioms (We give the axiomatization of Heyting [1930]): a1 A ⊃ A ∧A a2 A ∧B ⊃ B ∧A a3 A ⊃ B ⊃ (A ∧B ⊃ B ∧ C) a4 (A ⊃ B) ∧ (B ⊃ C) ⊃ (A ⊃ C) a5 B ⊃ (A ⊃ B) a6 A ∧ (A ⊃ B) ⊃ B a7 A ⊃ A ∨B a8 A ∨B ⊃ B ∨A a9 (A ⊃ C) ∧ (B ⊃ C) ⊃ (A ∨B ⊃ C) a10 ¬A ⊃ A ⊃ B a11 (A ⊃ B) ∧ (A ⊃ ¬B) ⊃ ¬A and the following rules (in usual notation): A B A ∧B r1 A A ⊃ B B r2 5.3. singular approximations 66 5.3 Singular Approximations First of all, we are interested in the relationship of many-valued logics to arbitrary logics in respect to the formulas they prove. In other words, what we are interested in is the containment relation between Taut(M) and Thm(L): 5.3.1. Definition A many-valued logic M approximates a logic L iff Thm(L) ⊆ Taut(M). M is called a (singular) approximation of L. 5.3.2. Example Every m-valued logic having the set of formulas Frm(L) as its set of tautologies is a singular approximation of any logic. Less trivially, the m-valued Gödel logic Gm is an m-valued singular approximation of H. If we have an approximation of a given logic L, we have a simple test if a formula is not a theorem of L: If F is not valid in M, then F cannot be a theorem of L. Since we cannot hope for the converse to hold, unless L is a many-valued logic, we want M to have as few theorems as possible. In fact, we would like to be able to find the best m-valued approximation for a given logic L. This is often possible, since (1) all m-valued propositional logics for a language L can be enumerated mechanically, and there is only a finite number of them, (2) it is decidable of two m-valued logics given by their truth tables, which of them has fewer theorems. The additional condition which needs to be satisfied is this: It is decidable whether an m-valued logic M approximates a given propositional logic L. The naive test for this would be to check that M satisfies all axioms and rules of L. However, this test might fail even though M actually is an approximation of L, e.g., if there are redundant rules. Therefore, we restrict attention to many-valued logics where this test is positive. Such many-valued logics are called sound approximations of L. Set inclusion of the sets of theorems orders all m-valued logics for a language into a finite complete lattice. The minimal elements in the lattice restricted to (sound) approximations of a given logic L then are the optimal approximations of L. We make all this precise. 5.3.3. Proposition There are ∏r j=1 m mnj many m-valued logics. Proof. The number of different truth functions V nj → V equals ∣∣V V nj ∣∣ = mmnj . 5.3.4. Proposition An m-valued logic M approximates the logic L if (not iff) (A) for every rule r in R(L): if a valuation of the variables in r makes all the premises of r true, it also makes the conclusion true. Proof. If: Let L ` F . We show that M |= F by induction on the length l of the derivation in L: l = 1: This means F is a substitution instance of an axiom A. By hypothesis, M |= A. Let I be a valuation of the variables in F . Let X1, X2, . . . , Xn be all the variables in A, let F = A[B1/X1, . . . , Bn/Xn] and let vi = I(Bi) (1 ≤ i ≤ n). By hypothesis, M |= A and, a fortiori, {X1 7→ v1, . . . , Xn 7→ vn} |= A. But this means that I |= F . Hence, M |= F . 5.3. singular approximations 67 l > 1. F is the conclusion of a rule r ∈ R(L). If r is 〈A1, . . . , Ak, A〉, and X1, X2, . . . , Xn are all the variables in A, A1, . . . , Ak, then the inference has the form A1[B1/X1, . . . , Bn/Xn] . . . Ak[B1/X1, . . . , Bn/Xn] F = A[B1/X1, . . . , Bn/Xn] r Let I be a valuation of the variables in F , and let vi = I(Bi) (1 ≤ i ≤ n). By induction hypothesis, the premises of r are valid. This implies that, for 1 ≤ i ≤ k, we haveu {X1 7→ v1, . . . , Xn 7→ vn} |= Ai. By hypothesis then, {X1 7→ v1, . . . , Xn 7→ vn} |= A. But this means that I |= F . Hence, M |= F . 5.3.5. Definition An m-valued logic M soundly approximates a given logic L iff condition (A) of Proposition 5.3.4 holds. 5.3.6. Corollary It is decidable if a given m-valued logic M soundly approximates a given propositional logic L. Proof. There are only finitely many rules to check, and condition (A) of Proposition 5.3.4 can be checked in finite time. 5.3.7. Remark The adjective sound has been chosen because a many-valued logic M which is an approximation but not a sound approximation of a given logic L has the following paradoxical property: For M we can construct the positive sequent calculus as in Chapter 3. Every rule of L can be written as a sequent calculus rule as follows: Let 〈F1, F2, . . . , Fn〉 be the rule in question. The corresponding sequent calculus rule is: Γ, [V +: F1] . . . Γ, [V +: Fn−1] Γ, [V +: Fn] The system resulting from combining these rules with the sequent calculus for M becomes unsound when all truth value constants are added. 5.3.8. Example The m-valued Gödel logics Gm are sound approximations of H. For instance, take axiom a3: B ⊃ A ⊃ B. This is a tautology in Gm, for assume we assign some truth values a and b to A and B, respectively. We have two cases: If a ≤ b, then (A ⊃ B) takes the value m−1. Whatever b is, it certainly is ≤ m−1, hence B ⊃ A ⊃ B takes the designated value m − 1. Otherwise, A ⊃ B takes the value b, and again (since b ≤ b), B ⊃ A ⊃ B takes the value m− 1. Modus ponens (r2) also passes the test: Assume A and A ⊃ B both take the value m− 1. This means that a ≤ b. But a = m− 1, hence b = m− 1. 5.3.9. Remark In connection with the relation of approximations and sound approximations the following questions are open, which have an influence on the applicability and range of the results in this chapter: By Proposition 5.3.4 it is decidable whether a given m-valued logic soundly approximates a given logic. Is the corresponding problem decidable or undecidable for approximations in general? Are there reasonable classes of logics where it is decidable? Is there a way to determine, for a given (class of) logic(s), whether it has only sound approximations? What are the reasons for the discrepancy between approximations and sound approximations? We have the following situation: An m-valued logic M and a logic L are given. M singularly approximates L, but the test fails. This may be due to (at least) two reasons: (1) The axiom system for L contains a rule for which the test fails, but which is never actually used in a deduction in L. (2) If the test fails, 5.3. singular approximations 68 this means that there is a rule r in L and an assignment to the variables occurring in r, s.t. the premises of r are true but the conclusion is true. Now it is prima facie conceivable that those formulas which, if substituted for the variables in the premises, make those premises provable, actually never take the truth values of the counterexample. For instance, r might be F (X) ` G(X), and for an interpretation with val(X) = v we have that F (X) is true but G(X) is false. But what we would in fact want to test if whether it is possible for a theorem of L of the form F (H) to be true in M while G(H) is false. But it is conceivable that no formula H for which F (H) is a theorem of L ever takes the truth value v in M. The questions now are: Do such effects ever happen? Do they happen in "interesting" logics? Is it decidable, given a set of rules, whether such things happen? 5.3.10. Definition An m-valued logic M1 is better than M2, M1  M2, iff Taut(M1) ⊂ Taut(M2). 5.3.11. Theorem Let two m-valued logics M1, M2 be given. It is decidable whether M1  M2. Proof. It suffices to show the decidability of the following property: There is a formula A, s.t. (*) M2 |= A but M1 6|= A. We show this by giving an upper bound on the depth of a minimal formula A satisfying the above property. Since the set of formulas of L is enumerable, bounded search will produce such a formula iff it exists. Note that the property (*) is decidable by enumerating all assignments. Let A be a formula that satisfies (*), i.e., there is a valuation I s.t. I 6|=M1A. W.l.o.g. we can assume that A contains at most m different variables: if it contained more, some of them must be evaluated to the same truth value in the counterexample I for M1 6|= A. Unifying these variables leaves (*) intact. Let B = {B1, B2, . . .} be the set of all subformulas of A. Every formula Bj defines an m-valued truth function f(Bj) of m variables where the values of the variables which actually occur in Bj determine the value of f(Bj) via the matrix of M2. On the other hand, every Bj evaluates to a single truth value t(Bj) in the countermodel I. Consider the formula A′ constructed from A as follows: Let Bi be a subformula of A and Bj be a proper subformula of Bi (and hence, a proper subformula of A). If f(Bi) = f(Bj) and t(Bi) = t(Bj), replace Bi in A with Bj . A′ is shorter than A, and it still satisfies (*). By iterating this construction until no two subformulas have the desired property we obtain a formula A∗. This procedure terminates, since A′ is shorter than A; it preserves (*), since A′ remains a tautology under M2 (we replace subformulas behaving in exactly the same way under all valuations) and the countermodel I is also a countermodel for A′. The depth of A∗ is bounded above by mm m+1−1. This is seen as follows: If the depth of A∗ is d, then there is a sequence A∗ = B′0, B ′ 1, . . . , B ′ d of subformulas of A∗ where B′k is an immediate subformula of B ′ k−1. Every such B ′ k defines a truth function f(B′k) of m variables in M2 and a truth valued t(B ′ k) in M1 via I. There are mm m m-ary truth functions of m truth values. The number of distinct truth function-truth value pairs then is mm m+1. If d ≥ mmm+1, then two of the B′k, say B′i and B ′ j where B ′ j is a subformula of B ′ i define the same truth function and the same truth value. But then B′i could be replaced by B ′ j , contradicting the way A ∗ is defined. 5.3.12. Corollary It is decidable if two m-valued logics define the same set of tautologies. The non-strict relation  between m-valued logics is decidable. 5.4. sequential approximations 69 Proof. Taut(M1) = Taut(M2) iff neither M1  M2 nor M2  M1. M1  M2 iff M1  M2 or Taut(M1) = Taut(M2). Let ' be the equivalence relation on m-valued logics defined by: M1 ' M2 iff Taut(M1) = Taut(M2), and let MVLm be the set of all m-valued logics over L. By Mm we denote the set of all sets Taut(M) of tautologies of m-valued logics M . The partial order 〈Mm,⊆〉 is isomorphic to 〈MVLm/ ',〉. 5.3.13. Proposition 〈Mm,〉 is a finite complete partial order. Proof. The set of m-valued logics MVLm is obviously finite, since there are at most mn1mn2 * * *mnc different m-valued matrices for C.  is a partial order on MVLm/ ' with the smallest element ⊥ := Frm(L) and the largest element > := ∅. The "best" logic is the one without theorems, generated by a matrix where no connective takes a designated truth value anywhere. The "worst" logic is the one where every formula of L is a tautology, it is generated by a matrix where every connective takes a designated truth value everywhere. In every complete partial order over a finite set, there exist lub and glb for every two elements of the set. Hence, 〈M,4,5,⊥,>〉 is a finite complete lattice, where 4 is the lub in , and 5 is the glb in . Since  is decidable and M can be automatically generated the functions 4 and 5 are computable. 5.3.14. Proposition The optimal (i.e., minimal under ⊆) m-valued sound approximations of a logic L are computable. Proof. Consider the set A(L) of m-valued sound approximations of L. Since A(L) is finite and partially ordered by , A(L) contains minimal elements. The relation  is decidable, hence the minimal sound approximations can be computed. 5.3.15. Remark Of course, the best approximation one could wish for is a manyvalued logic M whose tautologies coincide with the theorems of L. L then provides an axiomatization of M. This of course is not always possible, at least for finitevalued logics. Lindenbaum has shown that any logic (in our sense, given by a set of rules and with substitution) can be characterized by an infinite-valued logic, see Lukasiewicz and Tarski [1930]. For a discussion of related questions see also Rescher [1969], § 24. Note that Rescher's notions of captures and adequate correspond to our singular approximation and sound approximation, respectively. 5.4 Sequential Approximations In the previous section we have shown that it is always possible to obtain the best sound m-valued approximation of a given logic, but there is no way to tell how good these approximations are. For all we know, the best approximations for a given logic L might only be those M having Frm(L) as their sets of tautologies, even though Thm(L) ⊂ Frm(L). It is hardly appropriate, then to call M an approximation of L, since it is far from being "close" to L. But how can we measure this "closeness" of many-valued logics to arbitrary logics? Clearly, a sensible measure is not easily defined by considering single many-valued logics. It is certainly more promising to take a step back and look at collections, say infinite sequences of many-valued logics. 5.4.1. Definition Let a logic L be given and let A = 〈M1, M2, M3, . . . , Mj , . . .〉 (j ∈ ω) be a sequence of many-valued logics s.t. 5.4. sequential approximations 70 (1) Mi  Mj iff i ≥ j, and (2) Mi is a singular approximation of L. A is called a sequential approximation of L iff Thm(L) = ⋂ j∈ω Taut(Mj). We say L is approximable, if there is such a sequential approximation for L. 5.4.2. Example Consider the sequence G = 〈G2, G3, G4, . . .〉 of Gödel logics and intuitionistic propositional logic H. Taut(Gi) ⊃ Thm(H), since Gi is a sound approximation of H. Furthermore, Gi+1  Gi. This has been pointed out by Gödel [1932], for a detailed proof see Gottwald [1989], Satz 3.4.1. However, it is not a sequential approximation of H: The formula (A ⊃ B) ∧ (B ⊃ A), while not a theorem of H, is a tautology of all Gi. In fact, ⋂ j≥2 Taut(Gi) is the set of tautologies of the infinite-valued Gödel logic Gא, which is axiomatized by the rules of H plus the above formula. This has been shown in Dummett [1959] (see also Gottwald [1989], § 3.4). Hence, G is a sequential approximation of Gא = H + (A ⊃ B) ∧ (B ⊃ A). Jaśkowski [1963] gave a sequential approximation of H. For this see also Rose [1958], Surma [1973], as well as Surma et al. [1975]. That H is approximable is also a consequence of Theorem 5.4.14, with the proof adapted to Kripke semantics for intuitionistic propositional logic, since H has the fmp (see Gabbay [1981], Ch. 4, Theorem 4(a)). The notion of approximabilty is of interest in itself, but moreover it tells us something about the quality of singular approximations: As has been pointed out above, there is no way to guarantee that we can come as close to a given logic with our singular approximation as we want to (if we only allow sufficiently many truth values). However, if L is approximable, then every formula false in L will also be falsified in a singular approximation of L with sufficiently many truth values. The natural question to ask is: Which logics are approximable? Here we can give a negative answer for undecidable logics, and a positive answer for (a class of) logics with the finite model property. 5.4.3. Proposition Let L be an undecidable propositional logic. Then L is not approximable. Proof. Since L is given by an axiom system, its set of theorems is r.e. Consequently, its set of non-theorems cannot be r.e. If L were approximable, there were a sequence A = 〈M1, M2, M3, . . .〉 s.t. ⋂ j≥2 Taut(Mj) = Thm(L). If N is a nontheorem of L, then there would be an index i s.t. N is false in Mi. But this would yield a semi-decision procedure for non-theorems of L: Try for each j whether N is false in Mj . If N is a non-theorem, this will be established at j = i, if not, we may go on forever. This contradicts the fact that the non-theorems are not r.e. 5.4.4. Example This shows that a result similar to that of Jaśkowski [1963] cannot be obtained for full propositional linear logic. This logic is undecidable, as has been shown by Lincoln et al. [1990] (see also Troelstra [1992], Ch. 20). For the positive result, note that a basic fact about sequential approximations is that every non-theorem of L is falsified in some Mi. Compare this to the finite model property of modal logics, which says that every non-theorem A of a modal logic L is falsified in some finite Kripke model. These countermodels can be coded into many-valued logics which are singular approximations of L but which also falsify A, too. By a product construction, this gives a singular approximation of L. The following definitions are taken from Chellas [1980]. 5.4. sequential approximations 71 5.4.5. Definition A modal logic L has as its language L the usual propositional connectives plus two unary modal operators: 2 (necessary) and 3 (possible). A Kripke model for L is a triple 〈W,R, P 〉, where (1) W is any set: the set of worlds, (2) R ⊆ W 2 is a binary relation on W : the accessibility relation, (3) P is a mapping from the propositional variables to subsets of W . A modal logic L is characterized by a class of Kripke models for L. This is called the standard semantics for modal logics (see Chellas [1980], Ch. 3). The semantics of formulas in standard models is defined as follows: 5.4.6. Definition Let L be a modal logic, KL be its characterizing class of Kripke models. Let K = 〈W,R, P 〉 ∈ KL be a Kripke model and A be a modal formula. If α ∈ W is a possible world, then we say A is true in α, α |=L A, iff the following holds: (1) A is a variable: α ∈ P (X) (2) A ≡ ¬B: not α |=L B (3) A ≡ B ∧ C: α |=L B and α |=L C (4) A ≡ B ∨ C: α |=L B or α |=L C (5) A ≡ 2B: for all β ∈ W s.t. α R β it holds that β |=L B (6) A ≡ 3B: there is a β ∈ W s.t. α R β and β |=L B We say A is true in K, K |=L A, iff for all α ∈ W we have α |=L A. A is valid in L, L |= A, iff A is true in every Kripke model K ∈ KL. By Taut(L) we denote the set of all formulas valid in L. 5.4.7. Example The modal logic S5 is characterized by the class of universal models, i.e., Kripke models where R = W 2. It can be shown that it is also characterized by the class of models where R is an equivalence relation (see Chellas [1980], Theorem 3.13). Modal logics as considered in the literature are also propositional logics in the sense of Definition 5.2.4, i.e., they can be axiomatized by a finite set of rules. If this is the case, then the valid formulas of La.k.a. theorems are r.e. 5.4.8. Example The modal logic S5 is axiomatized by the following axioms (see Chellas [1980], § 1.2): T 2A ⊃ A 5 3A ⊃ 23A K 2(A ⊃ B) ⊃ (2A ⊃ 2B) Df3 3A ≡ ¬2¬A in addition to the usual axioms for classical propositional logic, and the following rules (in usual notation): A 2A RN A A ⊃ B B MP 5.4. sequential approximations 72 A common method for proving that a modal logic is decidable is connected with the so-called finite model property: 5.4.9. Definition Let L be a modal logic characterized by K. L has the finite model property (fmp) iff for every A s.t. L 6|= A, there is a finite Kripke model K = 〈W,R, P 〉 ∈ K (i.e., W is finite), s.t. K 6|=L A. We would like to exploit the fmp of a modal logic L to construct a sequential approximation of L. First, we define a product operator on many-valued logics: 5.4.10. Definition Let M and M′ be m and m′-valued logics, respectively. Then M×M′ is the mm′-valued logic where V (M×M′) = V (M)×V (M′), V +(M×M′) = V +(M) × V (M′), and truth functions are defined component-wise. I.e., if 2 is an n-ary connective, then 2M×M′(w1, . . . , wn) = 〈2M, 2M′〉. For convenience, we define the following: Let I and I′ be valuations of M and M′, respectively. I×I′ is the valuation of M×M′ defined by: (I×I′)(X) = 〈I(X), I′(X)〉. If I× is a valuation of M×M′, then the valuations π1I× and π2I× of M and M′, respectively, are defined by π1I×(X) = v and π2I×(X) = v′ iff I×(X) = 〈v, v′〉. 5.4.11. Lemma Taut(M×M′) = Taut(M) ∩ Taut(M′) Proof. Let A be a tautology of M × M′ and I and I′ be valuations of M and M′, respectively. Since I × I′ |=M×M′ A, we have I |=M A and I′ |=M′ A by the definition of ×. Conversely, let A be a tautology of both M and M′, and let I× be a valuation of M × M′. Since π2I× |=M A and π2I× |=M′ A, it follows that I× |=M×M′ A. The definition and lemma are easily generalized to the case of finite products∏ i Mi by induction. 5.4.12. Definition Let K = 〈W,R, P 〉 be a finite Kripke model. We define the many-valued logic MK as follows: (1) V (MK) = {0, 1}W , the set of 0-1-sequences with indices from W . (2) V +(MK) = {1}W , the singleton of the sequence constantly equal to 1. (3) ¬MK , ∨MK , ∧MK , ⊃MK are defined componentwise from the classical truth functions (4) 2MK is defined as follows: 2MK (〈wα〉α∈W )β = { 1 if for all γ s.t. β R γ, wγ = 1 0 otherwise (5) 3MK is defined as follows: 3MK (〈wα〉α∈W )β = { 1 if there is a γ s.t. β R γ and wγ = 1 0 otherwise Furthermore, IK is the valuation defined by IK(X)α = 1 iff α ∈ P (X) and = 0 otherwise. 5.4.13. Lemma Let L and K be as in Definition 5.4.12. Then the following hold: (1) Every valid formula of L is a tautology of MK . 5.4. sequential approximations 73 (2) If K 6|=L A then IK 6|=MK A. Proof. Let B be a modal formula, and K ′ = 〈W,R, P ′〉. We prove by induction that valIK′ (B)α = 1 iff K ′ |=L B: B is a variable: P ′(B) = W iff IK(B)α = 1 for all α ∈ W by definition of IK . B ≡ ¬C: By the definition of ¬MK , valIK (B)α = 1 iff valIK (C)α = 0. By induction hypothesis, this is the case iff α 6|=L C. This in turn is equivalent to α |=K B. Similarly if B is of the form C ∧D, C ∨D, and C ⊃ D. B ≡ 2C: valIK (B)α = 1 iff for all β with α R β we have valIK (C)β = 1. By induction hypothesis this is equivalent to β |=L C. But by the definition of 2 this obtains iff α |=L B. Similarly for 3. (1) Every valuation I of MK defines a function PI via PI(X) = {α | I(X)α = 1}. Obviously, I = IPI . If L |= B, then 〈W,R, PI〉 |=L B. By the preceding argument then valI(B)α = 1 for all α ∈ W . Hence, B takes the designated value under every valuation. (2) A is not true in K. This is the case only if there is a world α at which it is not true. Consequently, valIK (A)α = 0 and A takes a non-designated truth value under IK . 5.4.14. Theorem Let L be a modal logic with the fmp, and 〈A1, A2, . . .〉 an enumeration of its non-theorems. A sequential approximation of L is given by 〈M1, M2, . . .〉 where M1 = MK1 , and Mi+1 = Mi ×MKi+1 where Ki is the smallest finite model s.t. Ki 6|=L Ai Proof. (1) Taut(Mi) ⊇ Taut(L): By inducton on i: For i = 1 this is Lemma 5.4.13 (1). For i > 1 the statement follows from Lemma 5.4.11, since Taut(Mi−1) ⊇ Taut(L) by induction hypothesis, and Taut(MKi) ⊇ Taut(L) again by Lemma 5.4.13 (1). (2) Mi  Mi+1 from A ∩B ⊆ A and Lemma 5.4.11. (3) Taut(L) = ⋂ i≥1 Taut(Mi). The ⊆-direction follows immediately from (1). Furthermore, by Lemma 5.4.13 (2), no non-tautology of L can be a member of all Taut(Mi), whence ⊇ holds. 5.4.15. Remark Note that Theorem 5.4.14 does not hold in general if L is not finitely axiomatizable. This follows from Proposition 5.4.3 and the existence of an undecidable recursively axiomatizable modal logic which has the fmp (see Urquhart [1981]). 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